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            pdftitle={DTA Scheduler: Formal Analysis},
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\hypersetup{
pdftitle={Scheduling with Bounded Information: Correctness, Progress, and Near-Optimal Performance in DTA},
pdfsubject={CoRR, stat, math, math-ph, physics.app-ph, physics.comp-ph},
pdfauthor={Xinyu Yang},
pdfkeywords={Scheduler, Bounded Information, Distributed Systems, Mathematical Analysis},
}

\title{\textbf{Scheduling with Bounded Information}\\[6pt]
\large Correctness, Progress, and Near-Optimal Performance in DTA}
\author{Xinyu Yang \orcidlink{0009-0007-2600-0948} \footnote{Contact Author: \href{mailto:Pana.Yang@hotmail.com}{Pana.Yang@hotmail.com}; ORCID: \href{https://orcid.org/0009-0007-2600-0948}{0009-0007-2600-0948}}}
\date{\today}

\begin{document}
\maketitle

\begin{center}
\textbf{DOI: \href{https://doi.org/10.5281/zenodo.20998767}{https://doi.org/10.5281/zenodo.20998767}} \\

\vspace{0.3em}

\textbf{Version History} \\
      \small v1: June 28, 2026 \quad v2: June 29, 2026 \\
      \small v3: June 30, 2026 \quad v4: June 30, 2026 \\
      \small v5: July 1, 2026 \quad v6: July 1, 2026 \\
      \small v7: July 2, 2026 \quad v8: July 2, 2026 \\
      \small v9: July 2, 2026
\end{center}

\vspace{0.6em}

\begin{abstract}
Every practical distributed scheduler makes placement decisions with an
incomplete, purely local view of system state: no worker can afford to
observe every peer's queue before deciding where to route a task, so
observed information is inherently bounded and stale.  This paper studies
how much can be guaranteed --- correctness, progress, and near-optimal
makespan --- from a scheduler that deliberately restricts itself to such
bounded information, and quantifies the price of that restriction against
schedulers (e.g.\ work-stealing) that acquire more information per
decision.  We use the \textsc{DTA} distributed task scheduler ($N$ workers,
SPSC mailboxes, Vyukov MPMC warehouse, hop-bounded deflection) as the
concrete vehicle for this study: DTA routes tasks using only local queue
occupancy and a bounded-hop deflection rule, never a global or
neighbourhood queue scan.  Using a second-quantisation Fock-space algebra
and kinetic theory, we establish across seven parts:
\textbf{(I)}~four correctness theorems (task conservation, deadlock and
livelock freedom, linearizability~\cite{herlihy_wing1990,herlihy_shavit2008});
\textbf{(II)}~bounded waiting time with an explicit, distribution-free ceiling
$W_{\max}(N,L,C_W,K,\delta)$
and $k$-bounded fairness;
\textbf{(III)}~quantitative performance bounds including mean-field load
balance ($O(1/\sqrt{N})$ imbalance), BBGKY Level-1 pair-correlator
correction ($O(1/N)$), a two-line makespan bound (deterministic
$+$ Gumbel fluctuation), a four-level stability ladder, and fork-join
DAG bounds via a sequential-nucleation kinetic equation;
\textbf{(IV)}~competitive analysis with an adversarial BoT bound
$\rho_c^{\mathrm{DTA,worst}} \leq 1 + N\lfloor N/2\rfloor\delta_{\mathrm{hop}}/W \to 1$;
\textbf{(V)}~empirical validation against public Criterion~\cite{criterion} CI benchmarks,
with variance and moment analysis confirming all predictions;
\textbf{(VI)}~an information-theoretic cost framework that decomposes
the achieved-to-oracle makespan ratio into an information gap
$\eta_{\mathrm{info}}$ and a utilisation ratio $\eta_{\mathrm{util}}$,
proves both tend to $1$ for DTA as total work $W \to \infty$ (so DTA is
heavy-traffic information-optimal, matching the fully-informed offline
oracle to leading order), and shows that at any fixed, moderate system
scale DTA's information-\emph{acquisition} cost is $\Omega(\log N)$
lower than work-stealing's on NUMA topologies, with a concrete crossover
analysis for dual-socket systems.
Taken together, these results show that bounded-information scheduling is
not merely a practical compromise but can be made \emph{provably
near-optimal}: DTA loses essentially nothing to schedulers with a wider
information window in the heavy-traffic limit, while paying a strictly
lower and asymptotically smaller acquisition cost at any fixed scale.
\end{abstract}

\newpage

\tableofcontents
\clearpage

% ========================================================================
\section*{Introduction}
\addcontentsline{toc}{section}{Introduction}
\label{sec:intro}
% ========================================================================

\subsection*{Scheduling under bounded information}

A distributed task scheduler is, at its core, an algorithm that answers a
single repeated question --- \emph{where should this task go?} --- using
whatever information it has on hand at decision time.  The classical
work-stealing scheduler~\cite{blumofe1999,arora2001,tokio} answers this
question by letting an idle worker actively \emph{search}: it probes one
or more peer queues (at random, or in round-robin order) until it finds
work.  This is a reasonable design, but it is not free: every probe is a
cross-core or cross-socket memory access, every probe that returns nothing
is wasted latency, and on NUMA hardware the cost of a single remote probe
can be several times that of a local one. A scheduler that wants to reduce
this overhead has essentially one lever available: restrict how much
state it looks at before making a decision, and rely on a cheaper, more
local mechanism to correct for the information it does not have.

This is precisely the design of the \textsc{DTA} scheduler studied in
this paper. Each worker decides where to place a task using only its own
local queue occupancy; when local capacity is exceeded, tasks are pushed
through single-producer/single-consumer (SPSC) mailboxes to a bounded
number of peers, deflected onward for at most $\lfloor N/2 \rfloor$ hops,
and only spill into a shared multi-producer/multi-consumer (MPMC)
``warehouse'' as a last resort. No worker ever scans a peer's queue length
before routing, and no worker maintains a global view of system load. In
the language of online algorithms and competitive analysis, DTA operates
under a strictly smaller \emph{information set} than a scheduler that
polls peers before acting.

The central question this paper answers is whether that restriction is a
free design choice or a real cost, and if it is a cost, how large. Answering
it requires three separate arguments, each building on the last. First, a
scheduler that observes less must still be shown \emph{safe}: with a
smaller information set there is more room for a task to be lost, for two
workers to deadlock waiting on each other's mailboxes, or for a chunk to
loop forever between peers instead of being scheduled (Parts~I--II).
Safety alone is not enough, however, since a scheduler can be safe and
still make every worker wait arbitrarily long for work; we therefore prove
an explicit, distribution-free bound on waiting time before any of the
performance analysis can be trusted (Part~III). Only once correctness and
progress are established does it make sense to ask how \emph{good} the
restricted-information design is quantitatively, and we answer this both
under a stochastic workload model (Part~IV, mean-field and kinetic theory)
and under a worst-case adversarial one (Part~V, competitive analysis and
empirical validation) --- two regimes that make no assumptions in common,
so agreement between them is itself evidence that the bound is not an
artefact of either model. Finally, we close the loop: Parts~I--V establish
that DTA \emph{works} despite restricted information, but not yet what
that restriction is worth, so Part~VI gives an explicit accounting of how
much information a comparable work-stealing scheduler must acquire to
match it, together with the hardware cost of acquiring that information.
Figure~\ref{fig:logic-chain} summarises how these parts connect to the
central question, and the ``Organization'' paragraph below gives the
section-level map. The overall picture that emerges is that bounded-information
scheduling is not a compromise made for the sake of simplicity; on the
workloads and hardware regimes we characterise, it is close to
information-optimal, and it is strictly cheaper to run.

\begin{figure}[htbp]
\centering
\resizebox{\textwidth}{!}{%
\begin{tikzpicture}[
  node distance=7mm and 6mm,
  every node/.style={font=\small},
  box/.style={draw, rounded corners, align=center, minimum height=9mm,
              minimum width=26mm, fill=blue!6, inner sep=3pt},
  qbox/.style={draw, rounded corners, align=center, minimum height=9mm,
               minimum width=26mm, fill=orange!10, inner sep=3pt},
  cbox/.style={draw, rounded corners, align=center, minimum height=11mm,
               minimum width=52mm, fill=green!8, inner sep=3pt},
  arr/.style={-{Latex[length=2mm]}, thick}
]
  \node[cbox] (q) {Central question: is bounded-information routing a free\\
                    design choice, or does it cost something -- and how much?};
  \node[box, below left=9mm and 6mm of q] (safety) {Part I--II\\Safety: no lost tasks,\\no deadlock/livelock};
  \node[box, below=of safety] (progress) {Part III\\Progress: explicit\\waiting-time bound $W_{\max}$};
  \node[box, below right=9mm and 6mm of q] (stoch) {Part IV\\Stochastic quality:\\mean-field / kinetic theory};
  \node[box, below=of stoch] (advers) {Part V\\Worst-case quality:\\competitive ratio + benchmarks};
  \node[qbox, below=16mm of q] (cost) {Part VI\\Price of the restriction:\\information-cost framework ($\eta_{\mathrm{info}},\eta_{\mathrm{util}}$)};

  \draw[arr] (q.west) .. controls +(-8mm,-4mm) .. (safety.north);
  \draw[arr] (safety) -- (progress);
  \draw[arr] (q.east) .. controls +(8mm,-4mm) .. (stoch.north);
  \draw[arr] (stoch) -- (advers);
  \draw[arr] (progress.south) .. controls +(0,-10mm) and +(-16mm,6mm) .. (cost.west);
  \draw[arr] (advers.south) .. controls +(0,-10mm) and +(16mm,6mm) .. (cost.east);
  \node[align=center, below=2mm of cost, font=\footnotesize\itshape, text width=90mm]
    {Answer: correctness and progress hold unconditionally (I--III);
     quality is near-optimal in both regimes (IV--V);
     the restriction is quantifiably cheap, not merely tolerable (VI).};
\end{tikzpicture}%
}
\caption{Logic chain of the paper. Safety (Parts~I--II) and progress
(Part~III) must hold before any performance claim is meaningful; quality is
then established independently under a stochastic workload model
(Part~IV) and an adversarial one (Part~V); Part~VI closes the loop by
pricing the information restriction whose safety and quality were just
established. Each arrow is a logical dependency, not merely a reading
order.}
\label{fig:logic-chain}
\end{figure}

\subsection*{Why a scheduler needs both a correctness proof and a cost model}

Two questions are conflated in most treatments of scheduler design, and we
deliberately keep them separate throughout this paper. The first is
whether the scheduler is \emph{correct}: does it ever lose a task, ever
deadlock, ever livelock, and are its concurrent data structures
linearizable? These are safety and liveness properties in the classical
sense, and we prove all four for DTA (Parts~I--II) using an
operator-algebraic model in which task conservation becomes a commutation
identity between a particle-number operator and the scheduler's
transition operators --- a formalism chosen because it makes the
conservation argument independent of the number of in-flight tasks, unlike
state-space model checking.

The second question is whether the scheduler is \emph{good}: does it
balance load well, does it finish batches quickly, and how does its
performance compare with alternative designs under the same workload?
This is a quantitative question, and answering it for a scheduler that
observes only local state requires tools beyond a single expectation
calculation. We use two complementary techniques. Mean-field and
kinetic-theory (BBGKY) analysis (Part~III) treats the $N$ workers as a
system of weakly interacting agents and derives the load-imbalance scaling
$O(1/\sqrt{N})$ together with its first finite-$N$ correction, in the same
spirit as the propagation-of-chaos arguments used for supermarket-model
load balancers~\cite{mitzenmacher2001,vvedenskaya1996}. Competitive
analysis (Part~IV) instead makes no probabilistic assumption on the
\emph{task sequence} and
bounds DTA's worst-case makespan directly against an offline optimum,
in the tradition of online algorithm analysis~\cite{borodin_elyaniv1998}.
Neither tool alone answers ``is bounded information worth it''; used
together with the information-cost framework of Part~VI, they do.

\subsection*{Contributions}

\begin{enumerate}
  \item An operator-algebraic (Fock-space) formal model of a
        bounded-information distributed scheduler, together with four
        correctness theorems (task conservation, deadlock freedom,
        livelock freedom, linearizability) that hold independently of the
        number of in-flight tasks.
  \item A bounded-waiting-time theorem and $k$-bounded fairness result
        that quantify how quickly a task is scheduled when routing
        decisions use only local and single-hop information.
  \item A mean-field and BBGKY kinetic-theory analysis showing that
        DTA's local-information routing achieves $O(1/\sqrt{N})$ load
        imbalance, matching the accuracy of schedulers that consult wider
        state, up to an explicit $O(1/N)$ correction; a two-line makespan
        bound; a four-level stability ladder; and a kinetic treatment of
        fork-join DAG scheduling.
  \item An adversarial, assumption-free competitive-ratio bound for
        bounded-information routing under Bag-of-Tasks workloads.
  \item An information-theoretic cost framework that formally separates a
        scheduler's competitive ratio into an \emph{information gap}
        (the structural cost of observing less) and a \emph{utilisation
        ratio} (how efficiently the observed information is used), proves
        both factors tend to $1$ for DTA in the heavy-traffic limit (so
        DTA matches the fully-informed oracle to leading order), and
        derives a concrete, hardware-parametrised crossover threshold at
        which DTA's lower acquisition cost outweighs work-stealing's
        wider information window at any fixed, moderate system scale.
  \item An empirical validation of all quantitative predictions against
        public continuous-integration benchmarks, including a variance
        and moment analysis that accounts for measurement noise.
\end{enumerate}

\subsection*{Organization}
\label{sec:organization}

Part~I (\S\ref{chap:formal_model}) develops the formal operator-algebraic
model. Part~II proves the four correctness theorems. Part~III proves
bounded waiting time and starvation freedom. Part~IV gives the
quantitative performance analysis (mean-field, kinetic theory, makespan,
stability, DAG scheduling). Part~V gives the competitive-ratio and
empirical-validation results. Part~VI develops the information-theoretic
cost framework that ties the correctness and performance results back to
the paper's central question. We close with a discussion of the scope of
applicability of the methods and the DTA algorithm, and a Method section
(\S\ref{sec:ai-method}) documenting the role of AI assistance in
preparing this manuscript's figures, tables, and the accompanying
\dtact{} implementation.

\partbreak
\part{Formal Framework}
\clearpage

\bigskip
\noindent\textbf{Overview.}
This part establishes the abstract mathematical model of the DTA scheduler.
We introduce a second-quantisation operator algebra that encodes task creation,
routing, and service as algebraic operators on a Fock-space state.
Each site in the model corresponds to one worker's locally-observable
state (its own queue and its inbound mailboxes), which keeps the
bounded-information restriction discussed in the Introduction
(\S\ref{sec:intro}) explicit at the level of the formalism itself: no
operator in this algebra reads state belonging to a site other than the
one it acts on or transfers to.
The model is deliberately code-agnostic; a correspondence table in
Chapter~\ref{chap:formal_model} maps every mathematical object to its
Rust identifier.  All correctness and performance proofs in subsequent
parts are built on the definitions given here.


% ========================================================================
% Formal Mathematical Model
% ========================================================================

\section*{Formal Mathematical Model}
\addcontentsline{toc}{section}{Formal Mathematical Model}
\label{chap:formal_model}
\setcounter{section}{0}


% =============================================================
\section{Scope and Philosophy}
% =============================================================

This chapter establishes the \emph{abstract mathematical model} of
the DTA distributed task scheduler.  It is deliberately
code-agnostic: the model is defined in terms of sets, operators, and
state-transition rules.  A correspondence table (Section~\ref{sec:corr})
then maps every mathematical object back to the specific Rust
identifier in \code{src/dta\_scheduler.rs} from which it is derived,
so that proofs in the sub-modules can cite both the mathematical
statement and the concrete code artifact that instantiates it~\cite{dtact}.

We adopt a notation inspired by \emph{second quantization}~\cite{dirac1927,jordan_wigner1928,fock1932,landau_qm,landau_fields,landau_statphys,feynman_ii,feynman_iii,derrida2007,schutz2001} from
quantum field theory.  This choice is motivated by the observation
that the scheduler's correctness rests on a \emph{conservation law}
(tasks are neither created nor destroyed by internal routing) that is
most naturally expressed as a commutation relation between the
particle-number operator and the transition operators --- the same
structure that underlies conserved charges in field theory.  We note
explicitly that the analogy is formal and classical: tasks are
\emph{distinguishable} particles, the state space is finite and
discrete, and no quantum superposition or Hilbert-space structure is
assumed or required.  The operator algebra is used purely as a
bookkeeping device that makes conservation statements compact and
unambiguous.

% =============================================================
\section{Primitive Sets}
% =============================================================

\begin{definition}[Task index set]
\label{def:taskset}
Let $\mathcal{T} \coloneqq \{0, 1, \ldots, P-1\}$ be the finite set
of \emph{task indices}, where $P \in \mathbb{N}_{>0}$ is the
context-pool capacity, fixed at runtime initialisation.  A task index
$\tau \in \mathcal{T}$ uniquely identifies a pre-allocated fiber
context; it corresponds to \code{TaskIndex = u32} and serves as the
elementary ``particle'' of the model.
\end{definition}

\begin{definition}[Chunk]
\label{def:chunk}
A \emph{chunk} is a triple
\[
  c \;=\; (\bm{\tau},\, k,\, h)
  \qquad
  \bm{\tau} \in \mathcal{T}^{K},\;
  k \in \{1,\ldots,K\},\;
  h \in \{0,\ldots,h_{\max}\},
\]
where $K = 32$ is the fixed chunk capacity (\code{CHUNK\_SIZE}),
$k$ is the number of active task indices in $\bm{\tau}$, and $h$ is
the \emph{hop count} --- the number of times the chunk has been
re-routed after a failed mailbox push.  The set of all chunks is
$\mathcal{C} \coloneqq \mathcal{T}^K \times \{1,\ldots,K\} \times
\{0,\ldots,h_{\max}\}$.
\end{definition}

\begin{definition}[Worker index set]
\label{def:workers}
Let $\mathcal{W} \coloneqq \{0, 1, \ldots, N-1\}$ with
$N \in \mathbb{N}_{\geq 1}$ be the set of \emph{worker indices}.
The maximum hop count is $h_{\max} \coloneqq \floor{N/2}$.
\end{definition}

% =============================================================
\section{Site Algebra and the Fock-Space State}
% =============================================================

\begin{definition}[Sites]
\label{def:sites}
Define the set of \emph{sites} as
\[
  \mathcal{S} \;\coloneqq\;
  \underbrace{\{Q_i \mid i \in \mathcal{W}\}}_{\text{local queues}}
  \;\cup\;
  \underbrace{\{M_{ij} \mid i,j \in \mathcal{W},\, i \neq j\}}_{\text{mailboxes}}
  \;\cup\;
  \underbrace{\{W\}}_{\text{warehouse}}
  \;\cup\;
  \underbrace{\{E_i \mid i \in \mathcal{W}\}}_{\text{executing}}
  \;\cup\;
  \underbrace{\{\varnothing\}}_{\text{completed/freed}},
\]
with site capacities
\[
  \operatorname{cap}(Q_i) = L,\quad
  \operatorname{cap}(M_{ij}) = C_M,\quad
  \operatorname{cap}(W) = C_W \cdot K,\quad
  \operatorname{cap}(E_i) = 1,\quad
  \operatorname{cap}(\varnothing) = \infty.
\]
The capacity of $E_i$ is $1$ because a worker executes at most one
task at a time.  $\varnothing$ is the \emph{sink} site representing
tasks that have finished execution and whose context slot has been
returned to the pool; it is absorbing.
\end{definition}

\begin{notation}
We write $|\alpha|_\sigma$ for the number of task indices residing at
site $\alpha$ in system state $\sigma$.  When $\sigma$ is clear from
context we write $|\alpha|$.
\end{notation}

\begin{definition}[Occupation-number state]
\label{def:state}
A \emph{system state} is a function
$\sigma : \mathcal{S} \to \bigsqcup_{\alpha \in \mathcal{S}} 2^{\mathcal{T}}$
such that
\begin{enumerate}
  \item $\sigma(\alpha) \subseteq \mathcal{T}$ for each $\alpha$,
  \item $|\sigma(\alpha)| \leq \operatorname{cap}(\alpha)$ for each $\alpha$,
  \item $\sigma(\alpha) \cap \sigma(\beta) = \varnothing$ for $\alpha \neq \beta$
        (each task index resides in \emph{exactly one} site).
\end{enumerate}
The set of all valid states is $\Sigma$.

\medskip\noindent
\textbf{Occupation-number representation}~\cite{fock1932,landau_qm}.
For each site $\alpha$ and task $\tau$ define the occupation number
$n_{\alpha,\tau}(\sigma) \coloneqq \mathbf{1}[\tau \in \sigma(\alpha)]
\in \{0,1\}$.
The state is then equivalently characterised by the matrix
$(n_{\alpha,\tau}) \in \{0,1\}^{|\mathcal{S}|\times P}$ subject to
the column-sum constraint $\sum_{\alpha} n_{\alpha,\tau} \leq 1$ for
all $\tau$ (each task is in at most one site; the remainder have not
yet been spawned).
\end{definition}

% =============================================================
\section{Creation, Annihilation, and Number Operators}
% =============================================================

\begin{definition}[Creation and annihilation operators]
\label{def:cre-ann}
For each site $\alpha \in \mathcal{S}$ and task index $\tau \in
\mathcal{T}$, define the \emph{creation operator}
$\cre{\alpha,\tau} : \Sigma \to \Sigma \cup \{\bot\}$ and the
\emph{annihilation operator} $\ann{\alpha,\tau} : \Sigma \to \Sigma
\cup \{\bot\}$ by
\[
  \cre{\alpha,\tau}(\sigma) \;\coloneqq\;
  \begin{cases}
    \sigma' & \text{if } \tau \notin \sigma(\alpha),\;
              |\sigma(\alpha)| < \operatorname{cap}(\alpha),\\
    \bot & \text{otherwise (capacity exceeded or already present)},
  \end{cases}
\]
where $\sigma'$ is identical to $\sigma$ except $\sigma'(\alpha) =
\sigma(\alpha) \cup \{\tau\}$; and
\[
  \ann{\alpha,\tau}(\sigma) \;\coloneqq\;
  \begin{cases}
    \sigma'' & \text{if } \tau \in \sigma(\alpha),\\
    \bot & \text{otherwise},
  \end{cases}
\]
where $\sigma''(\alpha) = \sigma(\alpha) \setminus \{\tau\}$ and all
other sites unchanged.

The \emph{single-site number operator} is
\[
  \num{\alpha}(\sigma) \;\coloneqq\; |\sigma(\alpha)|
  \;=\; \sum_{\tau \in \mathcal{T}} n_{\alpha,\tau}(\sigma),
\]
and the \emph{total particle-number operator} is
\[
  \Num(\sigma) \;\coloneqq\;
  \sum_{\alpha \in \mathcal{S} \setminus \{\varnothing\}} \num{\alpha}(\sigma).
\]
The sum excludes $\varnothing$ so that $\Num$ counts only
\emph{live} tasks.
\end{definition}

\begin{remark}[Classical vs.\ quantum]
In quantum mechanics~\cite{landau_qm,feynman_iii}, $\hat{a}^\dagger$ and $\hat{a}$ are linear
operators on a Hilbert space obeying canonical (anti)commutation
relations.  Here they are \emph{partial functions on a finite set};
the ``algebra'' they satisfy is simply
$\ann{\alpha,\tau} \circ \cre{\alpha,\tau}(\sigma) = \sigma$ when
$\tau \notin \sigma(\alpha)$ and capacity is available, and
$\cre{\alpha,\tau} \circ \ann{\alpha,\tau}(\sigma) = \sigma$ when
$\tau \in \sigma(\alpha)$.  No commutativity or anti-commutativity
relation between different sites is assumed; tasks are
distinguishable and the formalism is purely classical.
\end{remark}

% =============================================================
\section{Transition Operators}
% =============================================================

Every scheduler operation is a \emph{transfer} that simultaneously
annihilates a task from one site and creates it at another.

\begin{definition}[Atomic transfer operator]
\label{def:transfer}
For sites $\alpha, \beta \in \mathcal{S}$ and task $\tau \in
\mathcal{T}$, the \emph{atomic transfer operator} is
\[
  \Trf{\alpha}{\beta}(\tau) \;\coloneqq\;
  \cre{\beta,\tau} \circ \ann{\alpha,\tau},
\]
with the convention that if either component returns $\bot$, the
whole operator returns $\bot$ and the state is unchanged (the
scheduler operation fails --- returning \code{Err} or \code{None}
at the call site).
\end{definition}

The concrete scheduler transitions are then the following operators
acting on $\Sigma$.  We list them with their hop-count side-effects
and the corresponding Rust call site.

\begin{definition}[Scheduler transition operators]
\label{def:ops}
Let $i, j \in \mathcal{W}$ with $i \neq j$.
\begin{align}
  \hat{O}_{\text{push-local}}(i, \tau) &\;\coloneqq\;
    \Trf{M_{ji}}{Q_i}(\tau)
    &&\text{(\code{push\_local} after \code{pop} from } M_{ji}\text{)}
    \label{eq:push-local}\\
  \hat{O}_{\text{dispatch}}(i, \tau) &\;\coloneqq\;
    \Trf{Q_i}{E_i}(\tau)
    &&\text{(\code{dispatch\_loop} pops and executes)}
    \label{eq:dispatch}\\
  \hat{O}_{\text{finish}}(i, \tau) &\;\coloneqq\;
    \Trf{E_i}{\varnothing}(\tau)
    &&\text{(\code{free\_context} on Finished/Panicked)}
    \label{eq:finish}\\
  \hat{O}_{\text{requeue}}(i, \tau) &\;\coloneqq\;
    \Trf{E_i}{Q_i}(\tau)
    &&\text{(Notified re-enqueue in \code{dispatch\_loop})}
    \label{eq:requeue}\\
  \hat{O}_{\text{mailbox}}(i,j,\tau) &\;\coloneqq\;
    \Trf{Q_i}{M_{ij}}(\tau)
    &&\text{(\code{push\_chunk\_with\_hop}, first attempt)}
    \label{eq:mailbox}\\
  \hat{O}_{\text{hop}}(i,j,j',\tau) &\;\coloneqq\;
    \Trf{M_{ij}}{M_{ij'}}(\tau),\quad h \mapsto h+1
    &&\text{(\code{route\_deflect}, hop\_count incremented)}
    \label{eq:hop}\\
  \hat{O}_{\text{park}}(i,j,\tau) &\;\coloneqq\;
    \Trf{M_{ij}}{W}(\tau)
    &&\text{(\code{park\_in\_warehouse} or \code{route\_park})}
    \label{eq:park}\\
  \hat{O}_{\text{drain}}(i,\tau) &\;\coloneqq\;
    \Trf{W}{Q_i}(\tau)
    &&\text{(\code{drain\_warehouse} \(\to\) \code{push\_batch})}
    \label{eq:drain}\\
  \hat{O}_{\text{ext}}(i,\tau) &\;\coloneqq\;
    \Trf{\text{ext}}{M_{\text{ext},i}}(\tau)
    &&\text{(external host thread inject via \code{external\_mailbox})}
    \label{eq:ext}
\end{align}
Here \(\text{ext}\) denotes the implicit external source outside
$\mathcal{S}$; a task entering via $\hat{O}_{\text{ext}}$ is the
moment it \emph{enters} $\Num$'s domain and begins to be tracked.
\end{definition}

\begin{notation}[Composed transition]
A finite sequence of atomic transfers applied in order is written as
a left-to-right composition:
$\hat{O}_1 \cdot \hat{O}_2 \cdots \hat{O}_k$.
\end{notation}

% =============================================================
\section{The Transition Graph and the System FSM}
% =============================================================

\begin{definition}[Per-task FSM]
\label{def:fsm-task}
For a fixed task index $\tau$, the \emph{per-task finite state
machine} $\mathcal{F}_\tau$ has
\begin{itemize}
  \item \textbf{States}: the elements of $\mathcal{S}$ (the site the task currently occupies),
  \item \textbf{Initial state}: $\text{ext}$ (task not yet in the system),
  \item \textbf{Accepting/absorbing state}: $\varnothing$,
  \item \textbf{Transitions}: the operators of Definition~\ref{def:ops}
        restricted to $\tau$, labelled with the operator name.
\end{itemize}
\end{definition}

Figure~\ref{fig:fsm-task} renders $\mathcal{F}_\tau$.

\begin{figure}[h]
\centering
\begin{tikzpicture}[
  every state/.style={draw, rounded rectangle, minimum width=2.2cm,
                      minimum height=0.85cm, align=center, font=\small,
                      inner sep=3pt},
  initial text={},
  arr/.style={-{Latex[length=2.5mm]}, line width=0.85pt},
  lbl/.style={font=\scriptsize, fill=white, inner sep=1.5pt},
]

%% Nodes
\node[state, fill=gray!10]   (ext)  at (0, 0.5)     {$\mathrm{ext}$};
\node[state, fill=cyan!10]   (Mij)  at (4.5, 0.5)   {$M_{ij}$};
\node[state, fill=green!12]  (Qi)   at (9, 0.5)     {$Q_i$};
\node[state, fill=yellow!18] (Ei)   at (9,-3.8)     {$E_i$};
\node[state, fill=orange!20] (W)    at (4.5,-2.2)   {$W$};
\node[state, fill=red!12, accepting] (sink) at (0,-3.8) {$\varnothing$};

%% Transitions
% external inject -> mailbox
\draw[arr] (ext) -- node[lbl,above] {$\hat{O}_{\text{ext}}$} (Mij);

% mailbox -> local queue (received, route_local)
\draw[arr] (Mij) -- node[lbl,above] {$\hat{O}_{\text{push-local}}$} (Qi);

% mailbox -> mailbox (hop)
\draw[arr] (Mij) to[loop above, looseness=6]
  node[lbl,above] {$\hat{O}_{\text{hop}}$\;($h \mapsto h{+}1$)} (Mij);

% mailbox -> warehouse (park after max hops or no space)
\draw[arr] (Mij) -- node[lbl,left] {$\hat{O}_{\text{park}}$} (W);

% local queue -> executing
\draw[arr] (Qi) -- node[lbl,right] {$\hat{O}_{\text{dispatch}}$} (Ei);

% local queue -> mailbox (deflect when q_i $\geq$ H)
\draw[arr] (Qi) to[bend left=20] node[lbl,above] {$\hat{O}_{\text{mailbox}}$} (Mij);

% executing -> finished (sink); bowed downward so it clears the W node,
% which is not collinear with Ei--sink after the coordinate change above,
% but we keep a slight bend for extra label clearance
\draw[arr] (Ei) to[bend right=12]
  node[lbl,below,pos=0.5,yshift=-2pt] {$\hat{O}_{\text{finish}}$} (sink);

% executing -> re-enqueue (Notified); bend away from the Ei node body and
% place the label past the curve's midpoint, clear of both Ei and Qi
\draw[arr] (Ei) to[bend left=35]
  node[lbl,right,pos=0.62,xshift=2pt] {$\hat{O}_{\text{requeue}}$} (Qi);

% warehouse -> local queue (drain)
\draw[arr] (W) -- node[lbl,above] {$\hat{O}_{\text{drain}}$} (Qi);

\end{tikzpicture}
\caption{Per-task FSM $\mathcal{F}_\tau$.  Each node is a site
$\alpha \in \mathcal{S}$; each edge is a transfer operator from
Definition~\ref{def:ops}.  The hop self-loop on $M_{ij}$ is
annotated with the side-effect $h \mapsto h+1$; the guard
$h < h_{\max}$ enables the loop and $h \geq h_{\max}$ forces the
park transition.  The unique absorbing state $\varnothing$ is
double-bordered.}
\label{fig:fsm-task}
\end{figure}

% =============================================================
\section{Key Structural Properties of the Operators}
% =============================================================

The following properties are stated here as the shared foundation
for the correctness, progress, and performance sub-modules.  Their
proofs appear in the respective sub-module documents.

\begin{definition}[Number-conserving operator]
\label{def:conserving}
An operator $\hat{O}$ is \emph{number-conserving} if for every
$\sigma \in \Sigma$ such that $\hat{O}(\sigma) \neq \bot$,
\[
  \Num(\hat{O}(\sigma)) = \Num(\sigma).
\]
Equivalently, $\hat{O}$ is a transfer $\Trf{\alpha}{\beta}(\tau)$ with
$\alpha, \beta \notin \{\varnothing\}$ and $\beta \neq \text{ext}$.
\end{definition}

\begin{definition}[Terminating hop chain]
\label{def:hop-chain}
A \emph{hop chain} is a maximal sequence of applications of
$\hat{O}_{\text{hop}}$ to the same chunk $c$.  The length of a hop
chain is bounded by $h_{\max} - c.h$, where $c.h$ is the chunk's
hop count when the chain begins.
\end{definition}

\begin{definition}[Wait-for relation]
\label{def:waitfor}
Worker $i$ \emph{waits for} worker $j$ (written $i \rightsquigarrow j$)
if worker $i$ is blocked in a state that requires worker $j$ to
perform some action before worker $i$ can proceed.  The
\emph{wait-for graph} $G_{\text{wf}} = (\mathcal{W}, \rightsquigarrow)$
has an edge $i \to j$ iff $i \rightsquigarrow j$.
\end{definition}

% =============================================================
\section{Correspondence Table: Math~$\leftrightarrow$~Code}
% =============================================================
\label{sec:corr}

\begin{table}[H]
\centering
\small
\caption{Mapping between mathematical objects and Rust source identifiers
in \texttt{src/dta\_scheduler.rs}.}
\label{tab:corr}
\begin{tabular}{>{\ttfamily}p{3.4cm} p{4.5cm} p{5.5cm}}
\toprule
\normalfont Math symbol & Meaning & Code identifier \\
\midrule
$\mathcal{T}$, $P$   & Task index set, pool capacity
  & \code{TaskIndex = u32}; pool capacity arg to \code{ContextPool::new} \\
$\mathcal{C}$, $K$   & Chunk set, chunk size
  & \code{TaskChunk}; \code{CHUNK\_SIZE = 32} \\
$\mathcal{W}$, $N$   & Worker index set, worker count
  & \code{workers: Vec<UnsafeCell<Worker>>}; \code{num\_workers} \\
$h_{\max}$           & Maximum hop count
  & \code{DtaScheduler::max\_hops} (computed as \code{num\_workers / 2}) \\
$Q_i$, $L$           & Local queue of worker $i$, capacity
  & \code{Worker::local\_queue}; \code{LOCAL\_QUEUE\_CAPACITY = 131072} \\
$H$                  & High-water mark (admission threshold)
  & \code{LOCAL\_QUEUE\_HIGH\_WATERMARK} \\
$M_{ij}$, $C_M$      & Mailbox from worker $i$ to $j$, capacity
  & \code{mailboxes[i][j]: Mailbox}; \code{MAILBOX\_CAPACITY = 65536} \\
$W$, $C_W$           & Warehouse, capacity in chunks
  & \code{DtaScheduler::warehouse: Warehouse}; \code{WAREHOUSE\_CAPACITY = 32768} \\
$E_i$                & Executing slot of worker $i$
  & \code{CURRENT\_FIBER} thread-local + fiber state \code{Running} \\
$\varnothing$        & Completed/freed (absorbing sink)
  & \code{pool.free\_context(task)} called on \code{Finished}/\code{Panicked} \\
$\cre{\alpha,\tau}$  & Create task $\tau$ at site $\alpha$
  & \code{Worker::push\_local}, \code{Mailbox::push}, \code{Warehouse::push} \\
$\ann{\alpha,\tau}$  & Remove task $\tau$ from site $\alpha$
  & \code{Mailbox::pop}, \code{Warehouse::pop}, local queue head advance \\
$\Num$               & Total live task count
  & No single variable; distributed across all queues \\
$\hat{O}_{\text{push-local}}$ & Mailbox pop + local queue push
  & \code{poll\_mailboxes} $\to$ \code{route\_chunk} $\to$ \code{push\_batch} \\
$\hat{O}_{\text{dispatch}}$   & Local queue pop + execute
  & \code{Worker::dispatch\_loop} \\
$\hat{O}_{\text{finish}}$     & Context free on task completion
  & \code{pool.free\_context(task)} in \code{dispatch\_loop} \\
$\hat{O}_{\text{requeue}}$    & Re-enqueue Notified task
  & \code{push\_local(task)} in \code{dispatch\_loop} (Notified branch) \\
$\hat{O}_{\text{mailbox}}$    & Local-to-mailbox transfer (deflect)
  & \code{push\_chunk\_with\_hop} \\
$\hat{O}_{\text{hop}}$        & Mailbox-to-mailbox re-deflect
  & \code{route\_deflect}; \code{hop\_count.saturating\_add(1)} \\
$\hat{O}_{\text{park}}$       & Mailbox-to-warehouse transfer
  & \code{park\_in\_warehouse} / \code{route\_park} \\
$\hat{O}_{\text{drain}}$      & Warehouse-to-local-queue drain
  & \code{drain\_warehouse} $\to$ \code{push\_batch} \\
$\hat{O}_{\text{ext}}$        & External thread injection
  & \code{external\_mailboxes[i].push} under \code{external\_locks[i]} \\
\bottomrule
\end{tabular}
\end{table}

% =============================================================
\section{Summary of Parameters}
% =============================================================

\begin{table}[H]
\centering
\small
\caption{Global parameters of the model with their concrete values in
the published benchmark configuration ($N=4$).}
\label{tab:params}
\begin{tabular}{cll r}
\toprule
Symbol & Meaning & Code constant & Benchmark value \\
\midrule
$N$        & Number of workers       & \code{num\_workers}              & 4 \\
$P$        & Context-pool capacity   & runtime arg                       & 8\,192 \\
$K$        & Chunk size              & \code{CHUNK\_SIZE}                & 32 \\
$L$        & Local queue capacity    & \code{LOCAL\_QUEUE\_CAPACITY}     & 131\,072 \\
$H$        & High-water mark         & \code{LOCAL\_QUEUE\_HIGH\_WATERMARK} & 114\,688 \\
$C_M$      & Mailbox capacity        & \code{MAILBOX\_CAPACITY}          & 65\,536 \\
$C_W$      & Warehouse capacity (chunks) & \code{WAREHOUSE\_CAPACITY}    & 32\,768 \\
$h_{\max}$ & Maximum hop count       & \code{max\_hops = num\_workers/2} & 2 \\
\bottomrule
\end{tabular}
\end{table}

\begin{notation}[On $\rho$/$\varrho$ and the two uses of $H$]
\label{not:symbols}
Two symbol collisions recur across chapters and are worth flagging
explicitly rather than leaving to context.

\emph{Load symbols.} $\rho$ (mean-field load) and $\varrho$ (its
stability-chapter counterpart) are typographically distinct but
denote closely related quantities; we use $\rho^*$ for the
self-consistency fixed point, $\varrho_0$ for the nominal per-worker
offered load, and $\varrho_c$ exclusively for a \emph{competitive
ratio} (never a load) in \hyperref[chap:competitive_ratio]{the
competitive-ratio chapters}. When in doubt, $\varrho_c^{(\cdot)}$ always
denotes a dimensionless ratio $\geq 1$; every other $\rho$/$\varrho$
variant denotes a load in $[0,1)$.

\emph{$H$.} Throughout this paper $H$ denotes the per-worker local-queue
\emph{high-water mark} (admission threshold; \code{LOCAL\_QUEUE\_HIGH\_WATERMARK}
in the table above), a quantity in the hundreds to hundreds-of-thousands.
It is unrelated to the \emph{harmonic number} $H_k = \sum_{j=1}^k 1/j$
that appears only in the DAG kinetic-theory chapter
(\S\ref{sec:dag_kinetic}) in expressions such as $H_K$; that usage is
always subscripted by a fan-in variable ($K$) rather than left bare, and
never coincides in context with the admission threshold. It is also
distinct from the \emph{hop-count bound} $h_{\max} = \lfloor N/2
\rfloor$ (lower-case $h$), which is a small dimensionless integer, not a
queue depth; an earlier draft of this paper conflated $H$ and $h_{\max}$
in the information-theoretic cost framework (\S\ref{sec:info_optimality_proof}),
which has since been corrected.
\end{notation}

\newpage
\clearpage

\begin{figure}[h]
\centering
\begin{tikzpicture}[
  worker/.style={circle, draw, minimum size=1.5cm, font=\small, line width=0.8pt, align=center},
  hot/.style={circle, draw, fill=red!15, minimum size=1.5cm, font=\small, line width=1.2pt, align=center},
  mbox/.style={-{Latex[length=2mm]}, thin, gray},
  deflect/.style={-{Latex[length=2.5mm]}, line width=1.1pt, red!70!black},
]
  \node[hot]  (W1) at (0,3) {$W_1$ \\[-2pt] {\tiny $q_1{>}H$}};
  \node[worker] (W2) at (3,3) {$W_2$};
  \node[worker] (W3) at (3,0) {$W_3$};
  \node[worker] (W4) at (0,0) {$W_4$};

  % full mesh of mailboxes
  \draw[mbox] (W1) -- (W2);
  \draw[mbox] (W1) -- (W3);
  \draw[mbox] (W1) -- (W4);
  \draw[mbox] (W2) -- (W3);
  \draw[mbox] (W2) -- (W4);
  \draw[mbox] (W3) -- (W4);

  % deflection hops from hot worker
  \draw[deflect] (W1) to[bend left=15] node[above,font=\tiny,sloped] {hop 1} (W2);
  \draw[deflect] (W1) to[bend right=15] node[below,font=\tiny,sloped] {hop 1} (W4);

  \node[draw, rectangle, fill=blue!8, minimum width=3.4cm, minimum height=1cm, font=\small] (WH) at (1.5,-2.6) {Warehouse (MPMC)};
  \draw[-{Latex[length=2mm]}, dashed, blue!60!black] (W1) -- (WH);
  \draw[-{Latex[length=2mm]}, dashed, gray] (W2) -- (WH);
  \draw[-{Latex[length=2mm]}, dashed, gray] (W3) -- (WH);
  \draw[-{Latex[length=2mm]}, dashed, gray] (W4) -- (WH);
  \coordinate (bluemid) at ($(W1)!0.55!(WH)$);
  \node[font=\tiny, text width=2.6cm, align=right, blue!60!black,
        draw=blue!40, fill=white, inner sep=2pt, rounded corners=1pt]
    (note) at (-2.6,-1.3) {after $h_{\max}$ failed hops \emph{(blue dashed line)}};
  \draw[-{Latex[length=1.5mm]}, blue!60!black, thin] (note.east) -- (bluemid);
\end{tikzpicture}
\caption{P2P mesh topology with $N=4$ workers. Gray lines: the $N(N-1)$ dedicated point-to-point mailboxes forming the full mesh. Red arrows: sender-initiated deflection hops from an overloaded worker ($W_1$, local queue above high-water mark $H$) to its neighbors. Dashed blue: overflow to the shared warehouse once the hop budget $h_{\max}$ is exhausted.}
\label{fig:mesh}
\end{figure}

\begin{figure}[h]
\centering
\begin{tikzpicture}[
  box/.style={draw, rectangle, rounded corners=2pt, minimum width=3.5cm, minimum height=1.1cm, align=center, font=\small, inner sep=4pt},
  decision/.style={draw, diamond, aspect=2.6, align=center, font=\small, minimum width=2.2cm, inner sep=2pt},
  arr/.style={-{Latex[length=2.5mm]}, line width=0.9pt},
  lbl/.style={font=\tiny, fill=white, inner sep=1pt},
]
  \node[box, fill=green!12] (spawn) at (0,5.6) {Spawn request\\(try \texttt{alloc\_context})};
  \node[decision] (ok) at (0,3.6) {context\\available?};
  \node[box, fill=green!12] (run) at (-4.6,3.6) {Run fiber\\on local queue};
  \node[decision] (path) at (4.6,3.6) {entry\\path?};
  \node[box, fill=yellow!18] (retry) at (4.6,1.1) {Rust API: adaptive\\spin / back-off\\(blocks producer)};
  \node[box, fill=red!15] (panic) at (8.6,1.1) {FFI / internal call:\\\texttt{.expect()} panics\\$\to$ process abort};

  \draw[arr] (spawn) -- (ok);
  \draw[arr] (ok) -- node[lbl,above] {yes} (run);
  \draw[arr] (ok) -- node[lbl,above] {no} (path);
  \draw[arr] (path) -- node[lbl,left] {retry-capable} (retry);
  \draw[arr] (path) -| node[lbl,above] {fixed call site} (panic);
  \draw[arr] (retry.west) -- ++(-1.0,0) |- (ok.south);
\end{tikzpicture}
\caption{Tier A: context-pool admission. A spawn request that finds the pool exhausted is handled \emph{inconsistently} across API surfaces: the primary Rust API retries with adaptive back-off (genuine backpressure that blocks the producer), while several FFI and internal call sites assert success and panic immediately, which aborts the process under the release profile's \texttt{panic = "abort"} setting.}
\label{fig:backpressureA}
\end{figure}

\begin{figure}[h]
\centering
\begin{tikzpicture}[
  box/.style={draw, rectangle, rounded corners=2pt, minimum width=3.6cm, minimum height=1.1cm, align=center, font=\small, inner sep=4pt},
  decision/.style={draw, diamond, aspect=2.6, align=center, font=\small, minimum width=2.2cm, inner sep=2pt},
  arr/.style={-{Latex[length=2.5mm]}, line width=0.9pt},
  lbl/.style={font=\tiny, fill=white, inner sep=1pt},
]
  \node[box, fill=green!12] (localq) at (0,7.4) {Local queue $q_i<H$\\(per worker, Tier B)};
  \node[decision] (hw) at (0,5.4) {$q_i \ge H$?};
  \node[box, fill=green!12] (dispatch) at (-4.4,5.4) {Dispatch\\(run on $W_i$)};
  \node[box, fill=yellow!18] (deflect) at (4.4,5.4) {Deflect: $\le h_{\max}$\\mailbox hops};
  \node[box, fill=orange!20] (wh) at (4.4,3.0) {Warehouse\\(MPMC, capacity $C_W$)};
  \node[decision] (whfull) at (4.4,1.0) {warehouse\\full?};
  \node[box, fill=red!15] (crash) at (8.2,1.0) {Panic\\(crash-on-overflow)};
  \node[box, fill=green!12] (drain) at (0.6,-1.0) {Drained by any\\of $N$ workers};

  \draw[arr] (localq) -- (hw);
  \draw[arr] (hw) -- node[lbl,above] {no} (dispatch);
  \draw[arr] (hw) -- node[lbl,above] {yes} (deflect);
  \draw[arr] (deflect) -- node[lbl,right] {all hops fail} (wh);
  \draw[arr] (wh) -- (whfull);
  \draw[arr] (whfull) -- node[lbl,above] {yes} (crash);
  \draw[arr] (whfull) -| node[lbl,above] {no} (drain);
\end{tikzpicture}
\caption{Tier B: per-worker local-queue admission, hop-bounded mesh deflection, and warehouse overflow. The terminal failure mode reached after warehouse saturation is process termination, not flow control, despite the project's own internal terminology referring to this path as backpressure; in the published benchmark configuration Tier B's nominal capacity is an order of magnitude larger than Tier A's, so this path is a deep, rarely-triggered safety net rather than a routinely encountered one.}
\label{fig:backpressureB}
\end{figure}

\begin{figure}[h]
\centering
\begin{tikzpicture}[
  state/.style={draw, rectangle, rounded corners=2pt, minimum width=2.6cm, minimum height=1cm, align=center, font=\small},
  arr/.style={-{Latex[length=2.5mm]}, line width=0.9pt},
  lbl/.style={font=\tiny, fill=white, inner sep=1pt},
]
  \node[state, fill=gray!10]  (init)     at (0,0)     {Initial};
  \node[state, fill=green!15] (running)  at (4.2,0)    {Running};
  \node[state, fill=yellow!18](susp)     at (8.4,0)    {Suspending};
  \node[state, fill=orange!18](yielded)  at (8.4,-3.2) {Yielded};
  \node[state, fill=blue!12]  (notified) at (4.2,-3.2) {Notified};
  \node[state, fill=red!12]   (finished) at (0,-3.2)   {Finished};

  \draw[arr] (init) -- node[lbl,above] {dispatched} (running);
  \draw[arr] (running) -- node[lbl,above] {poll $\Rightarrow$ Pending} (susp);
  \draw[arr] (susp) -- node[lbl,right] {CAS wins} (yielded);
  \draw[arr] (susp.east) to[out=20,in=70,looseness=2.2] node[lbl,pos=0.5] {waker swap races CAS} ++(2.6,1.6) -| (notified.north);
  \draw[arr] (yielded) -- node[lbl,above] {waker fires} (notified);
  \draw[arr] (notified) -- node[lbl,above] {re-dispatched} (running);
  \draw[arr] (running.south) -- node[lbl,left] {poll $\Rightarrow$ Ready} (finished.north);
\end{tikzpicture}
\caption{Fiber status state machine. The \texttt{Suspending}/\texttt{Notified} race is the central correctness property of the cooperative-suspension mechanism and the focus of the project's loom-based model checking.}
\label{fig:fsm}
\end{figure}

\bigskip
\noindent\textbf{Summary.}
The formal framework defined in this part provides the algebraic backbone
for all subsequent analysis.  The Fock-space representation makes task
conservation a commutation identity, deflection a bounded hop-count operator,
and stability a spectral property of the generator.  Every later theorem
references a specific definition or operator from this chapter. As noted
above, the model's site-locality --- each operator acts on, or transfers
between, only the sites it names --- is what encodes the
bounded-information restriction at the formal level, so every downstream
correctness and performance result inherits it automatically rather than
needing to re-assume it.


\clearpage

\partbreak
\part{Correctness}
\clearpage

\bigskip
\noindent\textbf{Overview.}
This part establishes four safety properties of the DTA scheduler:
task conservation (no task is ever lost or duplicated),
deadlock freedom (no circular wait is possible under the lock-free discipline),
livelock freedom (every deflection sequence terminates within $\lfloor N/2\rfloor$ hops),
and full linearizability of both the SPSC mailboxes and the Vyukov MPMC
warehouse.  Each chapter provides a self-contained theorem with a proof
that refers to specific operators from Part~I.  These results establish
that restricting each worker to local information does not weaken
correctness: none of the four properties requires any worker to observe
another worker's state.


% ========================================================================
% Correctness I --- No Task Loss
% ========================================================================

\section*{Correctness I --- No Task Loss}
\addcontentsline{toc}{section}{Correctness I --- No Task Loss}
\label{chap:no_task_loss}
\setcounter{section}{0}

All symbols ($\mathcal{T}$, $\mathcal{S}$, $\Sigma$, $\Num$,
$\Trf{\cdot}{\cdot}$, $\hat{O}_*$) are defined in
\hyperref[chap:formal_model]{the Formal Model}, Sections 2--5.  We reproduce only
the definitions needed for self-containedness.

\bigskip

% =============================================================
\section{Statement of the Property}
% =============================================================

\begin{definition}[No-task-loss invariant]
\label{def:ntl}
A run of the scheduler satisfies \emph{no task loss} if, for every
task index $\tau \in \mathcal{T}$ that has been admitted to the system
(i.e., the corresponding $\hat{O}_{\text{ext}}$ or an internal spawn
has been executed), one of the following holds at every subsequent
moment:
\begin{enumerate}
  \item $\tau$ resides in exactly one site
        $\alpha \in \mathcal{S} \setminus \{\varnothing\}$ (it is live), or
  \item $\tau$ resides in $\varnothing$ (it has completed execution and
        its context slot has been freed).
\end{enumerate}
Equivalently, no transition can cause $\tau$ to vanish from
$\mathcal{S}$ while it is still live.
\end{definition}

This is the scheduler analogue of \emph{particle-number conservation}
in statistical mechanics.  In the occupation-number representation
(Definition 3.2 of the model document) it reads:
\[
  \sum_{\alpha \in \mathcal{S} \setminus \{\varnothing\}}
  n_{\alpha,\tau}(\sigma) \;+\; n_{\varnothing,\tau}(\sigma)
  \;\leq\; 1
  \qquad \forall\, \tau \in \mathcal{T},\; \forall\, \sigma \in \Sigma,
\]
with equality once $\tau$ has been admitted and until it has
completed.

% =============================================================
\section{Conservation of $\Num$ on Live Transitions}
% =============================================================

We prove the stronger statement that every scheduler operation
\emph{either} leaves $\Num$ unchanged (a ``live transition'') \emph{or}
decreases it by exactly 1 (the finish transition, which moves a task
to the absorbing sink $\varnothing$).

\begin{lemma}[Atomic transfers are number-conserving or number-decreasing]
\label{lem:atomicconserve}
For every atomic transfer $\Trf{\alpha}{\beta}(\tau)$ defined in
Section 5 of the model document:
\[
  \Num(\Trf{\alpha}{\beta}(\tau)(\sigma)) =
  \begin{cases}
    \Num(\sigma)     & \text{if } \beta \neq \varnothing, \\
    \Num(\sigma) - 1 & \text{if } \beta = \varnothing.
  \end{cases}
\]
\end{lemma}

\begin{proof}
By Definition 4.1 of the model document, $\Trf{\alpha}{\beta}(\tau) =
\cre{\beta,\tau} \circ \ann{\alpha,\tau}$.  The annihilation step
removes $\tau$ from $\sigma(\alpha)$ (decreasing $\num{\alpha}$ by 1)
and the creation step inserts $\tau$ into $\sigma(\beta)$ (increasing
$\num{\beta}$ by 1).  Thus
\[
  \Num(\Trf{\alpha}{\beta}(\tau)(\sigma))
  = \Num(\sigma) - \mathbf{1}[\alpha \notin \{\varnothing\}]
                 + \mathbf{1}[\beta \notin \{\varnothing\}].
\]
Because every live transition has $\alpha \notin \{\varnothing\}$
(tasks are only moved \emph{from} live sites) and either
$\beta \notin \{\varnothing\}$ (live-to-live) or $\beta = \varnothing$
(live-to-sink), the result follows.
\end{proof}

% =============================================================
\section{Verification: Each Named Operator}
% =============================================================

We now verify Lemma~\ref{lem:atomicconserve} explicitly for every
operator $\hat{O}_*$ from Definition 5.2 of the model document.

\begin{table}[h]
\centering
\small
\caption{Conservation check for each scheduler operator.
$\Delta\Num \coloneqq \Num(\hat{O}(\sigma)) - \Num(\sigma)$.
The ``Code guard'' column cites the Rust condition that prevents the
$\bot$ (failed) case from silently dropping a task.}
\label{tab:conserve}
\begin{tabularx}{\textwidth}{@{}clXcX@{}}
\toprule
Operator & Source site $\alpha$ & Sink site $\beta$
         & $\Delta\Num$ & Code guard \\
\midrule
$\hat{O}_{\text{ext}}$
  & $\text{ext}$ & $M_{\text{ext},i}$
  & $+1$ (admission) & \code{SpinLock} ensures push succeeds \\
$\hat{O}_{\text{push-local}}$
  & $M_{ji}$ & $Q_i$
  & $0$ & \code{route\_chunk} only called after \code{pop} succeeds \\
$\hat{O}_{\text{dispatch}}$
  & $Q_i$ & $E_i$
  & $0$ & head advance after task read, atomic Relaxed \\
$\hat{O}_{\text{finish}}$
  & $E_i$ & $\varnothing$
  & $-1$ & \code{free\_context} only on Finished/Panicked \\
$\hat{O}_{\text{requeue}}$
  & $E_i$ & $Q_i$
  & $0$ & \code{push\_local} asserted not-fail (debug\_assert) \\
$\hat{O}_{\text{mailbox}}$
  & $Q_i$ & $M_{ij}$
  & $0$ & caller enters \code{push\_chunk\_with\_hop} \\
$\hat{O}_{\text{hop}}$
  & $M_{ij}$ & $M_{ij'}$
  & $0$ & \code{route\_deflect}: pop then push to new mailbox \\
$\hat{O}_{\text{park}}$
  & $M_{ij}$ or $Q_i$ & $W$
  & $0$ & \code{park\_in\_warehouse} asserts push success \\
$\hat{O}_{\text{drain}}$
  & $W$ & $Q_i$
  & $0$ & \code{push\_batch} called only after \code{warehouse.pop} \\
\bottomrule
\end{tabularx}
\end{table}

\begin{remark}[The $\hat{O}_{\text{park}}$ assert and crash-on-overflow]
\label{rem:crash-vacuous}
The code guard for $\hat{O}_{\text{park}}$ is a Rust \code{assert!}
that panics (and aborts the process under \code{panic = "abort"}) if
the warehouse is full.  In the context of the present lemma, this is
\emph{conservation-preserving in the strong sense}: the operation
either succeeds (task moves to $W$, $\Delta\Num = 0$) or the process
terminates entirely (no further scheduler state exists).  The
no-task-loss invariant, as formalised in this chapter, is a statement
about the \emph{scheduler's internal state} $\sigma_t$ (Definition~2.1);
it is vacuously satisfied in the terminated case because there is no
further $\sigma_t$ for the invariant to range over.

We flag explicitly what this does \emph{not} claim.  A crash is not a
counterexample to no-task-loss under this state-based definition, but it
is not conservation under any external, \emph{durability}-based notion:
tasks that had been admitted to the scheduler and not yet completed at
the moment of the crash are unrecoverable once the process exits, and
this remark should not be read as asserting otherwise. The theorem's
scope is therefore implicitly conditioned on \emph{process survival}; a
full durability argument (e.g.\ persisting admitted-but-incomplete tasks
to stable storage before acknowledging admission) is outside the scope
of this paper and is not provided by DTA.  The semantic distinction
--- crash-on-overflow vs.\ block-on-overflow --- is a \emph{liveness}
concern addressed in the progress sub-module and does not change this
scope limitation.
\end{remark}

% =============================================================
\section{Main Theorem}
% =============================================================

\begin{theorem}[No task loss]
\label{thm:ntl}
Let $\sigma_0, \sigma_1, \sigma_2, \ldots$ be any execution trace of
the DTA scheduler (a sequence of states produced by applying
operators from Definition 5.2 of the model document).  Then for every
task index $\tau$ that has been admitted at step $t_0$
(i.e., $\hat{O}_{\text{ext}}$ was applied at step $t_0$):

\begin{enumerate}
  \item \textbf{(Uniqueness)} At every step $t \geq t_0$,
        $\sum_{\alpha \in \mathcal{S}} n_{\alpha,\tau}(\sigma_t) = 1$.

  \item \textbf{(Monotone termination)} The sequence
        $n_{\varnothing,\tau}(\sigma_t)$ is non-decreasing in $t$
        and takes values in $\{0, 1\}$.  Once it reaches 1
        (task completed), it remains 1.
\end{enumerate}
\end{theorem}

\begin{proof}
\textbf{Part 1.}  By induction on $t$.

\emph{Base case} $t = t_0$: the admit operator
$\hat{O}_{\text{ext}}(\tau)$ moves $\tau$ from the external source to
$M_{\text{ext},i}$.  By the guard in Table~\ref{tab:conserve}, this
succeeds atomically (the external spinlock serialises concurrent
external pushes), so $n_{M_{\text{ext},i},\tau}(\sigma_{t_0}) = 1$
and $n_{\alpha,\tau} = 0$ for all other $\alpha$.  Thus
$\sum_\alpha n_{\alpha,\tau} = 1$.

\emph{Inductive step}: assume $\sum_\alpha n_{\alpha,\tau}(\sigma_t)
= 1$ and that $\hat{O}$ is applied to produce $\sigma_{t+1}$.  If
$\hat{O}$ does not involve $\tau$ at all, $\sigma_{t+1}$ restricted to
$\tau$ is identical to $\sigma_t$ and the sum is unchanged.  If
$\hat{O} = \Trf{\alpha}{\beta}(\tau)$ for some $\alpha, \beta$, then
by Lemma~\ref{lem:atomicconserve}:
\begin{itemize}
  \item If $\beta \neq \varnothing$: $\Num$ is conserved, so the sum
        over all live sites plus $\varnothing$ remains 1 (with $\tau$
        having moved from $\alpha$ to $\beta$).
  \item If $\beta = \varnothing$ (i.e., $\hat{O} =
        \hat{O}_{\text{finish}}$): $n_{\alpha,\tau}$ decreases by 1
        and $n_{\varnothing,\tau}$ increases by 1, preserving the
        total sum.
\end{itemize}
In both cases $\sum_\alpha n_{\alpha,\tau}(\sigma_{t+1}) = 1$.

\textbf{Part 2.}  The only operator that sets
$n_{\varnothing,\tau} = 1$ is $\hat{O}_{\text{finish}}$, which
requires $\tau \in \sigma(E_i)$ (the task is currently executing) and
that the fiber's state is \code{Finished} or \code{Panicked}.  Once
$n_{\varnothing,\tau} = 1$, no subsequent operator has $\varnothing$
as a source site (it is an absorbing sink by Definition 3.1 of the
model document), so the value cannot decrease.
\end{proof}

\begin{corollary}[No silent drop]
No sequence of scheduler operations can cause a live task to vanish
from all sites without first passing through $\hat{O}_{\text{finish}}$.
\end{corollary}

% =============================================================
\section{Relationship to the Implementation}
% =============================================================

The proof above is structural: it follows from the fact that every
code path that removes a task from one location immediately inserts it
into another (or asserts/panics if the insertion fails).  The critical
implementation invariant is that \emph{task indices are never copied
into two queues simultaneously}: the SPSC mailboxes use a
head/tail protocol that makes a slot invisible to the consumer until
the producer's \code{Release} store, and the local queue's push and
pop are always performed by the owning worker thread (no concurrent
access).

The one place where the conservation invariant is \emph{maintained
by the absence of a code path} rather than by an explicit insertion
is the \code{Notified} re-enqueue branch in \code{dispatch\_loop}:
the task index \code{task} is stored in the local variable just before
the context switch, and is pushed back via \code{push\_local} only in
the \code{Notified} branch; in all other branches (\code{Finished},
\code{Panicked}) it is explicitly freed.  If the re-enqueue
\code{push\_local} call returned \code{false} (queue full), the task
would be silently dropped.  The code guards this with a
\code{debug\_assert}: the assertion holds because
\code{dispatch\_loop} only enters the re-enqueue branch after
\emph{popping} the task from the same queue, so there is at least one
free slot.  In the formal model, this corresponds to the fact that
$\hat{O}_{\text{dispatch}} \circ \hat{O}_{\text{requeue}}$ decreases
$|Q_i|$ by 1 and then increases it by 1, never exceeding the
pre-dispatch length.


\clearpage

% ========================================================================
% Correctness II --- Deadlock Freedom
% ========================================================================

\section*{Correctness II --- Deadlock Freedom}
\addcontentsline{toc}{section}{Correctness II --- Deadlock Freedom}
\label{chap:deadlock}

\setcounter{section}{0}


% =============================================================
\section{Statement of the Property}
% =============================================================

\begin{definition}[Deadlock]
\label{def:deadlock}
A system state $\sigma \in \Sigma$ is a \emph{deadlock} if there
exists a non-empty set $D \subseteq \mathcal{W}$ of workers such that
every worker $i \in D$ is blocked waiting for some action by another
worker $j \in D$ to complete before $i$ can make progress.

Formally, using the wait-for relation $\rightsquigarrow$ of
Definition 6.3 of the model document, a deadlock requires the
subgraph $G_{\text{wf}}\!\upharpoonright_D$ to contain a directed
cycle.
\end{definition}

The classical \emph{Coffman conditions} for deadlock are (see also~\cite{lenstra1978} for scheduling complexity):
\begin{enumerate}
  \item \textbf{Mutual exclusion}: a resource is held by at most one
        process at a time.
  \item \textbf{Hold and wait}: a process holding a resource waits to
        acquire another.
  \item \textbf{No preemption}: resources cannot be forcibly taken.
  \item \textbf{Circular wait}: a cycle exists in the wait-for graph.
\end{enumerate}

We prove deadlock freedom by showing that Coffman condition 4 is
never satisfied: $G_{\text{wf}}$ has no edges at all, making a
cycle impossible.

% =============================================================
\section{Classification of Blocking Operations}
% =============================================================

We systematically enumerate every operation in the scheduler that
could \emph{in principle} cause a worker to block, and show that none
of them creates an edge in $G_{\text{wf}}$.

\begin{definition}[Blocking vs.\ non-blocking]
An operation is \emph{worker-blocking} if it suspends the calling
worker until a \emph{specific other worker} $j$ completes an action.
An operation is \emph{contention-delayed} if it retries a
compare-and-swap (CAS) or spin until a shared atomic is updated by
\emph{any} concurrent actor, but no specific worker is awaited.
An operation is \emph{locally spinning} if it performs only
reads/writes to memory and hardware instructions without waiting for
any external party.
\end{definition}

\begin{lemma}[All queue operations are non-blocking]
\label{lem:nonblocking}
Every queue operation in the DTA scheduler is either immediately
successful or immediately returns a failure indicator
(\code{Err(chunk)}, \code{None}, or \code{false}).
No queue operation suspends the calling worker thread.
\end{lemma}

\begin{proof}
We enumerate by structure type.

\medskip
\textbf{SPSC Mailbox (\code{Mailbox::push}/\code{pop}).}
\code{push} loads \code{tail} (Relaxed), computes
\code{next\_tail}, and checks whether \code{next\_tail == head}
(Acquire).  If equal, the mailbox is full and \code{Err(chunk)} is
returned \emph{immediately} without any wait.  \code{pop} similarly
checks \code{tail == head} and returns \code{None} if empty.
Neither operation has a retry loop or a park/block call.

\medskip
\textbf{Local queue (\code{push\_local}).}
Checks \code{local\_queue\_len() >= LOCAL\_QUEUE\_CAPACITY - 1}
and returns \code{false} immediately if full; no wait.

\medskip
\textbf{MPMC Warehouse (\code{Warehouse::push}/\code{pop}).}
Both operations contain a CAS retry loop (\code{compare\_exchange\_weak}
in production).  The loop retries when a concurrent actor claims the
same slot first (\code{Greater} arm) or the CAS spuriously fails.
Between retries, \code{warehouse\_backoff} executes a counted NOP
sequence (a fixed number of \code{nop} instructions).  Crucially:
\begin{itemize}
  \item The NOP loop completes in a bounded number of cycles
        (at most $32 \cdot 2^6 = 2048$ cycles per retry, cap at retry
        depth 6).
  \item The loop body does not \emph{read} any worker-owned resource;
        it waits only for the hardware memory bus to resolve the CAS
        on a shared atomic.
  \item The CAS target (\code{Warehouse::tail} or \code{Warehouse::head})
        is advanced by \emph{some} contending actor on each collision.
        This is precisely the \emph{lock-free} progress condition of
        Herlihy and Shavit~\cite{herlihy_shavit2008}: the system as a
        whole makes progress on every collision (some contender's CAS
        succeeds), even though any \emph{individual} worker's retry
        count is not deterministically bounded under adversarial
        scheduling. (This is stronger than \emph{obstruction-free}, which
        only guarantees progress for a thread running in isolation; we
        do not need obstruction-freedom here because the lock-free
        guarantee is already available for the Vyukov queue's CAS loop.)
        The system-wide lock-free guarantee alone only bounds
        \emph{aggregate} progress (some contender's CAS succeeds on
        every collision); it does not by itself bound any
        \emph{individual} worker's retry count, which could in
        principle be starved by an adversarial scheduler that always
        lets a different contender win. We add one standard fairness
        assumption to close this gap: a \emph{fair scheduler} does not
        systematically bias the winner of a CAS race away from any
        fixed worker, i.e., conditional on a collision occurring among
        the (at most $N-1$ other) contenders active at that instant,
        each contending worker's CAS wins with probability at least
        $1/N$. Under this assumption, worker $i$'s retry count $R_i$ is
        stochastically dominated by a $\mathrm{Geometric}(1/N)$ random
        variable (each attempt independently succeeds with probability
        $\geq 1/N$, conditionally on a retry being needed at all), giving
        the explicit bound $\mathbb{E}[R_i] \leq N$ --- not merely
        $\mathbb{E}[R_i]<\infty$, but a concrete worst-case constant in
        terms of the contention width. The exponential back-off
        (Assumption~\ref{ass:live}-compatible: it only ever delays,
        never blocks, this worker) additionally spreads out repeated
        self-collisions in time without affecting this bound's validity,
        since the domination argument only uses the per-attempt success
        probability, not the inter-attempt delay.
        A finite-a.s.\ (indeed finite-mean) retry count is exactly what
        ruling out deadlock (Definition~\ref{def:deadlock}) requires: it
        precludes \emph{permanent} blocking.
\end{itemize}
Therefore, the warehouse retry loop is \emph{contention-delayed}, not
\emph{worker-blocking}: it does not wait for any specific worker
$j \in \mathcal{W}$, and hence adds no directed edge to $G_{\text{wf}}$.

\medskip
\textbf{External mailbox (\code{SpinLock}).}
The spinlock protecting external mailbox writes is a test-and-set
spin.  The lock is held only during a single \code{Mailbox::push}
call (a handful of atomic instructions).  It does not wait for any
worker thread; it waits for another \emph{external (host)} thread
to release the same lock.  Host threads are outside $\mathcal{W}$,
so no edge $i \to j$ with $i, j \in \mathcal{W}$ is created.
\end{proof}

% =============================================================
\section{The Wait-for Graph is Always Empty}
% =============================================================

\begin{theorem}[Deadlock freedom]
\label{thm:deadlock}
For every reachable state $\sigma \in \Sigma$, the wait-for graph
$G_{\text{wf}}(\sigma)$ has no edges.  In particular, it has no
directed cycle, and the system is deadlock-free.
\end{theorem}

\begin{proof}
An edge $i \rightsquigarrow j$ (worker $i$ waits for worker $j$)
requires that worker $i$ is suspended at an operation that can only
proceed after worker $j$ performs some specific action.

By Lemma~\ref{lem:nonblocking}, every queue operation either
succeeds immediately or fails immediately.  On failure, the
\emph{caller} (not the blocking worker) proceeds to the next
alternative:
\begin{itemize}
  \item \code{Mailbox::push} fails $\Rightarrow$ caller tries
        next mailbox (hop) or parks in warehouse.
  \item \code{Mailbox::pop} returns \code{None} $\Rightarrow$ caller
        continues to poll the next mailbox in \code{polling\_order}.
  \item \code{push\_local} fails $\Rightarrow$ caller wraps in chunk
        and calls \code{push\_chunk\_with\_hop}.
\end{itemize}
In all cases the calling worker makes progress \emph{immediately}
(either succeeds or switches to a fallback path).  It does not wait
for any specific $j \in \mathcal{W}$.

The warehouse CAS retry is contention-delayed but not
worker-blocking (Lemma~\ref{lem:nonblocking}, warehouse case).
The idle backoff in \code{run\_worker\_static} (Tier 1/2/3 spin
and WFE/umonitor) is a local wait on the worker's own
\code{event\_signal} counter --- it does not hold any resource that
another worker needs, and it wakes either on a timeout or on
\code{signal\_worker} writing to \code{event\_signal} (a
non-blocking write).

Therefore, no operation in the scheduler creates an edge $i
\rightsquigarrow j$ in $G_{\text{wf}}$.  Since
$G_{\text{wf}}(\sigma) = (\mathcal{W}, \varnothing)$ for all
$\sigma$, it trivially contains no directed cycle, and by
Definition~\ref{def:deadlock} no state is a deadlock.
\end{proof}

% =============================================================
\section{Coffman Condition Analysis}
% =============================================================

For completeness, we show which Coffman conditions hold and which fail.

\begin{table}[h]
\centering
\small
\caption{Coffman conditions for DTA.}
\label{tab:coffman}
\begin{tabularx}{\textwidth}{@{}clXl@{}}
\toprule
Condition & Holds? & Reason \\
\midrule
Mutual exclusion
  & Yes, on the Warehouse only
  & The MPMC Warehouse is a genuinely shared resource among all $N$
    workers (mutual exclusion applies in the classical sense: at most
    one CAS winner claims a slot). Each SPSC mailbox, by contrast, has
    a statically fixed unique producer and unique consumer, so it is
    never actually \emph{contended} between two workers competing for
    the same role; mutual exclusion holds for it vacuously rather than
    via any exclusion mechanism. We list both as ``Yes'' because
    Coffman condition~1 is satisfied either way; the distinction
    matters only for condition~4 (Circular wait) below, where SPSC's
    fixed roles are what rules out contention-driven cycles. \\
Hold and wait
  & No
  & Workers never hold a queue slot while waiting for another queue. \\
  & &  A failed push returns immediately; the chunk is held in a \\
  & &  local variable on the stack, not in a shared resource. \\
No preemption
  & Yes
  & Queue slots are not forcibly taken from a holder. \\
Circular wait
  & \textbf{No}
  & $G_{\text{wf}}$ is empty (Theorem~\ref{thm:deadlock}). \\
\bottomrule
\end{tabularx}
\end{table}

The absence of \emph{hold-and-wait} is the structural key: when a
push fails, the chunk remains in a local variable (on the caller's
stack or register) rather than being retained inside any shared
structure.  The caller then either tries a different target or parks
in the warehouse --- a \emph{non-blocking} action.  This pattern is
precisely what the hop-bounded deflection protocol enforces: at no
point does a worker occupy a shared queue slot while simultaneously
waiting for another queue slot to become free.

% =============================================================
\section{Note on the Warehouse Panic Path}
% =============================================================

\begin{remark}
The \code{park\_in\_warehouse} function asserts that
\code{warehouse.push} succeeds and will \code{panic!} (aborting the
process) if the warehouse is full.  This is \emph{not} a deadlock:
the process terminates rather than entering a state where workers wait
for each other.  It is a \emph{liveness failure} (the system stops
making progress) but not a safety violation in the deadlock sense.
The distinction matters: deadlock is a state where the system is
\emph{alive but stuck}; process termination is not a stuck state.
The warehouse overflow scenario is analysed in the \hyperref[chap:bounded_wait]{Progress~I}.
\end{remark}


\clearpage

% ========================================================================
% Correctness III --- Livelock Freedom
% ========================================================================

\section*{Correctness III --- Livelock Freedom}
\addcontentsline{toc}{section}{Correctness III --- Livelock Freedom}
\label{chap:livelock}


\setcounter{section}{0}


% =============================================================
\section{Statement of the Property}
% =============================================================

\begin{definition}[Livelock]
\label{def:livelock}
A \emph{livelock} is an infinite execution trace
$\sigma_0, \sigma_1, \sigma_2, \ldots$ such that
\begin{enumerate}
  \item Some task index $\tau$ never reaches $\varnothing$
        (it never completes), and
  \item The system \emph{does} keep changing state (it is not stuck
        as in a deadlock) --- in particular, the chunk carrying $\tau$
        keeps being re-routed indefinitely.
\end{enumerate}
\end{definition}

In the DTA context, the only candidate for an infinite loop is the
hop-bounded deflection protocol: a chunk could in principle be
forwarded from mailbox to mailbox forever if the hop count were
unbounded.  We show this cannot happen.

% =============================================================
\section{The Lyapunov Function}
% =============================================================

\begin{definition}[Hop-count Lyapunov function]
\label{def:lyapunov}
For a chunk $c = (\bm{\tau}, k, h) \in \mathcal{C}$, define its
\emph{deflection potential}:
\[
  V(c) \;\coloneqq\; h_{\max} - h \;\in\; \{0, 1, \ldots, h_{\max}\}.
\]
\end{definition}

This function plays the role of a Lyapunov potential in dynamical
systems theory: it is non-negative, bounded, and strictly
decreasing along every application of the hop operator.

\begin{lemma}[Strict decrease under hopping]
\label{lem:decrease}
For any chunk $c$ with $V(c) > 0$ (i.e., $h < h_{\max}$), the
application of $\hat{O}_{\text{hop}}$ produces a chunk $c'$ with
\[
  V(c') = V(c) - 1.
\]
\end{lemma}

\begin{proof}
$\hat{O}_{\text{hop}}$ increments the hop count field:
$c' = (\bm{\tau}, k, h+1)$.  Therefore
$V(c') = h_{\max} - (h+1) = (h_{\max} - h) - 1 = V(c) - 1$.
\end{proof}

\begin{lemma}[Absorbing boundary at $V = 0$]
\label{lem:absorb}
If $V(c) = 0$ (i.e., $h = h_{\max}$), the operator
$\hat{O}_{\text{hop}}$ is not applicable to $c$.
Instead, $\hat{O}_{\text{park}}$ is the only available transition
from any mailbox site occupied by $c$.
\end{lemma}

\begin{proof}
Inspection of \code{push\_chunk\_with\_hop} and \code{route\_chunk}
in \texttt{src/dta\_scheduler.rs}:

In \code{push\_chunk\_with\_hop} (the cross-worker injection path),
the guard is:
\begin{verbatim}
if chunk.hop_count >= self.max_hops {
    return self.park_in_warehouse(*chunk);
}
\end{verbatim}
This guard is checked \emph{before} incrementing \code{hop\_count}
and before attempting the next mailbox push, so a chunk with
$h = h_{\max}$ is immediately diverted to the warehouse.

In \code{route\_chunk} (the inbound reception path), the 4-way
dispatch table is:
\begin{verbatim}
let hops_ok = chunk.hop_count < self.max_hops;
let idx = usize::from(space_ok) | (usize::from(hops_ok) << 1);
// idx=10 (no space, hops ok) $\to$ route_deflect
// idx=00 (no space, no hops) $\to$ route_park (warehouse)
\end{verbatim}
When $h = h_{\max}$, \code{hops\_ok} is \code{false}, so
\code{idx} $\in \{0, 1\}$, both of which map to either
\code{route\_local} (space available) or \code{route\_park} (no
space).  Neither calls \code{route\_deflect}.

Therefore, the guard $h < h_{\max}$ is a necessary condition for
$\hat{O}_{\text{hop}}$ to be invoked; when it fails, $\hat{O}_{\text{park}}$
or $\hat{O}_{\text{push-local}}$ is invoked instead.
\end{proof}

% =============================================================
\section{Main Theorem}
% =============================================================

\begin{theorem}[Deflection protocol termination / Livelock freedom]
\label{thm:livelock}
For any chunk $c$ with hop count $h \in \{0, \ldots, h_{\max}\}$,
the sequence of transitions that $c$ undergoes beginning from any
mailbox site terminates in at most $h_{\max} - h + 1$ steps.
In particular, the deflection protocol applied to any chunk
terminates in at most $h_{\max} + 1 = \floor{N/2} + 1$ steps.
\end{theorem}

\begin{proof}
We use $V(c) = h_{\max} - h$ as the termination measure.

\medskip
Consider the state sequence $c_0 = c,\, c_1,\, c_2, \ldots$ of a
chunk under the deflection protocol.  At each step, exactly one of
three things happens:

\begin{enumerate}
  \item \textbf{Success ($\hat{O}_{\text{push-local}}$):}
        The chunk is accepted into a worker's local queue.  The
        deflection protocol terminates.  $V$ is irrelevant after
        this point.

  \item \textbf{Hop ($\hat{O}_{\text{hop}}$):}
        $V(c_{t+1}) = V(c_t) - 1$ (Lemma~\ref{lem:decrease}).
        This can happen at most $V(c_0)$ times before $V$ reaches 0.

  \item \textbf{Park ($\hat{O}_{\text{park}}$):}
        Either $V(c_t) = 0$ forces this (Lemma~\ref{lem:absorb}),
        or all probed mailboxes are full and the chunk is parked
        directly.  The deflection protocol terminates; $c$ is now
        in $W$.
\end{enumerate}

At each step, either the protocol terminates (cases 1 and 3) or
$V$ strictly decreases by 1 (case 2).  Since $V \geq 0$ by
definition, case 2 can occur at most $V(c_0) \leq h_{\max}$ times.
Therefore, the protocol terminates in at most $h_{\max} + 1$ steps:
at most $h_{\max}$ hops (case 2) followed by exactly one terminating
action (case 1 or case 3), which is not itself a hop.
\end{proof}

\begin{remark}[Hops vs.\ protocol steps]
\label{rem:hops-vs-steps}
We are careful throughout the paper to distinguish the \emph{hop count}
of a chunk (the number of case-2 events, bounded by $h_{\max} =
\lfloor N/2\rfloor$, and the quantity that carries the
$\delta_{\mathrm{hop}}$ latency cost used in
\hyperref[chap:competitive_ratio]{the competitive-ratio analysis} and
elsewhere) from the \emph{protocol step count} of
Theorem~\ref{thm:livelock} and Corollary~\ref{cor:nodeflectloop} (hops
\emph{plus} the one terminating success/park action, bounded by
$h_{\max}+1$). Every downstream use of ``$\leq \lfloor N/2\rfloor$
hops'' or ``$\lfloor N/2\rfloor\,\delta_{\mathrm{hop}}$ overhead''
refers to the hop count, which is exactly $h_{\max}$, not $h_{\max}+1$;
the terminating action's cost (a park or a successful local push) is
accounted for separately wherever those costs are charged.
\end{remark}

\begin{corollary}[No indefinite looping in route\_deflect]
\label{cor:nodeflectloop}
The function \code{push\_chunk\_with\_hop} contains a \code{loop}
without an explicit upper-iteration bound.  Theorem~\ref{thm:livelock}
guarantees this loop terminates in at most
$h_{\max} + 1 = \floor{N/2} + 1$ iterations under any execution.
\end{corollary}

\begin{corollary}[Global livelock freedom]
\label{cor:global}
The DTA scheduler does not livelock (Definition~\ref{def:livelock}):
for every task $\tau$ and every valid execution trace, the chunk
carrying $\tau$ reaches one of the absorbing terminal states $Q_i$,
$W$, or $\varnothing$ in finite time (as measured in scheduler
operation steps).
\end{corollary}

\begin{proof}
Fix an arbitrary task $\tau$ and let $c$ be the chunk carrying it at
any point in the execution. Theorem~\ref{thm:livelock} applies to $c$
regardless of the state of any other chunk: its proof uses only $c$'s
own hop counter $h$ and the deterministic step count $h_{\max}-h+1$,
and never references the state of a different chunk $c' \neq c$ ---
each chunk's deflection sequence is driven purely by its own
per-chunk hop counter and the (bounded, but chunk-local) sequence of
mailbox-full/mailbox-available outcomes it individually encounters, so
distinct chunks' termination arguments do not interfere with one
another. Hence Theorem~\ref{thm:livelock} bounds $c$'s termination time
irrespective of what any other chunk in the system is doing, and since
$\tau$ and the execution trace were arbitrary, no task's chunk can
circulate indefinitely: this is precisely condition~(2) of
Definition~\ref{def:livelock} failing for every $\tau$, so no livelock
trace exists.
\end{proof}

% =============================================================
\section{Physical Analogy: The Lyapunov Interpretation}
% =============================================================

\begin{remark}[Connection to Lyapunov stability theory]
The argument of Theorem~\ref{thm:livelock} is the discrete analogue
of a standard Lyapunov stability argument in dynamical systems
theory.  Recall: a function $V : X \to \mathbb{R}_{\geq 0}$ is a
\emph{Lyapunov function} for a discrete-time system if
$V(x_{t+1}) < V(x_t)$ for all non-terminal states $x_t$.
The existence of such a function implies that the system cannot
cycle and must reach a terminal state in finite time.

Here, $V(c) = h_{\max} - h$ is exactly such a function for the
deflection sub-system:
\begin{itemize}
  \item $V \geq 0$ (bounded below),
  \item $V$ decreases by exactly 1 at each hop (strictly monotone),
  \item $V = 0$ forces termination (absorbing boundary).
\end{itemize}
The hop counter $h$ plays the role of an increasing \emph{action
variable}, and $V$ is its conjugate \emph{Lyapunov potential}.

In the language of statistical physics, the deflection protocol is
a \emph{finite-time dissipative process}: the potential $V$ is
``dissipated'' at a unit rate per hop, and the system must reach its
ground state $V=0$ --- the warehouse --- in finite time.  Unlike a
conservative Hamiltonian system, no energy is pumped back into the
potential, so recurrence is impossible.
\end{remark}

% =============================================================
\section{Scope and Limitations}
% =============================================================

Theorem~\ref{thm:livelock} proves that the \emph{deflection protocol
for a single chunk} terminates.  It does not prove:

\begin{itemize}
  \item That \emph{every task eventually executes} --- this is a
        \emph{starvation-freedom} or \emph{progress} property that
        requires additionally showing that tasks in $W$ or $Q_i$
        are eventually dispatched.  This is addressed in the
        \hyperref[chap:bounded_wait]{Progress~I}.
  \item That the system terminates for a batch of $T$ tasks in any
        bounded wall-clock time --- this requires a makespan bound and
        is addressed in \hyperref[chap:makespan]{Performance~III}.
\end{itemize}


\clearpage

% ========================================================================
% Correctness IV --- Atomicity and Linearizability
% ========================================================================

\section*{Correctness IV --- Atomicity and Linearizability}
\addcontentsline{toc}{section}{Correctness IV --- Atomicity and Linearizability}
\label{chap:linearizability}

\setcounter{section}{0}


% =============================================================
\section{Scope and Citation Strategy}
% =============================================================

Linearizability of concurrent data structures is a well-established
theory; we cite the definitive results and restrict our contribution
to three tasks:
\begin{enumerate}
  \item Identifying the \emph{linearization point} of each operation
        in the specific code (\S\ref{sec:spsc}--\ref{sec:mpmc}).
  \item Verifying that the \emph{memory-order annotations} in the
        Rust source satisfy the preconditions assumed by those results
        (\S\ref{sec:memorder}).
  \item Connecting the argument to the causal-order language used in
        the formal model (\S\ref{sec:causal}).
\end{enumerate}

\medskip\noindent\textbf{Primary references.}
\begin{itemize}
  \item Herlihy \& Wing (1990) --- canonical definition of
        linearizability~\cite{herlihy_wing1990}.
  \item Boehm \& Adve (2008) --- the C++ memory model; defines the
        happens-before relation $\xrightarrow{hb}$ for
        Release/Acquire pairs~\cite{boehm2008}.
  \item Vyukov (2010) --- the bounded MPMC ring-buffer algorithm whose
        per-slot sequence-number protocol is used verbatim in the
        warehouse~\cite{vyukov2010}.
  \item Haas et al.\ (2015) --- formal linearizability proof of the
        Vyukov queue~\cite{haas2015}.
\end{itemize}

% =============================================================
\section{Background: Linearizability and Linearization Points}
% =============================================================

\begin{definition}[Linearizability \cite{herlihy_wing1990}]
A concurrent execution history $H$ is \emph{linearizable} with
respect to a sequential specification $S$ if there exists a
\emph{linearization} --- a total order $\prec$ on the completed
operations in $H$ --- such that
\begin{enumerate}
  \item $\prec$ is consistent with the real-time order of
        non-overlapping operations in $H$, and
  \item each operation's result in $H$ matches its result under $S$
        when operations are executed in the order $\prec$.
\end{enumerate}
A sufficient condition (used in all proofs below) is the existence
of a \emph{linearization point} $\LP{op}$ for each operation $op$:
a single atomic step \emph{during} $op$'s execution such that the
operation appears to take effect instantaneously at $\LP{op}$.
\end{definition}

\begin{definition}[Happens-before \cite{boehm2008}]
\label{def:hb}
An operation $A$ \emph{happens-before} $B$ (written
$A \xrightarrow{hb} B$) if there exists a chain of
synchronization edges connecting them.  In the C++/Rust memory
model, a \emph{Release store} to an atomic variable $x$ at location
$\ell$ happens-before every subsequent \emph{Acquire load} of $x$
from $\ell$ that observes the stored value.  Every read performed
after an Acquire load is guaranteed to observe all writes that
happened-before the corresponding Release store.
\end{definition}

\begin{remark}[Physical interpretation]
The happens-before relation $\xrightarrow{hb}$ is the
discrete-memory-model analogue of the \emph{causal (light-cone)
order} in special relativity: two events are causally related
iff a signal (here: an observable value written then read) connects
them.  Events connected by $\xrightarrow{hb}$ have a
definite before/after order; events \emph{not} connected by any
$\xrightarrow{hb}$ chain are causally unrelated (``space-like
separated'') and their relative order is unobservable.
Linearizability is the statement that, despite the existence of
space-like-separated pairs of events, we can always construct a
consistent global time-ordering --- a ``foliation'' of the event
set --- that agrees with the causal order wherever the latter is
defined.
\end{remark}

% =============================================================
\section{SPSC Mailbox}
\label{sec:spsc}
% =============================================================

The mailbox (\code{struct Mailbox}) is a bounded single-producer
single-consumer (SPSC) ring buffer.  The canonical linearizability
result for SPSC queues is folklore~\cite{treiber1986} (see also \cite{boehm2008}
\S{}6 or any concurrent-programming textbook); we state it for
our specific implementation.

\subsection{Data structure}

The mailbox has two atomic index fields, \code{head} and
\code{tail}, each on its own 64-byte cache line to prevent false
sharing.  The invariant is: slots $[\code{head}, \code{tail})$
(mod $C_M$) contain valid chunks; the remaining slots are free.
Only the producer writes \code{tail}; only the consumer writes
\code{head}.

\subsection{Linearization points}

\begin{theorem}[SPSC mailbox is linearizable]
\label{thm:spsc}
Each \code{Mailbox::push} and \code{Mailbox::pop} operation is
linearizable.  The linearization points are:
\[
  \LP{\code{push}(c)}    \;=\; \code{tail.store(next\_tail,\,Release)}
\]
\[
  \LP{\code{pop}()}      \;=\; \code{head.store(next\_head,\,Release)}
\]
\end{theorem}

\begin{proof}[Proof sketch]
\textbf{Push.}  The producer reads \code{tail} (Relaxed --- local
variable, no other writer), computes \code{next\_tail}, checks
fullness via \code{head.load(Acquire)}, writes the chunk into the
buffer slot, then stores \code{next\_tail} to \code{tail} with
Release ordering.  The Release store is atomic; before it executes,
the chunk is written into the buffer (sequenced-before in the
producer's thread).  After the Release store, any consumer
performing \code{tail.load(Acquire)} will observe \code{next\_tail}
and, by the Release/Acquire synchronization edge
($\xrightarrow{hb}$, Definition~\ref{def:hb}), will also observe
the chunk write.  Thus the chunk becomes visible to the consumer
\emph{atomically at} $\LP{\code{push}}$.

\textbf{Pop.}  Symmetric: the consumer reads the chunk, then
advances \code{head} with Release ordering.  The producer's
subsequent \code{head.load(Acquire)} observes the advance and sees
the freed slot.  The pop takes effect atomically at
$\LP{\code{pop}}$.

\textbf{Non-overlap.}  Because there is exactly one producer and
one consumer, the producer's Release on \code{tail} and the
consumer's Acquire on \code{tail} form a single synchronization
edge --- no concurrent producers race on \code{tail}, and no
concurrent consumers race on \code{head}.  The total order on
operations is therefore unique and consistent with the
FIFO sequential specification.
\end{proof}

% =============================================================
\section{MPMC Warehouse (Vyukov Queue)}
\label{sec:mpmc}
% =============================================================

The warehouse (\code{struct Warehouse}) implements the bounded MPMC
queue due to Vyukov~\cite{vyukov2010}, using per-slot sequence
numbers to coordinate multiple concurrent producers and consumers.

\subsection{Algorithm recap}

Each slot $s$ has an atomic sequence number \code{seq}.
The protocol (reproduced from the source):
\begin{itemize}
  \item A producer at position $p$ may write slot $p \bmod C_W$
        iff \code{seq}$= p$ (slot is free for this round).
        It claims the slot via CAS on \code{tail} ($p \to p+1$),
        writes the payload, then publishes via
        \code{seq.store($p+1$, Release)}.
  \item A consumer at position $p$ may read slot $p \bmod C_W$
        iff \code{seq}$= p+1$ (slot has been published).
        It claims via CAS on \code{head} ($p \to p+1$),
        reads the payload, then releases the slot via
        \code{seq.store($p + C_W$, Release)}.
\end{itemize}

\subsection{Linearization points}

\begin{theorem}[Vyukov MPMC warehouse is linearizable \cite{haas2015}]
\label{thm:mpmc}
Each \code{Warehouse::push} and \code{Warehouse::pop} operation is
linearizable.  The linearization points are:
\[
  \LP{\code{push}(c)}    \;=\;
    \code{slot.seq.store($p$+1,\,Release)}
    \quad\text{(after successful CAS on \code{tail})}
\]
\[
  \LP{\code{pop}()}      \;=\;
    \code{slot.seq.store($p$+$C_W$,\,Release)}
    \quad\text{(after successful CAS on \code{head})}
\]
\end{theorem}

\begin{proof}[Proof by reference and parameter check]
Haas et al.~\cite{haas2015} give a full linearizability proof of the
Vyukov queue under the assumption that sequence-number updates use
Release/Acquire ordering and that the CAS operations on \code{tail}
and \code{head} use at minimum Relaxed ordering for the success case.
We verify that the DTA source satisfies these preconditions:

\begin{center}
\small
\begin{tabularx}{\textwidth}{@{} c X l @{}}
\toprule
Operation & Memory order in source & Required by \cite{haas2015} \\
\midrule
\texttt{tail} CAS (success)   & \texttt{Relaxed}/\texttt{Relaxed} & $\geq$ Relaxed \checkmark \\
\texttt{seq.load} (probe)     & \texttt{Acquire}                & Acquire \checkmark \\
\texttt{seq.store} (publish)  & \texttt{Release}                & Release \checkmark \\
\texttt{head} CAS (success)   & \texttt{Relaxed}/\texttt{Relaxed} & $\geq$ Relaxed \checkmark \\
\texttt{seq.store} (reclaim)  & \texttt{Release}                & Release \checkmark \\
\bottomrule
\end{tabularx}
\end{center}

All preconditions are satisfied.  The result of \cite{haas2015}
therefore applies directly, establishing linearizability of
\code{Warehouse::push} and \code{Warehouse::pop} with the
linearization points above.
\end{proof}

% =============================================================
\section{Memory-Order Annotation Audit}
\label{sec:memorder}
% =============================================================

For completeness, we audit every atomic operation in the two
structures against the C++ memory model's ordering hierarchy
(Relaxed $\subset$ Acquire/Release $\subset$ AcqRel $\subset$
SeqCst).  An annotation is \emph{correct} if it provides the
minimum ordering required for the property proved and does not
weaken any synchronization edge relied upon by
Theorems~\ref{thm:spsc}--\ref{thm:mpmc}.

\begin{table}[h]
\centering\small
\caption{Memory-order audit for SPSC Mailbox and MPMC Warehouse.
``Min.\ required'' is the weakest ordering that preserves
linearizability; ``Source'' is the actual annotation.}
\label{tab:memorder}
\begin{tabularx}{\textwidth}{@{}llXllc@{}}
\toprule
Structure & Operation & Variable & Source ordering & Min.\ required & OK? \\
\midrule
\multirow{4}{*}{Mailbox}
  & push: read \code{tail}   & \code{tail}  & Relaxed  & Relaxed  & \checkmark \\
  & push: read \code{head}   & \code{head}  & Acquire  & Acquire  & \checkmark \\
  & push: write \code{tail}  & \code{tail}  & Release  & Release  & \checkmark \\
  & pop:  read \code{tail}   & \code{tail}  & Acquire  & Acquire  & \checkmark \\
  & pop:  write \code{head}  & \code{head}  & Release  & Release  & \checkmark \\
\midrule
\multirow{9}{*}{Warehouse}
  & push: load \code{tail}   & \code{tail}    & Relaxed     & Relaxed  & \checkmark \\
  & push: CAS \code{tail}    & \code{tail}    & Relaxed/Relaxed & Relaxed/Relaxed & \checkmark \\
  & push: load \code{seq}    & \code{seq}     & Acquire     & Acquire  & \checkmark \\
  & push: store \code{seq}   & \code{seq}     & Release     & Release  & \checkmark \\
  & push: \code{fetch\_add}  & \code{backlog} & Release     & Relaxed$^\dagger$  & \checkmark \\
  & pop:  load \code{head}   & \code{head}    & Relaxed     & Relaxed  & \checkmark \\
  & pop:  CAS \code{head}    & \code{head}    & Relaxed/Relaxed & Relaxed/Relaxed & \checkmark \\
  & pop:  load \code{seq}    & \code{seq}     & Acquire     & Acquire  & \checkmark \\
  & pop:  store \code{seq}   & \code{seq}     & Release     & Release  & \checkmark \\
  & pop:  \code{fetch\_sub}  & \code{backlog} & Release     & Relaxed$^\dagger$  & \checkmark \\
  & \code{is\_busy}: load    & \code{backlog} & Relaxed     & Relaxed  & \checkmark \\
\bottomrule
\end{tabularx}
\end{table}

The claim of \emph{correctness} above depends on precisely
identifying where the linearization point of each Warehouse
operation actually falls, so we spell this out explicitly rather
than leaving it implicit in the ``OK?'' column.

The ring-index CAS on \code{tail} (push) and \code{head} (pop) is
annotated \emph{Relaxed/Relaxed} in the implementation, not
Release/Release: the CAS itself only arbitrates \emph{which}
producer/consumer wins a given slot index and carries no payload.
The actual publication of the payload --- the event with which a
concurrent \code{pop} must synchronize --- is the subsequent,
separate \code{Release} store to the per-slot \code{seq} field
(line~503/570 of \code{dta\_scheduler.rs}), which is read with
\code{Acquire} at the top of the opposing operation's loop
(line~476/547). This Acquire/Release pair on \code{seq}, not the
index CAS, is therefore the true linearization point used in the
proofs of Theorem~\ref{thm:mpmc} and \S\ref{sec:causal}; the index
CAS may safely be Relaxed/Relaxed because correctness never relies
on it establishing a happens-before edge on its own.

$^\dagger$The \code{backlog} counter is a liveness/scheduling hint
(read by \code{is\_busy} to decide whether a worker should drain the
Warehouse), not a correctness-critical synchronization variable: no
proof in this paper relies on an ordering guarantee for
\code{backlog} beyond eventual consistency with the CAS-protected
\code{head}/\code{tail} indices. The \code{Release} used on the
producer/consumer side is stronger than the Relaxed minimum required
and is retained only because it costs nothing on the store side; the
consumer-side \code{is\_busy} read is correctly Relaxed.

With this correction, all annotations are at least as strong as the
minimum required, and none is unnecessarily strong in a way that
would impose unintended sequentially-consistent ordering.

% =============================================================
\section{Connection to the Formal Model}
\label{sec:causal}
% =============================================================

Linearizability of the SPSC mailbox and MPMC warehouse ensures that
each atomic transfer operator $\hat{T}_{\alpha \to \beta}(\tau)$
(Definition 4.1 of \hyperref[chap:formal_model]{the Formal Model}) is well-defined as an
\emph{instantaneous} event in the execution history: the
linearization point of each queue operation is the unique instant at
which $\tau$ moves from site $\alpha$ to site $\beta$.

This justifies the occupation-number state formalism of the formal
model: the system state $\sigma(t)$ is well-defined at every instant
(not just between operations), because each concurrent operation
appears to take effect at a single atomic point.  Without
linearizability, the occupation numbers $n_{\alpha,\tau}$ could be
in an indeterminate superposition of 0 and 1 during a concurrent
push/pop --- an ill-defined state that would break the foundation of
the no-task-loss proof.

\begin{remark}[Loom coverage]
The concurrent state machine encoded in \code{loom} tests exercises
the SPSC mailbox and MPMC warehouse under all valid interleavings
explored by loom's exhaustive scheduler, providing mechanical
corroboration of Theorems~\ref{thm:spsc}--\ref{thm:mpmc} within the
preemption bound configured for those tests.
\end{remark}

% =============================================================
%  References (inline for standalone compilation)
% =============================================================



\bigskip
\noindent\textbf{Summary.}
The four correctness theorems collectively guarantee that DTA is a
safe concurrent scheduler: tasks are never created or destroyed by
routing, no thread can block indefinitely on a lock it will never acquire,
no deflection chain cycles, and every queue access is observably equivalent
to a single atomic operation.  These properties are independent of load
and hold in the worst case, forming a correctness floor on which the
probabilistic performance guarantees of Part~IV rest. Notably, none of
the four proofs required strengthening a worker's information set beyond
its own local state and single-hop mailbox traffic: bounded information
is a safe restriction, not merely a fast one.


\clearpage

\partbreak
\part{Progress}
\clearpage

\bigskip
\noindent\textbf{Overview.}
This part proves two progress properties: every task is eventually
scheduled (starvation freedom), and the waiting time before a task
begins execution is bounded by an explicit, purely structural function
of the system's capacity parameters ($N$, $L$, $C_W$, $K$, $\delta$) ---
a worst-case bound that requires no distributional assumption on
$\lambda$ or $\mu$, holding for every arrival pattern subject to one
pathwise rate condition on the warehouse (Assumption~\ref{ass:wh-net-drain}:
the warehouse's net inflow during any busy interval must not exceed its
collective drain rate, needed so that a worker's mailbox polling cannot
be delayed indefinitely by sustained peer overflow). Both bounds are
derived from the local dispatch-loop structure and the deterministic
polling order, without any assumption that a worker knows the state of
its peers --- a direct consequence of the bounded-information design.
The bounded-wait theorem
also serves as the starting point for the makespan analysis in Part~IV.


% ========================================================================
% Progress I --- Bounded Waiting Time and $k$-Bounded Fairness
% ========================================================================

\section*{Progress I --- Bounded Waiting Time and $k$-Bounded Fairness}
\addcontentsline{toc}{section}{Progress I --- Bounded Waiting Time and $k$-Bounded Fairness}
\label{chap:bounded_wait}
\setcounter{section}{0}


% =============================================================
\section{Setting and Assumptions}
% =============================================================

\begin{assumption}[Finite task execution time]
\label{ass:smax}
Every task $\tau \in \mathcal{T}$ completes execution in at most
$s_{\max} < \infty$ time units once a worker begins executing it.
\end{assumption}

\begin{assumption}[Finite context-switch overhead]
\label{ass:c}
The overhead of one context switch (save/restore register set,
update fiber state, prefetch) is bounded by a constant $c < \infty$.
\end{assumption}

\begin{assumption}[Workers are not stalled]
\label{ass:live}
No worker thread is preempted or blocked by the OS for more than
a constant interval $t_{\text{OS}} < \infty$ during any finite
window of interest.
\end{assumption}

\begin{definition}[Processing quantum]
\label{def:delta}
Define $\delta \coloneqq s_{\max} + c$.  This is the worst-case
time for a worker to process one task: execute it to completion
and perform the associated context-switch bookkeeping.
\end{definition}

\begin{remark}[$\delta$ is reused with a related but distinct meaning
from Performance~I onward]
\label{rem:delta-overload}
Here (and throughout Progress~I--II) $\delta$ is a deterministic
\emph{worst-case} bound, used only in distribution-free guarantees.
Starting in \hyperref[chap:load_balance]{Performance~I}, the same symbol
$\delta$ is reused for the queueing-theoretic \emph{mean} service time of
the M/M/1/$H$ model, via $\mu \coloneqq 1/\delta$ (service times
themselves i.i.d.\ exponential, not deterministic, under
Assumption~\ref{ass:poisson}); and in
\hyperref[chap:makespan]{Performance~III}'s steady-state makespan
definition, $\delta$ is further treated as a \emph{deterministic unit
service time per task} for the purpose of computing
$C_{\max} = \max_i q_i^* \cdot \delta$. These three uses coincide
\emph{exactly} only in the (unrealistic) special case of deterministic,
constant-time task execution; in general the worst-case
$\delta = s_{\max}+c$ of this definition upper-bounds the mean service
time $1/\mu$ used from Performance~I onward. We do not attempt to
track the resulting approximation error explicitly (it is subsumed by
the mean-field/M/M/1/$H$ approximation itself, whose validity is
discussed in \S\ref{chap:kinetic_theory}); readers should treat
$\delta$ in the performance chapters as the \emph{nominal} per-task time
$1/\mu$, not the worst-case quantum of this section, unless a bound is
explicitly labelled deterministic/worst-case.
\end{remark}

\begin{definition}[Waiting time]
\label{def:wait}
The \emph{waiting time} $W(\tau)$ of a task $\tau$ is the elapsed
time from its admission into the scheduler (execution of
$\hat{O}_{\text{ext}}$ or an internal spawn) to the moment worker
$i$ begins executing it ($\hat{O}_{\text{dispatch}}$ applied to
$\tau$).
\end{definition}

% =============================================================
\section{Anatomy of the Worker Main Loop}
% =============================================================

The structure of \code{run\_worker\_static} (the worker's main loop)
determines the time budget between any two consecutive
\code{poll\_mailboxes} calls.  We record it here for reference.

\begin{definition}[Worker iteration]
\label{def:iteration}
One \emph{iteration} of worker $i$'s main loop consists of:
\begin{enumerate}
  \item \textbf{Warehouse check.}  Load \code{warehouse.backlog}
        (Relaxed, single atomic read).  Let $b \coloneqq 1$ if the
        warehouse is busy, $0$ otherwise.
  \item \textbf{Drain (if $b=1$).}  Call \code{drain\_warehouse},
        which pops at most $D = 64$ chunks and pushes their tasks
        into $Q_i$ via \code{push\_batch}.
  \item \textbf{Dispatch.}  Call \code{dispatch\_loop}, which runs
        every task currently in $Q_i$ to completion (or until a
        Notified re-enqueue).
  \item \textbf{Poll mailboxes (if $b=0$).}  Call
        \code{poll\_mailboxes}, which scans all $N-1$ peer mailboxes
        in \code{polling\_order} and the external mailbox, accepting
        chunks up to the high-water mark $H$.
  \item \textbf{Adaptive idle.}  If no activity was detected in
        steps 2--4, enter a tiered backoff (spin $\to$ pause $\to$
        WFE/umonitor).  This adds at most $t_{\text{idle}}$ before
        the next iteration; in the worst case for our bounds,
        $t_{\text{idle}} = 0$ (worker spins without any
        time-consuming idle).
\end{enumerate}
\end{definition}

% =============================================================
\section{Building Blocks}
% =============================================================

\begin{lemma}[Dispatch duration]
\label{lem:dispatch}
One call to \code{dispatch\_loop} on worker $i$ takes at most
$L\,\delta$ time, where $L$ is the local queue capacity.
\end{lemma}

\begin{proof}
\code{dispatch\_loop} processes tasks sequentially until $Q_i$ is
empty.  Since $|Q_i| \leq L$ at all times, and each task takes at
most $\delta$ to process, the total time is $\leq L\delta$.
\end{proof}

\begin{lemma}[Drain duration]
\label{lem:drain}
One call to \code{drain\_warehouse} on worker $i$ adds at most
$DK$ tasks to $Q_i$ (where $D=64$ chunks, $K=32$ tasks/chunk)
and takes at most $DK\,\delta$ additional dispatch time in the
subsequent \code{dispatch\_loop} call.
\end{lemma}

\begin{proof}
\code{drain\_warehouse} pops at most $D$ chunks (code cap:
\code{let cap = 64}), each of size $\leq K$, stopping early if
$|Q_i| + K > H$.  The tasks are pushed via \code{push\_batch}
and dispatched in the same iteration's \code{dispatch\_loop}.
Total processing time for the drained batch: $\leq DK\delta$.
\end{proof}

\begin{assumption}[Warehouse net-drain condition]
\label{ass:wh-net-drain}
During any maximal interval in which the warehouse is continuously
observed busy (i.e., \code{warehouse.is\_busy()} holds at the start of
every worker iteration throughout the interval), the aggregate rate at
which \emph{other} workers push new chunks into the warehouse does not
exceed the collective drain rate $N/(K\delta)$ chunks per unit time.
This is a pathwise, worst-case rate cap on the arrival pattern (not a
distributional/steady-state average): it holds automatically whenever
the mean-field warehouse-stability condition
$\Lambda_W < \Lambda_D$ of Theorem~\ref{thm:wh-stable}
(\hyperref[chap:stability]{Performance~IV}) holds, but is stated here
directly as a worst-case condition so that this chapter's bound remains
distribution-free. Without some such condition, sustained concurrent
overflow from other workers could in principle keep a worker's
observed backlog nonzero indefinitely, and since
Definition~\ref{def:iteration} gates \code{poll\_mailboxes} on
\code{warehouse.is\_busy()}$=$false, that worker would never resume
polling --- so the bound below is conditional on
Assumption~\ref{ass:wh-net-drain}, not unconditional over every
conceivable arrival pattern.
\end{assumption}

\begin{lemma}[Warehouse drain time (collective)]
\label{lem:wdrain}
Under Assumption~\ref{ass:wh-net-drain}, if the warehouse contains at
most $C_W \cdot K$ tasks at the start of a busy interval, all $N$
workers collectively empty it within that interval in at most
\[
  T_{\mathrm{drain}} \;=\; \frac{C_W \cdot K}{N}\,\delta
\]
time, assuming each worker drains and dispatches at full speed with
no idle intervals.
\end{lemma}

\begin{proof}
The warehouse holds at most $C_W$ chunks at the start of the interval,
and by Assumption~\ref{ass:wh-net-drain} the net inflow during the
interval is non-positive relative to the collective drain rate. Each
worker drains $D$ chunks per iteration of the drain phase, spending
$DK\delta$ dispatching them.  The total chunk-drain capacity of all $N$
workers per unit time is $N \cdot D / (DK\delta) = N/(K\delta)$
chunks per time unit, so the time to drain (at most) $C_W$ net chunks
is $C_W K\delta / N$.  (The factor $K$ accounts for the dispatch
time of the $K$ tasks per chunk.)
\end{proof}

% =============================================================
\section{Waiting Time: Case Analysis}
% =============================================================

A task $\tau$ admitted to the scheduler enters one of two
initial sites (see the per-task FSM in \hyperref[chap:formal_model]{the Formal Model}):
\begin{itemize}
  \item A mailbox $M_{\text{ext},i}$ (external injection) or
        $M_{ij}$ (intra-worker deflection), or
  \item The warehouse $W$ (direct warehouse diversion when
        \code{warehouse.is\_busy()}).
\end{itemize}

\begin{figure}[h]
\centering
\begin{tikzpicture}[
  box/.style={draw, rounded corners, minimum width=3.2cm,
              minimum height=0.9cm, align=center, font=\small, inner sep=4pt},
  arr/.style={-{Latex[length=2.5mm]}, line width=0.85pt},
  lbl/.style={font=\scriptsize, fill=white, inner sep=1.5pt},
]
  \node[box, fill=gray!10]   (ext) at (0,5)    {Admitted\\ $M$ or $W$};
  \node[box, fill=yellow!18] (M)   at (-3,3)   {Mailbox $M_{ji}$\\wait $\leq W_M$};
  \node[box, fill=orange!18] (W)   at (3,3)    {Warehouse $W$\\wait $\leq W_W$};
  \node[box, fill=green!15]  (Q)   at (0,1)    {Local queue $Q_i$\\wait $\leq W_Q$};
  \node[box, fill=blue!12]   (E)   at (0,-1)   {Executing $E_i$\\(start of execution)};

  \draw[arr] (ext) -- node[lbl,left]  {mailbox path} (M);
  \draw[arr] (ext) -- node[lbl,right] {warehouse path} (W);
  \draw[arr] (M)   -- node[lbl,left]  {poll\_mailboxes} (Q);
  \draw[arr] (W)   -- node[lbl,right] {drain\_warehouse} (Q);
  \draw[arr] (Q)   -- node[lbl,right] {dispatch\_loop} (E);
\end{tikzpicture}
\caption{Waiting-time case structure.  $W_M$, $W_W$, and $W_Q$ are
the per-stage worst-case waiting times derived in \S\ref{sec:cases}.}
\label{fig:wait}
\end{figure}

\subsection{Per-stage waiting times}
\label{sec:cases}

\paragraph{Stage Q: waiting in the local queue.}

When $\tau$ enters $Q_i$, it is placed at the tail.  The
\code{dispatch\_loop} processes tasks in FIFO order.  The number of
tasks ahead of $\tau$ is at most $L-1$ (the queue was at most full
minus the newly added task).

\[
  W_Q \;\leq\; L\,\delta
\]

\paragraph{Stage M: waiting in a mailbox (normal operation).}

When the warehouse is not busy ($b=0$ in every iteration),
worker $i$ calls \code{poll\_mailboxes} after every
\code{dispatch\_loop}.  One \code{dispatch\_loop} takes at most
$L\delta$ (Lemma~\ref{lem:dispatch}).  Therefore the gap between
two consecutive polls of $M_{ji}$ is at most:

\[
  W_M^{(\text{normal})} \;\leq\; L\,\delta
\]

\paragraph{Stage M: waiting in a mailbox (warehouse active).}

When the warehouse is busy, worker $i$ skips \code{poll\_mailboxes}
for the duration of the drain phase.  Under
Assumption~\ref{ass:wh-net-drain}, Lemma~\ref{lem:wdrain} bounds this
drain phase at $T_{\mathrm{drain}} = C_W K\delta/N$, after which
\code{warehouse.is\_busy()} reports false and polling resumes on the
next iteration, so $\tau$ is picked up within the next poll cycle
($\leq L\delta$).

\[
  W_M \;\leq\; \frac{C_W K}{N}\,\delta \;+\; L\,\delta
\]

\paragraph{Stage W: waiting in the warehouse.}

$\tau$ enters $W$ at some position $p \leq C_W K - 1$ (behind at
most $C_W K - 1$ other tasks).  All $N$ workers simultaneously
drain the warehouse, collectively processing $NDK$ tasks per drain
round.  The time to drain past $\tau$'s position:

\[
  W_W \;\leq\; \ceil{\frac{C_W K}{NDK}} \cdot DK\delta
       \;=\; \ceil{\frac{C_W}{ND}} \cdot DK\delta
       \;\leq\; \left(\frac{C_W}{N} + D\right)\delta K
\]
using $\lceil x \rceil \le x+1$ termwise. When $ND$ divides $C_W$
exactly (as in every configuration considered in this paper, e.g.\ the
benchmark parameters of \S\ref{sec:validity}), the ceiling is exact and
this reduces to $W_W \leq \frac{C_WK}{N}\delta$, the value used in the
main theorem below; in general the extra additive slack is at most
$D\delta K$, which is dominated by the $L\delta$ term already present
in the theorem's bound whenever $D \leq L$ (true for all published
configurations, since $D \le C_W/N \ll L$).
After being drained, $\tau$ enters $Q_i$ and waits $W_Q$.

% =============================================================
\section{Main Theorem}
% =============================================================

\begin{theorem}[Bounded waiting time]
\label{thm:bounded-wait}
Under Assumptions~\ref{ass:smax}--\ref{ass:live} and~\ref{ass:wh-net-drain},
for any task $\tau$ admitted to the DTA scheduler:
\[
  \boxed{
    W(\tau) \;\leq\; \delta\!\left(2L + \frac{C_W K}{N}\right)
  }
\]
\end{theorem}

\begin{proof}
The worst case is the combination of all three stages:
\begin{align*}
  W(\tau)
  &\leq W_M + W_Q
  &&\text{(mailbox path, warehouse active)} \\
  &\leq \left(\frac{C_W K}{N}\,\delta + L\,\delta\right) + L\,\delta \\
  &= \delta\!\left(2L + \frac{C_W K}{N}\right).
\end{align*}
The warehouse path ($W_W + W_Q$) is subsumed:
$W_W + W_Q = C_W K\delta/N + L\delta < W_M + W_Q$.

The hop overhead from the livelock-freedom result
(\hyperref[chap:livelock]{Correctness~III}, Theorem~3.1)
contributes at most $h_{\max}\,\delta_{\mathrm{hop}}$ where
$\delta_{\mathrm{hop}} \ll \delta$ (one failed push attempt is a
handful of atomic operations, not a full task execution); this term
is dominated by $L\delta$ and absorbed into the bound.
\end{proof}

\begin{corollary}[Parametric bound for benchmark configuration]
With $N=4$, $L=131{,}072$, $C_W=32{,}768$, $K=32$:
\[
  W(\tau) \;\leq\; 524{,}288\,\delta
\]
This is a \emph{conservative} worst-case ceiling --- it requires the
warehouse to be full ($\approx 10^6$ queued tasks) while every
local queue is simultaneously at capacity.  Under typical benchmark
loads ($P = 8{,}192$ concurrent fibers), the warehouse is rarely
active and the practical wait is $O(P\delta/N) \ll L\delta$.
\end{corollary}

% =============================================================
\section{$k$-Bounded Fairness}
% =============================================================

\begin{definition}[$k$-bounded fairness]
\label{def:kfair}
A scheduler is \emph{$k$-bounded fair} if every task admitted at
time $t_0$ begins execution before the scheduler completes $k$
full dispatch rounds after $t_0$.  A \emph{dispatch round} for
worker $i$ is one call to \code{dispatch\_loop} (processing all
tasks currently in $Q_i$).
\end{definition}

\begin{theorem}[$k$-bounded fairness of DTA]
\label{thm:kfair}
The DTA scheduler is $k$-bounded fair with
\[
  k \;\leq\; 2 + \ceil{\frac{C_W K}{NL}}
\]
dispatch rounds per worker.
\end{theorem}

\begin{proof}
One dispatch round processes at most $L$ tasks and takes at most
$L\delta$ time.  The total waiting time $W(\tau) \leq \delta(2L +
C_W K/N)$ spans at most
\[
  \frac{W(\tau)}{L\delta}
  \;\leq\; \frac{2L + C_W K/N}{L}
  \;=\; 2 + \frac{C_W K}{NL}
\]
dispatch rounds.  Since $k$ must be an integer, we take the ceiling.
\end{proof}

\begin{corollary}[Benchmark $k$-value]
With the published benchmark parameters:
\[
  k \;\leq\; 2 + \ceil{\frac{32{,}768 \times 32}{4 \times 131{,}072}}
    \;=\; 2 + \ceil{2}
    \;=\; 4
\]
Every admitted task starts execution within 4 dispatch rounds of
the receiving worker.  Under no-warehouse-active conditions
($C_W K \ll NL$), this collapses to $k \leq 2$: at most one round
to pick the task from the mailbox, and at most one round to dispatch
it once in $Q_i$.
\end{corollary}

% =============================================================

\begin{figure}[htbp]
\centering
\begin{minipage}{0.48\linewidth}\centering
\includegraphics[width=\linewidth]{figure/wait_bound_vs_N.png}
\caption{Upper bound $W(\tau) \le \delta(2L + C_W K/N)$ as a function of $N$ (parameters: $L=H=8192$, $K=32$, $C_W=500$, $\delta=1\,\mu\mathrm{s}$).}\label{fig:wait_bound}
\end{minipage}\hfill
\begin{minipage}{0.48\linewidth}\centering
\includegraphics[width=\linewidth]{figure/k_bound_vs_N.png}
\caption{$k$-Bounded fairness parameter $k$ as a function of $N$.}\label{fig:k_bound}
\end{minipage}
\end{figure}
\section{Scaling Behavior and Use in Makespan Analysis}
% =============================================================

The bound $W(\tau) \leq \delta(2L + C_W K/N)$ has two structurally
distinct components:

\begin{enumerate}
  \item $2L\delta$: the \emph{queue-depth component}, independent
        of $N$.  This is the cost of waiting behind tasks already
        in the local queue when $\tau$ arrives.  It is the dominant
        term when $N$ is small.

  \item $C_W K\delta/N$: the \emph{warehouse component}, which
        decreases as $1/N$.  This reflects the benefit of parallel
        warehouse draining: doubling the number of workers halves
        the per-task warehouse wait.  It is the dominant term when
        $N \gg C_W K / (2L) = 128$ (impractical at current
        parameter settings).
\end{enumerate}

\medskip\noindent\textbf{For the makespan bound} (to be derived in
\hyperref[chap:makespan]{Performance~III}): the waiting time $W(\tau)$
represents the \emph{per-task scheduling overhead} beyond the useful
work $s_{\max}$.  For a batch of $T$ tasks with total work
$T_1 = T \cdot s_{\max}$ and critical-path length $T_\infty$, the
makespan $C_{\max}$ satisfies (informally):
\[
  C_{\max} \;\leq\; \frac{T_1}{N} + T_\infty + W_{\max}
\]
where $W_{\max} = \delta(2L + C_W K/N)$ is the per-task scheduling
overhead term derived here.  The precise statement requires a
task-DAG model and is the subject of the performance sub-module.


\clearpage

% ========================================================================
% Progress II --- Starvation Freedom and $k$-Bounded Fairness
% ========================================================================

\section*{Progress II --- Starvation Freedom and $k$-Bounded Fairness}
\addcontentsline{toc}{section}{Progress II --- Starvation Freedom and $k$-Bounded Fairness}
\label{chap:starvation}

\setcounter{section}{0}

This chapter imports the bounded waiting time bound from
\hyperref[chap:bounded_wait]{Progress~I} (Theorem~4.1):
\[
  W(\tau) \;\leq\; W_{\max}
  \;\coloneqq\; \delta\!\left(2L + \frac{C_W K}{N}\right),
\]
and the per-task FSM and notation from
\hyperref[chap:formal_model]{the Formal Model}.  Assumptions~A1--A3
of \hyperref[chap:bounded_wait]{Progress~I} (finite $s_{\max}$, finite $c$,
workers not indefinitely stalled) are in force throughout.

% =============================================================
\section{Definitions}
% =============================================================

We give precise definitions suited to the DTA execution model
before stating the theorems.  The key distinction from classical
scheduling theory is that DTA is a \emph{work-conserving},
\emph{deterministic-polling} scheduler rather than a preemptive
or randomized one; this makes the relevant notion of fairness
quantitative (a finite bound on waiting rounds) rather than
probabilistic (positive probability of being served in finite time).

\begin{definition}[Starvation]
\label{def:starve}
A task $\tau$ \emph{starves} in an execution trace
$\sigma_0, \sigma_1, \ldots$ if $\tau$ is admitted at some finite
step $t_0$ but never reaches the executing site $E_i$ for any
$i \in \mathcal{W}$ in any subsequent step.  That is,
$n_{E_i, \tau}(\sigma_t) = 0$ for all $i$ and all $t > t_0$.
\end{definition}

\begin{definition}[Starvation freedom]
\label{def:starve-free}
The scheduler is \emph{starvation-free} if no task $\tau$ starves:
for every admitted $\tau$ and every valid execution trace,
there exists a finite step $t^* > t_0$ such that
$n_{E_i, \tau}(\sigma_{t^*}) = 1$ for some $i \in \mathcal{W}$.
\end{definition}

\begin{definition}[$k$-Bounded fairness]

The scheduler is \emph{$k$-bounded fair} with parameter $k \in
\mathbb{N}_{>0}$ if every admitted task $\tau$ begins execution
within at most $k$ complete dispatch rounds of its receiving
worker after admission.  A \emph{dispatch round} of worker $i$
is one complete execution of \code{dispatch\_loop} that empties
$Q_i$.

Formally: if $\tau$ is admitted at wall-clock time $t_0$, then
$n_{E_i, \tau}(\sigma_{t^*}) = 1$ for some $t^*$ satisfying
\[
  t^* - t_0 \;\leq\; k \cdot L\delta,
\]
where $L\delta$ is the maximum duration of one dispatch round
(Lemma~3.1 of \hyperref[chap:bounded_wait]{Progress~I}).
\end{definition}

\begin{remark}[Relation between the two definitions]
$k$-bounded fairness is strictly stronger than starvation freedom:
it asserts not merely that $\tau$ \emph{eventually} executes, but
that it does so within an \emph{explicit, finite} number of dispatch
rounds.  Starvation freedom is the qualitative limit as
$k \to \infty$; $k$-bounded fairness is its quantitative
refinement.  In the DTA context, we derive the stronger
$k$-bounded result directly, so starvation freedom follows as a
special case.
\end{remark}

\begin{remark}[Why ``bounded'' rather than ``probabilistic'']
Classical analyses of randomized schedulers (e.g., randomized
work-stealing) typically establish starvation freedom via ergodicity:
a task is served with positive probability on each attempt, so it
is served almost surely in finite time.  DTA's \code{polling\_order}
is a \emph{deterministic} round-robin scan of all peer mailboxes
(same-CCX first, then cross-CCX, but both groups fully covered).
This is closer to a \emph{periodic sweep} than a random walk, and
yields a deterministic bound rather than a probabilistic one.
A deterministic bound is both stronger and more practically useful:
it gives a hard latency guarantee rather than an almost-sure one.
\end{remark}

% =============================================================
\section{Main Results}
% =============================================================

\begin{theorem}[Starvation freedom]
\label{thm:starve-free}
Under Assumptions~A1--A3 of \hyperref[chap:bounded_wait]{Progress~I}, the
DTA scheduler is starvation-free.
\end{theorem}

\begin{proof}
By Theorem~4.1 of \hyperref[chap:bounded_wait]{Progress~I}, every admitted task
$\tau$ satisfies $W(\tau) \leq W_{\max} < \infty$.  That is,
$\tau$ begins execution at some time $t^* \leq t_0 + W_{\max}$,
which is a finite time.  Therefore $n_{E_i, \tau}(\sigma_{t^*}) = 1$
for some finite $t^*$, and $\tau$ does not starve by
Definition~\ref{def:starve}.  Since $\tau$ was arbitrary, the
scheduler is starvation-free.
\end{proof}

\begin{theorem}[$k$-Bounded fairness of DTA]

The DTA scheduler is $k$-bounded fair with
\[
  k \;=\; \ceil{2 + \frac{C_W K}{NL}}.
\]
For the published benchmark parameters $(N=4,\,L=131{,}072,\,
C_W=32{,}768,\,K=32)$, this gives $k = 4$.
\end{theorem}

\begin{proof}
From Theorem~4.1 of \hyperref[chap:bounded_wait]{Progress~I},
$W(\tau) \leq \delta(2L + C_W K/N)$.
One dispatch round lasts at most $L\delta$ (Lemma~3.1 of
\hyperref[chap:bounded_wait]{Progress~I}).
The number of dispatch rounds elapsed before $\tau$ begins
execution is therefore at most
\[
  \frac{W(\tau)}{L\delta}
  \;\leq\;
  \frac{\delta(2L + C_W K/N)}{L\delta}
  \;=\; 2 + \frac{C_W K}{NL}.
\]
Taking the ceiling (since round counts are integers) yields
$k = \ceil{2 + C_W K/(NL)}$.  For the benchmark parameters:
$k = \ceil{2 + 32{,}768 \times 32 / (4 \times 131{,}072)} = \ceil{4} = 4$.
\end{proof}

\begin{corollary}[Asymptotic fairness in $N$]
\label{cor:asymptotic}
As $N \to \infty$ with $L$, $C_W$, $K$ fixed:
\[
  k \;\to\; \ceil{2 + 0} \;=\; 2.
\]
That is, in the limit of many workers the scheduler achieves
$2$-bounded fairness: every task waits at most 2 dispatch rounds
before execution begins.  This floor of $2$ reflects the two
unavoidable stages --- one round to poll the mailbox and move $\tau$
into $Q_i$, and one round to dispatch $\tau$ from $Q_i$ --- and
cannot be reduced to 1 without merging the two stages into a single
atomic step, which would require priority insertion into an
already-executing dispatch loop.
\end{corollary}

\begin{corollary}[CCX-aware polling does not break fairness]
\label{cor:ccx}
The \code{polling\_order} construction (Definition~3.4 of
\hyperref[chap:formal_model]{the Formal Model}) places same-CCX peers before
cross-CCX peers.  Cross-CCX mailboxes therefore appear later
in each polling pass, incurring a higher \emph{expected} wait
than same-CCX mailboxes, but they are \emph{always present} in
the list.  Every \code{poll\_mailboxes} call scans all $N-1$
peer mailboxes without exception.

Consequently, the $k$-bound of Theorem~\ref{thm:kfair} applies
uniformly to all tasks regardless of their source worker's CCX
affinity: cross-CCX tasks may experience higher \emph{average}
latency within a poll pass (they are polled last), but the
worst-case $k$-bound is unchanged.  The CCX ordering is a
\emph{latency-optimization}, not a \emph{starvation risk}.
\end{corollary}

% =============================================================
\section{What This Does and Does Not Establish}
% =============================================================

\begin{remark}[Scope]
Theorems~\ref{thm:starve-free}--\ref{thm:kfair} establish:
\begin{itemize}
  \item Every admitted task \emph{begins} execution within $k$
        dispatch rounds.
  \item This bound is \emph{deterministic}, not probabilistic.
  \item The bound holds for all task arrival patterns (worst-case
        over all schedules compatible with Assumptions~A1--A3).
\end{itemize}
They do \emph{not} establish:
\begin{itemize}
  \item That tasks complete in bounded time --- completion requires
        additionally that $s_{\max} < \infty$ (Assumption~A1),
        which is assumed but not a scheduler property.
  \item Load balance across workers --- the $k$-bound is a per-task
        latency guarantee, not a statement about the variance of
        $|Q_i|$ across workers.  Load balance is addressed in
        \hyperref[chap:load_balance]{Performance~I}.
  \item Fairness under adversarial task injection --- if tasks are
        injected at rate $\lambda > N\mu$ (arrival faster than
        service), the warehouse will eventually fill and the
        process will abort (crash-on-overflow).  The $k$-bound
        assumes the system operates within its stability region,
        whose characterisation is the subject of
        \hyperref[chap:stability]{Performance~IV}.
\end{itemize}
\end{remark}


\bigskip
\noindent\textbf{Summary.}
The bounded-wait theorem gives an explicit, distribution-free per-task
latency ceiling as a function of the capacity parameters $N$, $L$,
$C_W$, $K$, and $\delta$ alone --- Starvation Freedom (proved separately
in Progress~II) then follows as a corollary of this finite ceiling. Both
results are obtained without any worker consulting a peer's state, which
is exactly what makes them worst-case, distribution-free guarantees
rather than probabilistic ones.
The $k$-bounded fairness corollary tightens this to a round-robin-style
guarantee over any window of $k$ consecutive scheduling decisions.
Both results will be invoked in Performance~III to convert per-task
latency bounds into whole-batch makespan bounds.


\clearpage

\partbreak
\part{Performance Analysis}
\clearpage

\bigskip
\noindent\textbf{Overview.}
This part derives quantitative performance bounds across five chapters.
Performance~I develops the mean-field load-balance theory and the
M/M/1/$H$ self-consistency equation.
Performance~II provides the BBGKY kinetic-theory derivation of the
mean-field result, together with an explicit $O(1/N)$ pair-correlator correction.
Performance~III proves a two-line makespan bound combining deterministic
work-conservation with Gumbel extreme-value fluctuations.
Performance~IV establishes a four-level stability ladder
(throughput, NESS convergence, burst absorption, and crash-free threshold).
Performance~V extends the kinetic-theory framework to
DAG scheduling, deriving mean-field makespan bounds and second-order
Gumbel corrections for fork-join graphs.
Read together, these five chapters answer the quantitative half of the
paper's central question: routing decisions based only on local queue
occupancy achieve $O(1/\sqrt{N})$ load imbalance, which is the same
leading-order scaling obtained by schedulers that consult a wider state
window, so the bounded-information restriction costs nothing at the
mean-field level, plus an explicit $O(1/N)$ correction beyond it. This
$N \to \infty$ statement should not be over-read at small, fixed $N$:
at the $N=4$ benchmark scale used elsewhere in this paper, the
mean-field approximation itself carries a worst-case error up to
$O(33\%)$ (\S\ref{chap:kinetic_theory}), so ``costs nothing'' is an
asymptotic, not a small-$N$, guarantee.


% ========================================================================
% Performance I --- Load Balance
% ========================================================================

\section*{Performance I --- Load Balance}
\addcontentsline{toc}{section}{Performance I --- Load Balance}
\label{chap:load_balance}
\setcounter{section}{0}

\begin{assumption}[Poisson arrivals]
\label{ass:poisson}
Unlike the worst-case, distribution-free bounds of Parts~II and~IV
(\hyperref[chap:bounded_wait]{Progress~I}, \hyperref[chap:makespan]{Performance~III}
Line~1, and \hyperref[chap:competitive_ratio]{the competitive-ratio analysis}),
every quantitative result in this chapter and in
\hyperref[chap:kinetic_theory]{Performance~II} that references the
steady-state distribution $f_1^*$, the self-consistency equation (SC),
or any M/M/1/$H$ quantity assumes that task arrivals to the system
form a Poisson process of rate $\lambda_0$ per worker (equivalently,
inter-arrival times are i.i.d.\ exponential), independent of service
times, which are themselves i.i.d.\ exponential with rate $\mu = 1/\delta$.
This is the standard mean-field tractability assumption for queueing
network analysis and is what makes $f_1^*$ an exact truncated-geometric
distribution (Proposition~3.2). Results that do \emph{not} carry this
assumption explicitly (e.g.\ the deterministic bounds of Progress~I) hold
for \emph{any} arrival process, Poisson or not.
\end{assumption}

% =============================================================
\section{Why There Is No Boltzmann Distribution}
% =============================================================

Before deriving the stationary distribution, we establish why the
standard equilibrium-statistical-mechanics approach --- finding an
energy functional $\mathcal{H}$ and writing $\pi(\bm{q}) \propto
e^{-\beta\mathcal{H}(\bm{q})}$ --- does not apply.

\begin{definition}[Detailed balance]
\label{def:db}
A Markov chain with state space $\Sigma$ and transition rates
$W(\sigma \to \sigma')$ satisfies \emph{detailed balance} with
respect to a distribution $\pi$ if
\[
  \pi(\sigma)\,W(\sigma \to \sigma')
  \;=\;
  \pi(\sigma')\,W(\sigma' \to \sigma)
  \qquad \forall\,\sigma, \sigma' \in \Sigma.
\]
Detailed balance is the necessary and sufficient condition for
$\pi$ to be an equilibrium (Gibbs) measure of the form
$\pi(\sigma) \propto e^{-\beta\mathcal{H}(\sigma)}$ for some
functional $\mathcal{H}$.
\end{definition}

\begin{lemma}[DTA violates detailed balance]
\label{lem:no-db}
The joint queue-length process $(q_1(t), \ldots, q_N(t))$ of
the DTA scheduler does not satisfy detailed balance, and
therefore has no Boltzmann-form stationary distribution.
\end{lemma}

\begin{proof}
Consider the two-worker case ($N=2$) with states
$\bm{q} = (q_1, q_2)$.  The deflection transition
\[
  (H, q_2) \;\xrightarrow{\lambda_0}\; (H, q_2 + 1)
  \qquad (q_1 \text{ above threshold, deflects to } W_2)
\]
has a forward rate $W((H, q_2) \to (H, q_2+1)) = \lambda_0 > 0$.
The reverse transition
\[
  (H, q_2+1) \;\to\; (H, q_2)
\]
can only occur via $W_2$ dispatching a task: rate $\mu$, \emph{not}
via any reverse-deflection (tasks are never pulled back from $W_2$
to $W_1$).  The forward transition is sourced by $W_1$'s overload;
the reverse is sourced by $W_2$'s service.  These are
structurally different mechanisms with different rates, and there
is no value of $\pi$ that makes both sides of the detailed-balance
equation equal simultaneously for all $(q_1, q_2)$.
\end{proof}

\begin{remark}[Physical interpretation]
The DTA scheduler is a \emph{non-equilibrium steady state (NESS)}
system: it sustains a persistent probability current in state space
(tasks flow directionally from overloaded to underloaded workers)
even in the stationary regime.  This is the scheduling analogue of
a driven diffusion process, not a relaxation to thermal equilibrium.
The standard toolkit for such systems --- master equation,
mean-field theory, or large-deviation theory --- applies; the
Gibbs-ensemble toolkit does not.
\end{remark}

% =============================================================
\section{Mean-Field Reduction}
% =============================================================

\begin{assumption}[Exchangeability]
\label{ass:exch}
All workers are statistically equivalent: they share the same
capacity $L$, threshold $H$, service rate $\mu = 1/\delta$, and
base arrival rate $\lambda_0$.  The deflection target is chosen
uniformly at random among the $N-1$ peers (mean-field
approximation to the hash routing;
cf.~\cite{mitzenmacher2001,vvedenskaya1996} for the power-of-two-choices mean-field analysis).
\end{assumption}

\begin{proposition}[Hash routing is exact mean-field in distribution]
\label{prop:hash_mf}
Let the routing target be $\tau(i) = (h_1(\mathrm{flow\_id}) + h_2(\mathrm{flow\_id})) \bmod N$
where $h_1, h_2$ are independent hash functions.
If $\mathrm{flow\_id}$ is uniform on $\{0, 1, \ldots, 2^k - 1\}$ and $N \mid 2^k$,
then $\tau(i)$ is uniform on $\{0, 1, \ldots, N-1\}$ and independent across
distinct flow IDs.
Hence the mean-field Assumption~\ref{ass:exch} (uniform random peer) is
\emph{exact in distribution} for i.i.d.\ flow IDs, not merely an approximation.
\end{proposition}

\begin{proof}
For $N \mid 2^k$, the map $x \mapsto x \bmod N$ is a surjection from
$\{0,\ldots,2^k-1\}$ onto $\{0,\ldots,N-1\}$ with each pre-image of size $2^k/N$.
If $x$ is uniform on $\{0,\ldots,2^k-1\}$, then $x \bmod N$ is uniform on
$\{0,\ldots,N-1\}$.  Since $h_1(\mathrm{flow\_id})$ and $h_2(\mathrm{flow\_id})$
are computed from independent bit-ranges of \code{flow\_id}~\cite{dtact_source},
their sum $\bmod N$ is also uniform by closure of the uniform distribution
under convolution on $\mathbb{Z}/N\mathbb{Z}$.
Across distinct flow IDs, the outputs are independent by the pseudorandomness
of the hash functions.
\end{proof}

\begin{remark}[Finite-$N$ correction for non-uniform flow IDs]
When \code{flow\_id} is \emph{not} uniformly distributed (e.g., biased
toward certain connection classes), the routing distribution deviates from
uniform by $\varepsilon_{\mathrm{bias}} \le \mathrm{TV}(\mathrm{flow\_id}, \mathrm{Uniform})$.
The BBGKY Level-1 correction (\hyperref[chap:kinetic_theory]{Performance~II},
Section~3) gives an $O(1/N)$ pair-correlation bound that dominates
$\varepsilon_{\mathrm{bias}}$ for $N=4$ in the DTA benchmark.
For CCX-topology modes (same-CCX peers preferred), the mean-field result
gives the \emph{topology-unaware} baseline; CCX-aware corrections are
treated as a finite-$N$ perturbation of size $O(1/N)$.
\end{remark}

Under Assumption~\ref{ass:exch}, the marginal distribution of any
single worker's queue length $q_i$ converges (in the limit $N \to
\infty$, by the propagation-of-chaos theorem for exchangeable
particle systems~\cite{sznitman1991}) to a common distribution
$\pi$, independent of the other workers.  For finite $N$ (including
the benchmark $N=4$), we treat $\pi$ as an approximation and
quantify the error by simulation in Section~\ref{sec:sim}.

\subsection{Effective arrival rate and the self-consistency equation}

In the mean-field approximation, the queue of worker $i$ evolves
as a birth-death chain with:
\begin{itemize}
  \item \textbf{Birth rate} (task arrival) at state $q < H$:
        $\lambda_{\text{eff}}$ (to be determined self-consistently).
  \item \textbf{Birth rate} at state $q \geq H$: $0$
        (worker $i$ deflects all incoming tasks; they do not join
        $Q_i$).
  \item \textbf{Death rate} (task dispatch) at state $q > 0$:
        $\mu = 1/\delta$.
\end{itemize}

The effective arrival rate $\lambda_{\text{eff}}$ accounts for:
\begin{enumerate}
  \item \textbf{Direct spawns} to worker $i$: rate $\lambda_0$.
        These are accepted iff $q_i < H$.
  \item \textbf{Deflected tasks from other workers}: each of the
        $N-1$ peers deflects at rate $\lambda_0 p_d$ (where
        $p_d \coloneqq P(q \geq H)$ is the deflection
        probability), and sends $1/(N-1)$ of those deflections
        to worker $i$.  Worker $i$ accepts them iff $q_i < H$.
\end{enumerate}
Therefore, when $q_i < H$:
\[
  \lambda_{\text{eff}}
  \;=\; \lambda_0 \;+\; (N-1)\cdot\frac{\lambda_0 p_d}{N-1}
  \;=\; \lambda_0(1 + p_d).
\]

\begin{definition}[Traffic intensity and self-consistency equation]
\label{def:rho}
Define the \emph{effective traffic intensity}
$\rho \coloneqq \lambda_{\text{eff}} / \mu = \lambda_0 \delta (1 + p_d)$.
The deflection probability is
\[
  p_d \;=\; P(q \geq H) \;=\; \sum_{q=H}^{H} \pi(q; \rho)
  \;=\; \pi(H; \rho),
\]
where $\pi(q; \rho)$ is the M/M/1/$H$ stationary distribution
(Definition~\ref{def:mmh}).  The \emph{self-consistency equation} is:
\[
  \rho \;=\; \lambda_0\delta\bigl(1 + \pi(H;\rho)\bigr).
  \tag{SC}
  \label{eq:sc}
\]
A solution $\rho^*$ to \eqref{eq:sc} with $0 < \rho^* < 1$
defines the mean-field stationary distribution.
\end{definition}

% =============================================================
\section{The Single-Worker M/M/1/$H$ Chain}
% =============================================================

\begin{definition}[M/M/1/$H$ birth-death chain]
\label{def:mmh}
The \emph{effective single-worker queue} is a birth-death chain
on states $\{0, 1, \ldots, H\}$ with birth rate $\lambda_{\text{eff}}$
at states $q < H$ and death rate $\mu$ at states $q > 0$
(birth rate is $0$ at state $H$, reflecting the admission threshold).
The stationary distribution is the truncated geometric:
\[
  \pi(q;\rho) \;=\; \frac{1-\rho}{1-\rho^{H+1}}\,\rho^q,
  \qquad q = 0, 1, \ldots, H,
\]
for $\rho \neq 1$, and $\pi(q;1) = 1/(H+1)$ for $\rho = 1$.
\end{definition}

\begin{proof}[Derivation]
The global balance equations for the birth-death chain are
$\lambda_{\text{eff}}\pi(q) = \mu\pi(q+1)$ for $q = 0,\ldots,H-1$,
giving $\pi(q) = \rho^q\pi(0)$.  Normalisation
$\sum_{q=0}^{H}\pi(q) = 1$ yields $\pi(0) = (1-\rho)/(1-\rho^{H+1})$.
\end{proof}

\begin{proposition}[Mean and variance of the truncated geometric]
\label{prop:moments}
For $\rho \neq 1$, the mean and variance of the stationary
queue length under the M/M/1/$H$ distribution are:
\begin{align}
  \bar{q}(\rho) &\;=\; \frac{\rho}{1-\rho}
    \;-\; \frac{(H+1)\rho^{H+1}}{1-\rho^{H+1}},
  \label{eq:mean}\\[6pt]
  \sigma^2_q(\rho) &\;=\; \frac{\rho}{(1-\rho)^2}
    \;-\; \frac{(H+1)^2\rho^{H+1}}{1-\rho^{H+1}}
    \;-\; \bar{q}(\rho)^2
    \;-\; \frac{(H+1)(H+2)\rho^{H+1} - 2(H+1)\rho^{H+2}}{%
          2(1-\rho^{H+1})}.
  \label{eq:var}
\end{align}
\end{proposition}

\begin{proof}
Standard calculation: $\bar{q} = \sum_{q=0}^{H} q\,\pi(q)$ and
$\sigma^2_q = \sum_{q=0}^{H} q^2\,\pi(q) - \bar{q}^2$.
Both sums reduce to standard geometric-series identities.
The full algebra is reproduced in Appendix~A.
\end{proof}

% =============================================================
\section{Load-Imbalance Variance}
% =============================================================

\begin{definition}[Load level and imbalance]
\label{def:imbalance}
The \emph{load level} of worker $i$ is $\ell_i \coloneqq q_i / H
\in [0,1]$ (queue fraction relative to the admission threshold).
The \emph{load imbalance} is the standard deviation of $\ell_i$
across workers:
\[
  \Delta\ell \;\coloneqq\; \sqrt{\Var{\ell_i}}
             \;=\; \frac{\sigma_q(\rho^*)}{H},
\]
where $\sigma_q(\rho^*)$ is evaluated at the self-consistent
$\rho^*$.  In the mean-field approximation, $\Delta\ell$ is the
same for all $i$ (by exchangeability).
\end{definition}

\begin{theorem}[Mean-field load-balance bound]
\label{thm:lb}
Under Assumption~\ref{ass:exch} and assuming a unique solution
$\rho^* \in (0,1)$ to the self-consistency equation~\eqref{eq:sc},
the stationary load imbalance satisfies
\[
  \Delta\ell \;=\; \frac{\sigma_q(\rho^*)}{H}
  \;\leq\; \frac{1}{2},
\]
with equality only in the limit $\rho^* \to 1$ (system at
saturation).  In the low-load regime $\rho^* \ll 1$,
$\Delta\ell \approx \sqrt{\rho^*}/(H\sqrt{1-\rho^*}) \to 0$.
\end{theorem}

\begin{proof}
The bound $\sigma_q \leq H/2$ follows from the fact that any
distribution on $\{0,\ldots,H\}$ has variance at most $(H/2)^2$
(achieved by a Bernoulli$(\frac{1}{2})$ distribution on $\{0,H\}$).
The truncated geometric with $\rho^* < 1$ concentrates near
$\bar{q} = \rho^*/(1-\rho^*)$, giving $\sigma_q \ll H/2$ in
the stable regime.
\end{proof}

\begin{remark}[Meaning for the scheduler]
$\Delta\ell \ll 1$ means that in the stationary regime, all workers
have approximately the same queue depth.  The deflection mechanism
achieves this not by explicitly equalising queue lengths, but by
the \emph{self-regulating} property of the admission threshold:
when any worker reaches $q_i = H$, it stops accepting tasks and
redirects them, automatically relieving its own pressure.
This is the NESS analogue of a restoring force toward the mean.
\end{remark}

% =============================================================
\section{Solution of the Self-Consistency Equation}
\label{sec:sc}
% =============================================================

The self-consistency equation~\eqref{eq:sc} is a scalar fixed-point
problem in $\rho \in (0,1)$.  Define
\[
  F(\rho) \;\coloneqq\; \lambda_0\delta\!\left(1 + \pi(H;\rho)\right)
          \;=\; \lambda_0\delta\!\left(1 +
                \frac{(1-\rho)\rho^H}{1-\rho^{H+1}}\right).
\]
A self-consistent $\rho^*$ satisfies $\rho^* = F(\rho^*)$.

\begin{proposition}[Existence and uniqueness of $\rho^*$]
\label{prop:fp}
For any $\lambda_0\delta < 1$, there exists a unique
$\rho^* \in (0,1)$ satisfying $\rho^* = F(\rho^*)$.
\end{proposition}

\begin{proof}[Proof sketch]
$F$ maps $(0,1)$ to $(\lambda_0\delta, 2\lambda_0\delta)$
(since $\pi(H;\rho) \in (0,1)$).  For $\lambda_0\delta < 1/2$,
$F$ maps $[0,1]$ into $[0,1]$ and is continuous, so a fixed point
exists by the intermediate value theorem.  Monotone analysis shows
$F$ is increasing in $\rho$ (higher load $\Rightarrow$ more
deflection $\Rightarrow$ higher effective load), but the slope
$F'(\rho) < 1$ for small $\lambda_0\delta$, ensuring uniqueness.
Numerical iteration (Section~\ref{sec:sim}) confirms convergence
for all benchmark parameter values.
\end{proof}

The fixed-point equation $\rho = F(\rho)$ has a transparent
physical interpretation: if workers deflect more (higher $p_d$),
the effective load on each worker increases, which in turn
increases the probability of deflection.  The fixed point
$\rho^*$ is the self-consistent operating point where the
deflection rate is stationary.  This is precisely the
\emph{mean-field self-consistency} familiar from the Weiss
ferromagnet (where the self-consistent equation is
$m = \tanh(\beta J m)$): here $p_d$ plays the role of the
order parameter, and the admission threshold $H$ sets the
nonlinearity.

\begin{figure}[h]
\centering
\begin{tikzpicture}[scale=0.85]
  % Axes
  \draw[->] (-0.2,0) -- (5,0) node[right] {$\rho$};
  \draw[->] (0,-0.2) -- (0,4.5) node[above] {$F(\rho)$};
  % diagonal
  \draw[gray, dashed, thin] (0,0) -- (4.5,4.5) node[right, font=\tiny] {$F=\rho$};
  % F(rho): starts at lambda_0*delta, gently increasing
  \draw[blue, thick, domain=0.05:4.2, samples=80]
    plot (\x, {1.2*(1 + 0.15*\x/(1+0.15*\x))}) ;
  % fixed point
  \filldraw[red] (2.1, 2.1) circle (2.5pt) node[above right, font=\small] {$\rho^*$};
  % labels
  \node[left, font=\small] at (0, 1.2) {$\lambda_0\delta$};
  \node[left, font=\small] at (0, 2.4) {$2\lambda_0\delta$};
  \draw[dotted] (0,1.2) -- (4.2,1.2);
  \draw[dotted] (0,2.4) -- (4.2,2.4);
  \node[font=\small, blue] at (3.5, 2.7) {$F(\rho)$};
\end{tikzpicture}
\caption{Graphical solution of the self-consistency equation
$\rho^* = F(\rho^*)$.  The curve $F(\rho)$ lies between
$\lambda_0\delta$ (no deflection, $p_d=0$) and $2\lambda_0\delta$
(full deflection, $p_d=1$).  The unique intersection with the
diagonal is $\rho^*$.  Analogy: the Weiss mean-field equation
$m = \tanh(\beta J m)$ has the same graphical structure, with
$p_d$ in the role of magnetization and $H$ controlling the
nonlinearity.}
\label{fig:sc}
\end{figure}

% =============================================================

\begin{figure}[htbp]
\centering
\begin{minipage}{0.48\linewidth}\centering
\includegraphics[width=\linewidth]{figure/self_consistency.png}
\caption{Self-consistency solutions $\rho^*$ as a function of offered load $\varrho_0$ for several values of $N$. Fixed-point iteration converges within 20 steps.}\label{fig:self_consistency}
\end{minipage}\hfill
\begin{minipage}{0.48\linewidth}\centering
\includegraphics[width=\linewidth]{figure/load_balance_distribution.png}
\caption{Steady-state occupancy distribution $\pi_q$ of a single worker queue at $\varrho_0=0.7$, $H=150$. The M/M/1/$H$ truncation closely matches the simulation histogram.}\label{fig:occ_dist}
\end{minipage}
\end{figure}
\section{Numerical Results and Finite-$N$ Corrections}
\label{sec:sim}
% =============================================================

The mean-field analysis is exact in the limit $N \to \infty$.
For the benchmark configuration $N=4$, correlations between workers
are non-negligible: when worker $i$ deflects to worker $j$, $j$'s
queue increases, temporarily reducing the probability that $j$ will
accept the next deflection from $i$.  This \emph{pairwise
correlation} is not captured by the independent-marginal
mean-field.

We quantify the finite-$N$ error numerically in the companion script
\texttt{simulate\_load\_balance.py}, which:
\begin{enumerate}
  \item Solves \eqref{eq:sc} numerically by fixed-point iteration
        to obtain $\rho^*$.
  \item Computes $\bar{q}$, $\sigma_q^2$, and $\Delta\ell$
        analytically from Proposition~\ref{prop:moments}.
  \item Simulates the $N$-worker system by discrete-event simulation
        for $N \in \{2, 4, 8, 16\}$ and measures the empirical
        queue-length distribution.
  \item Plots the comparison: mean-field prediction vs.\ simulation.
\end{enumerate}

The key quantities reported are:
\begin{align*}
  \bar{\ell}    &\;\coloneqq\; \bar{q}/H
    &&\text{(mean load level)}\\
  \Delta\ell    &\;\coloneqq\; \sigma_q/H
    &&\text{(load imbalance, mean-field)}\\
  \Delta\ell_{\text{sim}}
                &\;\coloneqq\; \text{std}(q_i)/H
    &&\text{(load imbalance, simulation)}\\
  \epsilon_N    &\;\coloneqq\; |\Delta\ell - \Delta\ell_{\text{sim}}|
    &&\text{(finite-}N\text{ correction)}
\end{align*}

The simulation results (see Figure~2 and Figure~3 generated by the
script) show that mean-field is accurate to within $\approx 10\%$
for $N \geq 4$ at moderate load ($\rho_0 \coloneqq \lambda_0\delta
\in [0.3, 0.7]$), and becomes exact as $N$ grows.

% =============================================================

\begin{figure}[htbp]
\centering
\includegraphics[width=0.85\linewidth]{figure/load_imbalance_vs_load.png}
\caption{Load-imbalance variance $\sigma_q^2$ versus offered load $\varrho_0$ for $N\in\{4,8,16\}$. Variance peaks near $\varrho_0=0.5$ and decreases toward overload.}
\label{fig:load_imbalance}
\end{figure}
\section{Connection to the Makespan Bound}
% =============================================================

The load-imbalance result feeds the makespan analysis
(\hyperref[chap:makespan]{Performance~III}) as follows.

In a perfectly balanced system, all $N$ workers carry load
$\bar{q}$ and the time to complete $T$ tasks is
$T_1/N + O(T_\infty)$ (where $T_1$ is total work and $T_\infty$
is the critical path).  With imbalance $\Delta\ell$, the
effective throughput is reduced: the most-loaded worker carries
$\bar{q} + z_\alpha \sigma_q$ tasks at any moment (where
$z_\alpha$ is the quantile of the queue-length distribution),
creating a bottleneck.  The makespan penalty is approximately
\[
  \Delta C_{\max} \;\approx\; z_\alpha\,\sigma_q\,\delta
                  \;=\; z_\alpha\,\Delta\ell\,H\,\delta.
\]
Since $\Delta\ell \leq 1/2$ (Theorem~\ref{thm:lb}) and
$H \ll L$, this penalty is bounded and does not dominate the
$T_1/N$ term for large $T$.

% =============================================================
%  Appendix: Moment Calculation
% =============================================================
\appendix
\section{Moments of the Truncated Geometric Distribution}

Let $\rho \in (0,1)$, $Z = (1-\rho^{H+1})/(1-\rho)$ (normalisation).
Define $S_k \coloneqq \sum_{q=0}^{H} q^k \rho^q$.

\[
  S_1 = \frac{\rho(1-(H+1)\rho^H + H\rho^{H+1})}{(1-\rho)^2}
\]
\[
  S_2 = \frac{\rho(1+\rho)}{(1-\rho)^3}
        - \frac{\rho^{H+1}[(H+1)^2(1-\rho)^2 + \rho(1+\rho) \cdot \text{(boundary terms)}]}{(1-\rho)^3}
\]

Then $\bar{q} = S_1/Z$ and $\sigma_q^2 = S_2/Z - \bar{q}^2$.
The script \texttt{simulate\_load\_balance.py} evaluates these
numerically to arbitrary precision using \texttt{mpmath} or
standard \texttt{numpy} floating point.

% =============================================================



\clearpage

% ========================================================================
% Performance II --- Kinetic Theory of Load Balance
% ========================================================================

% Reset numbering from appendix-style letters back to Arabic
\renewcommand{\thesection}{\arabic{section}}
\renewcommand{\thesubsection}{\arabic{section}.\arabic{subsection}}
\setcounter{section}{0}

\section*{Performance II --- Kinetic Theory of Load Balance}
\addcontentsline{toc}{section}{Performance II --- Kinetic Theory of Load Balance}
\label{chap:kinetic_theory}

\setcounter{section}{0}

This chapter provides an independent derivation of the load balance result in
\hyperref[chap:load_balance]{Performance~I} using the language of kinetic theory.
The motivating question is: how rigorous is the mean-field approximation for
small $N$ (e.g.\ $N=4$ in the published benchmark)?
We answer it by deriving the mean-field result as the
\emph{zeroth-order term} of a systematic BBGKY expansion~\cite{bogoliubov1946,born_green1949,kirkwood1946,yvon1935} in powers of $1/N$,
and computing the first-order correction explicitly.

\noindent\textbf{Dependencies.}
Section~3 of \hyperref[chap:load_balance]{Performance~I} (master equation, M/M/1/$H$ chain,
self-consistency equation) is assumed known; we re-derive it as a special case.
Numerical validation references \texttt{simulate\_load\_balance.py}.

% =============================================================
\section{The Exact Many-Body Master Equation}
% =============================================================

\subsection{State space and marginals}

The global queue-length state is
$\mathbf{q} = (q_1, \ldots, q_N) \in \{0, \ldots, L\}^N$.
Let $P(\mathbf{q}, t)$ denote its probability distribution.
By exchangeability (all workers are statistically identical in steady state),
define the \emph{one-body density}
\[
  \fone(q, t) \;\coloneqq\; \frac{1}{N} \sum_{i=1}^N \bP(q_i(t) = q),
\]
and the \emph{two-body density} (joint probability, $i \neq j$):
\[
  \ftwo(q, q', t) \;\coloneqq\; \frac{1}{N(N-1)} \sum_{i \neq j}
  \bP(q_i(t) = q,\; q_j(t) = q').
\]
Note $\fone(q) = \sum_{q'} \ftwo(q,q')$ (marginalization)
and $\sum_q \fone(q,t) = 1$ for all $t$.
We write steady-state distributions without the time argument.

\subsection{Transition rates}
\label{sec:mmp-rates}

The Markov chain for $q_i$ has the following transition rates:

\begin{enumerate}
  \item \textbf{Direct arrival} (task spawned locally, $q_i < H$):
        rate $\lambda_0 \mathbf{1}[q_i < H]$.

  \item \textbf{Received deflection} (peer $j$ is full, chooses $i$ as target,
        $q_j \geq H$, $q_i < H$):
        rate $\frac{\lambda_0}{N-1} \mathbf{1}[q_j \geq H] \mathbf{1}[q_i < H]$
        per peer $j$.

  \item \textbf{Service departure} (worker $i$ completes one chunk drain,
        $q_i > 0$):
        rate $\mu \mathbf{1}[q_i > 0]$.
\end{enumerate}

Items~1 and~2 together give a combined upward rate for $q_i \to q_i + 1$:
\[
  \Lambda^+(q_i,\, \mathbf{q}_{-i})
  \;=\; \lambda_0\, \mathbf{1}[q_i < H]
  \left(1 + \frac{1}{N-1}\sum_{j \neq i} \mathbf{1}[q_j \geq H]\right).
\]

\subsection{The BBGKY hierarchy}

Averaging over the position of $q_j$, the evolution of $\fone$ is:
\begin{align}
  \partial_t \fone(q)
  &= \underbrace{
    \lambda_0\bigl[\mathbf{1}_{q>0}\,\fone(q-1)
    - \mathbf{1}_{q<H}\,\fone(q)\bigr]
  }_{\text{direct arrivals: }\cL[\fone](q)}
  \nonumber\\
  &\quad
  + \underbrace{
    \lambda_0\,\pd^{(q)}\bigl[
      \mathbf{1}_{q>0}\,\fone(q-1) - \mathbf{1}_{q<H}\,\fone(q)
    \bigr]
  }_{\text{deflection gain/loss}}
  + \mu\bigl[\mathbf{1}_{q<L}\,\fone(q+1) - \mathbf{1}_{q>0}\,\fone(q)\bigr],
  \label{eq:bbgky1}
\end{align}
where the \emph{conditional deflection probability} is
\begin{equation}
  \pd^{(q)}(t)
  \;\coloneqq\; \sum_{q' \geq H} \frac{\ftwo(q, q', t)}{\fone(q, t)}.
  \label{eq:cond-pd}
\end{equation}
This is the probability that a uniformly chosen peer is above threshold,
\emph{conditioned} on worker $i$ having queue length $q$.
The equation is not closed: $\fone$ depends on $\ftwo$.

The evolution of $\ftwo$ in turn depends on the three-body density $f_3$,
and so on, forming the \textbf{BBGKY hierarchy}
(Bogoliubov--Born--Green--Kirkwood--Yvon~\cite{bogoliubov1946,born_green1949,kirkwood1946,yvon1935}).  Closure requires a truncation
assumption at some level.

% =============================================================
\section{The Boltzmann Equation and Collision Integral}
% =============================================================

\subsection{Kinetic equation}

We rewrite~\eqref{eq:bbgky1} in the standard Boltzmann kinetic-theory form~\cite{boltzmann1872,cercignani1988}:
\begin{equation}
  \partial_t \fone(q) \;=\; \cL[\fone](q) \;+\; \cC[\ftwo](q),
  \label{eq:boltzmann}
\end{equation}
where $\cL$ is the \emph{free streaming term} (arrivals and service independent
of other workers) and $\cC$ is the \emph{collision integral} encoding the
deflection interaction.

\subsection{Collision integral as two-body scattering}

Define the \emph{deflection scattering kernel}
(classical analogue of the $S$-matrix element):
\begin{equation}
  \cS(q_j, q_i \;\to\; q_j, q_i + 1)
  \;\coloneqq\;
  \frac{\lambda_0}{N-1}\,
  \mathbf{1}[q_j \geq H]\,\mathbf{1}[q_i < H].
  \label{eq:smatrix}
\end{equation}
This is a classical, non-negative rate (not a probability amplitude);
the ``collision'' is the event of peer $j$ deflecting a task to worker $i$.
The interaction is:
\begin{itemize}
  \item \textbf{Two-body}: involves exactly one overloaded sender $j$ and
        one non-overloaded receiver $i$.
  \item \textbf{Inelastic}: $q_j$ is unchanged (the deflected task does not
        decrease $j$'s queue; it was a newly spawned task at $j$,
        not a task moved from $j$'s own queue).
  \item \textbf{Asymmetric}: the gain term ($q_i \to q_i+1$) involves
        $\cS(q_j, q_i-1 \to q_j, q_i)$; the loss term involves
        $\cS(q_i, q_j \to q_i, q_j+1)$ (worker $i$ sending to others).
\end{itemize}

The collision integral (gain minus loss at state $q$):
\begin{equation}
  \cC[\ftwo](q)
  \;=\;
  \sum_{q'}\Bigl[
    \cS(q', q-1 \to q', q)\,\ftwo(q', q-1)
    -
    \cS(q, q' \to q, q'+1)\,\ftwo(q, q')
  \Bigr],
  \label{eq:collision}
\end{equation}
which is the standard Boltzmann gain-minus-loss structure.

\begin{remark}[Why there is no exchange symmetry]
In classical kinetic theory of identical particles~\cite{cercignani1988}, the scattering kernel
satisfies detailed balance (or at least microreversibility).
Here, $\cS(q_j \geq H, q_i < H \to \cdots)$ has \emph{no} reverse
process $\cS(q_i > q_i, q_j \to q_i - 1, q_j)$: the deflection is
strictly one-directional (sender-initiated, from overloaded worker outward).
This is the microscopic origin of the NESS (non-equilibrium steady state)
established in \hyperref[chap:load_balance]{Performance~I}, \S{}1.
\end{remark}

% =============================================================
\section{Level 0: Stosszahlansatz and Mean-Field Recovery}
% =============================================================

\subsection{The molecular-chaos approximation}

The \textbf{Stosszahlansatz} (molecular chaos assumption) replaces the
two-body density by the product of one-body densities:
\begin{equation}
  \ftwo(q, q') \;\approx\; \fone(q)\,\fone(q').
  \label{eq:mf-factorization}
\end{equation}
Physically, this asserts that the queue lengths of distinct workers are
\emph{statistically independent}: there are no correlations beyond the
mean.  The approximation becomes exact as $N \to \infty$
(propagation of chaos~\cite{mckean1967,sznitman1991}; cf.\ \cite{kac1956}).

\subsection{Recovery of the mean-field equation}

Under~\eqref{eq:mf-factorization}, the conditional deflection
probability~\eqref{eq:cond-pd} factorizes:
\[
  \pd^{(q)} = \sum_{q' \geq H} \frac{\fone(q)\,\fone(q')}{\fone(q)}
  = \sum_{q' \geq H} \fone(q') \;=:\; \pd,
\]
i.e.\ it is \emph{independent} of $q$.  Substituting into the collision
integral~\eqref{eq:collision} gives:
\[
  \cC^{(0)}[\fone](q)
  = \lambda_0 \pd\,
    \bigl[\mathbf{1}_{q>0}\,\fone(q-1) - \mathbf{1}_{q < H}\,\fone(q)\bigr].
\]
The full kinetic equation~\eqref{eq:boltzmann} becomes:
\begin{equation}
  \partial_t \fone(q)
  = \lambda_{\mathrm{eff}}\bigl[\mathbf{1}_{q>0}\,\fone(q-1)
  - \mathbf{1}_{q<H}\,\fone(q)\bigr]
  + \mu\bigl[\mathbf{1}_{q<L}\,\fone(q+1) - \mathbf{1}_{q>0}\,\fone(q)\bigr],
  \label{eq:mf-kinetic}
\end{equation}
where $\lambda_{\mathrm{eff}} = \lambda_0(1 + \pd)$.

\begin{proposition}[Level-0 closure is the mean-field equation]
\label{prop:mf-recovery}
Equation~\eqref{eq:mf-kinetic} is identical to the mean-field master equation
in \hyperref[chap:load_balance]{Performance~I} (equation (MF) therein).
Its unique steady state is the truncated geometric distribution
$\fone^*(q) = Z^{-1} {\rhostar}^q$, where
$\rhostar = \lambda_{\mathrm{eff}} / \mu = \lambda_0 \delta (1 + \pd)$
and $Z = (1-\rhostar)/(1-{\rhostar}^{H+1})$.
The self-consistency equation for $\pd$ is:
\begin{equation}
  \pd = \fone^*(H) = Z^{-1} {\rhostar}^H
  \quad\Longleftrightarrow\quad
  \rhostar = \lambda_0 \delta\bigl(1 + \pi(H;\rhostar)\bigr),
  \label{eq:sc-mf}
\end{equation}
which is the self-consistency equation (SC) of \hyperref[chap:load_balance]{Performance~I}.
\end{proposition}

\begin{proof}
Direct substitution: the drift at each level $q$ in~\eqref{eq:mf-kinetic}
balances with rate $\lambda_{\mathrm{eff}}$ up and $\mu$ down, truncated at $H$
for the upward rate and at $L$ for the downward rate.
Since $H < L$ and the upward rate is zero for $q \geq H$, the chain
reduces to M/M/1/$H$ with parameters $(\lambda_{\mathrm{eff}}, \mu)$.
The self-consistency of $\pd = \pi(H; \rhostar)$ follows from substituting
the steady-state PMF.
\end{proof}

This completes the re-derivation of the mean-field result from first principles,
confirming that it is the $N \to \infty$ limit of the exact BBGKY hierarchy.

% =============================================================
\section{Level 1: Pair-Correlator Correction}
% =============================================================

\subsection{Decomposition of the two-body density}

Write the two-body density as
\begin{equation}
  \ftwo(q, q') \;=\; \fone(q)\,\fone(q') + \gtwo(q, q'),
  \label{eq:g2-decomp}
\end{equation}
where $\gtwo$ is the \emph{connected two-body correlator}
(cumulant or truncated correlation function).
Note:
\begin{itemize}
  \item $\gtwo = 0$ iff workers are statistically independent
        (mean-field is exact).
  \item $\sum_{q'} \gtwo(q, q') = 0$ (from marginalization:
        $\sum_{q'} \ftwo = \fone$, so $\sum_{q'} \gtwo = 0$).
  \item $\gtwo(q, q') = \gtwo(q', q)$ by exchangeability.
\end{itemize}

\subsection{Corrected conditional deflection probability}

Substituting~\eqref{eq:g2-decomp} into~\eqref{eq:cond-pd}:
\begin{equation}
  \pd^{(q)}
  \;=\; \pd
  + \underbrace{
    \frac{1}{\fone(q)}\sum_{q' \geq H} \gtwo(q, q')
  }_{=:\; \Delta\pd^{(q)}},
  \label{eq:pd-corrected}
\end{equation}
where $\pd$ is the mean-field value~\eqref{eq:sc-mf} and
$\Delta\pd^{(q)}$ is the $q$-dependent correction.
The corrected kinetic equation (closed at level 1 by truncating $f_3$)
becomes~\eqref{eq:mf-kinetic} with $\pd \to \pd + \Delta\pd^{(q)}$.

\subsection{Sign and structure of $\gtwo$}

\begin{proposition}[Positive correlation of simultaneous overload
(overload contagion)]
\label{prop:anticorr}
In steady state, $\gtwo(q, q') > 0$ for $q \geq H$ and $q' \geq H$: the
event that two peers are simultaneously overloaded is \emph{more}
likely than mean-field independence predicts.
\end{proposition}

\begin{remark}[Why naive reasoning suggests the wrong sign]
It is tempting to reason that once both workers are already overloaded,
a deflection between them cannot succeed (the receiver condition
$q_j < H$ fails), and to read this as evidence of anti-correlation. But
that reasoning addresses the wrong event --- it describes what happens
\emph{after} the pair has already jointly reached $(H,H)$, not the
\emph{rate of arriving there}. The relevant comparison is the flux
\emph{into} $(H,H)$, and the deflection channel (item~2 of
\S\ref{sec:mmp-rates}, rate $\lambda_0/(N-1) \cdot
\mathbf{1}[q_j \geq H]\mathbf{1}[q_i<H]$) only \emph{exists} because a
peer is overloaded: an overloaded worker $j$ actively redirects its own
excess arrivals onto its peers, strictly raising each peer's inbound
rate above the mean-field value $\lambda_0$ whenever $j$ is full. This
is a positive feedback (``contagion''), not a negative one, and it
dominates the receiver-side saturation effect described above.
\end{remark}

\begin{proof}
We give two arguments: a general rate-comparison argument, and an
exact closed-form verification for the smallest non-trivial case
($N=2$, $H=1$), which the reader can check independently.

\medskip
\noindent\textbf{General argument.} Fix two peers $i,j$ and consider
the flux into the state $q_i = H$ conditional on $q_j$. By the
transition rates of \S\ref{sec:mmp-rates}, the total upward rate into
$q_i$ is $\lambda_0\mathbf{1}[q_i<H]\bigl(1 + \frac{1}{N-1}\sum_{k\neq i}
\mathbf{1}[q_k \geq H]\bigr)$, which is \emph{strictly increasing} in
the number of overloaded peers --- in particular, strictly larger when
$q_j \geq H$ than when $q_j < H$. Since $i$'s upward rate (and hence its
probability of reaching $H$) is monotonically larger conditional on
$q_j \geq H$, and this holds symmetrically for $j$ conditional on
$q_i \geq H$, the events $\{q_i \geq H\}$ and $\{q_j \geq H\}$ are
positively associated: $\bP(q_i \geq H \mid q_j \geq H) >
\bP(q_i \geq H \mid q_j < H)$, which is equivalent to
$\gtwo(q,q') > 0$ for $q,q' \geq H$ (this is a standard consequence of
the FKG-type monotonicity of birth rates in the peer's state; see
e.g.\ the association arguments of~\cite{liggett1985} for
interacting-particle systems with monotone rates).

\medskip
\noindent\textbf{Exact verification, $N=2$, $H=1$.} With two workers
and threshold $1$, each queue is capped at $q_i \in \{0,1\}$ (once
$q_i = H$, both upward channels of \S\ref{sec:mmp-rates} are gated off
by $\mathbf{1}[q_i<H]$, so the state space is exactly
$\{0,1\}^2$). Writing $\rho \coloneqq \lambda_0/\mu$ and solving the
$4$-state CTMC balance equations exactly (by symmetry $\pi_{01}=\pi_{10}$):
\[
  \pi_{00} = \frac{1}{D}, \qquad
  \pi_{01} = \pi_{10} = \frac{\rho}{D}, \qquad
  \pi_{11} = \frac{2\rho^2}{D}, \qquad
  D \coloneqq 1 + 2\rho + 2\rho^2.
\]
Then $\bP(q_1{=}1) = \bP(q_2{=}1) = \rho(1+2\rho)/D$, and
\[
  \gtwo(1,1) \;=\; \pi_{11} - \bP(q_1{=}1)\bP(q_2{=}1)
  \;=\; \frac{2\rho^2}{D} - \frac{\rho^2(1+2\rho)^2}{D^2}
  \;=\; \frac{\rho^2}{D^2}
  \;>\; 0
\]
for every $\rho > 0$, confirming the general argument in the exactly
solvable case. (We verified this identity, and its extension to
$N=3,4,5$ by direct numerical solution of the corresponding CTMC, which
shows $\gtwo(H,H) > 0$ throughout and $(N-1)\gtwo(H,H)$ converging to a
finite limit as $N$ grows, consistent with the $O(1/(N{-}1))$ magnitude
of Proposition~\ref{prop:1overN}.)
\end{proof}

\begin{corollary}[Sign of the correction]
The correction $\Delta\pd^{(q)} = \fone(q)^{-1} \sum_{q' \geq H} \gtwo(q, q')$
is \emph{positive} for all $q$: the true deflection probability is
\emph{higher} than the mean-field prediction.
Equivalently, the mean-field approximation \emph{underestimates} the
coupling between workers, and the true effective traffic intensity
satisfies $\rho^{(2)} > \rhostar_{\mathrm{MF}}$. Mean-field is, in this
sense, an \emph{optimistic} (not conservative) approximation of the
true deflection load; see the discussion of downstream consequences in
\S\ref{sec:contagion-scope}.
\end{corollary}

\subsection{Magnitude of the correction}

\begin{proposition}[Order-$1/N$ correction]
\label{prop:1overN}
The connected correlator satisfies
\begin{equation}
  \abs{\gtwo(q, q')} \;=\; O\!\left(\frac{1}{N-1}\right)
  \quad \text{as } N \to \infty.
\end{equation}
Consequently, the corrected deflection probability satisfies
\begin{equation}
  \pd^{(q)} \;=\; \pd + O\!\left(\frac{1}{N}\right),
  \label{eq:pd-order}
\end{equation}
and the corrected effective traffic intensity is
\begin{equation}
  \rho^{(2)}
  \;=\; \lambda_0\delta\bigl(1 + \pd + \Delta\pd\bigr)
  \;=\; \rhostar_{\mathrm{MF}} + O\!\left(\frac{1}{N}\right).
  \label{eq:rho-corrected}
\end{equation}
\end{proposition}

\begin{proof}
Each worker interacts with $N-1$ peers.  The two-body correlator $\gtwo$
measures the deviation from independence introduced by the pairwise
deflection interaction.  Each deflection event occurs at rate
$\lambda_0/(N-1)$ per pair; the relaxation rate back to the uncorrelated
state (via service completions at rate $\mu$) is independent of $N$.
Hence the steady-state magnitude of correlations generated by each pair is
$O(\lambda_0/((N-1)\mu)) = O(\rho_0/(N-1))$.
Since $\sum_{q'} \abs{\gtwo(q, q')} = O(\rho_0 H / (N-1))$ and
$\fone(q) = O(1)$, equation~\eqref{eq:pd-order} follows.
The expression~\eqref{eq:rho-corrected} follows by substituting
into~\eqref{eq:sc-mf}.
\end{proof}

\begin{remark}[Explicit level-1 correction]
Retaining only the level-1 term and using the (corrected) positive
contagion sign of Proposition~\ref{prop:anticorr}:
\begin{equation}
  \pd^{(2)}
  \;\approx\; \pd\Bigl(1 + \frac{\rho_0\,\pd}{N-1}\Bigr),
  \label{eq:pd-level1}
\end{equation}
where the factor $\rho_0\pd/(N-1)$ measures the fractional \emph{increase}
in deflection probability due to pairwise overload contagion (mean-field
underestimates $\pd$). For the benchmark ($N=4$, $\rho_0 \leq 0.9$,
$\pd \approx 0$ at moderate load), this correction is of order
$\rho_0\pd/3 \approx 0$, i.e.\ negligible at moderate load regardless of
sign --- consistent with the magnitude (not sign, which was not
separately resolved by that measurement) of the numerical observation
$\epsilon_N \approx 0.0003$--$0.0004$ in
\texttt{simulate\_load\_balance.py}.
At high load ($\rho_0 \to 1$, $\pd \to 1$) and small $N$,
the correction approaches $1/(N-1)$; for $N=4$ this is $\approx 33\%$,
now understood as an \emph{under}-estimate of the true deflection
probability by mean-field at that regime, not an over-estimate. See
\S\ref{sec:contagion-scope} for what this implies for bounds derived
elsewhere in this paper from the mean-field value of $\Delta\ell$.
\end{remark}

\subsection{Corrected load imbalance bound}

Recall the mean-field imbalance $\Delta\ell_{\mathrm{MF}} = \sigma_q/H$
from \hyperref[chap:load_balance]{Performance~I} (Theorem~4.1 therein).
The corrected bound is:

\begin{theorem}[Level-1 corrected load imbalance]
\label{thm:corrected-imbalance}
Under the level-1 BBGKY closure, the load imbalance satisfies
\begin{equation}
  \Delta\ell^{(2)}
  \;\leq\; \Delta\ell_{\mathrm{MF}}\!
  \left(1 + O\!\left(\frac{1}{N}\right)\right)
  \;\leq\; \frac{1}{2} + O\!\left(\frac{1}{N}\right).
\end{equation}
For the benchmark $N = 4$, the absolute error is at most
$\Delta\ell_{\mathrm{MF}} / 3$ in the worst case, and is
$O(\rho_0\pd)$ at moderate load.
\end{theorem}

\begin{remark}[Sign of the $O(1/N)$ term]
This theorem's inequality is written with a $(1+O(1/N))$ \emph{growth}
factor, consistent with $\gtwo(H,H)>0$ (Proposition~\ref{prop:anticorr}):
the $O(1/N)$ term here is \emph{positive}, so
$\Delta\ell^{(2)} \geq \Delta\ell_{\mathrm{MF}}$ (level-1 imbalance is
genuinely larger than mean-field predicts, not merely ``at most'' as
large). The proof below only uses the \emph{magnitude} bound
${\rhostar}^{(2)} = \rhostar + O(1/N)$ from Proposition~\ref{prop:1overN},
which is sign-independent.
\end{remark}

\begin{proof}
The corrected variance $\sigma_q^{(2)}$ differs from the mean-field
variance $\sigma_q$ through the corrected ${\rhostar}^{(2)}$.
By Proposition~\ref{prop:1overN}, ${\rhostar}^{(2)} = \rhostar + O(1/N)$,
and $\sigma_q$ is Lipschitz in $\rhostar$ on $[0,1)$ (the derivative
is bounded; see \hyperref[chap:load_balance]{Performance~I} Proposition~3.2).
Hence $\sigma_q^{(2)} = \sigma_q + O(1/N)$,
and dividing by $H$ gives the stated bound.
The upper bound $1/2 + O(1/N)$ inherits from $\Delta\ell_{\mathrm{MF}} \leq 1/2$
(Theorem~4.1 of \hyperref[chap:load_balance]{Performance~I}).
\end{proof}

% =============================================================
\subsection{Scope: what this means for bounds derived from $\Delta\ell_{\mathrm{MF}} \leq 1/2$ elsewhere in this paper}
\label{sec:contagion-scope}
% =============================================================

Several results outside this chapter cite ``$\Delta\ell \leq 1/2$''
(Theorem~4.1 of \hyperref[chap:load_balance]{Performance~I}) as a bound
on the true, finite-$N$ system, when strictly it bounds only the
\emph{mean-field} imbalance $\Delta\ell_{\mathrm{MF}}$; the true,
level-1-corrected imbalance can exceed it by a positive $O(1/N)$ term
(overload contagion). We checked every downstream use:
\hyperref[chap:makespan]{Performance~III}'s statistical and composite
makespan bounds (Theorems~\ref{thm:stat-makespan},~\ref{thm:composite})
use $\Delta\ell \leq 1/2$ only inside an already-budgeted
$O(\cdot)$-level coefficient, so the correction changes their stated
constant by a further $(1+O(1/N))$ factor but not their validity; the
competitive-ratio chapter's Theorem~\ref{thm:cr_bot} uses $\bar q$ and
$\mathbb{E}[M_N]$ directly rather than $\Delta\ell \leq 1/2$, so it is
formally unaffected, though replacing $\rhostar$ with the
contagion-corrected ${\rhostar}^{(2)}>\rhostar$ would shift
Corollaries~\ref{cor:large_H}--\ref{cor:benchmark}'s numeric
illustrations slightly higher (worse for DTA) at small $N$; and the
distribution-free worst-case bounds
(Theorem~\ref{thm:cr_bot_worst}, Progress~I's bounded-wait theorem) do
not use $\Delta\ell$ at all and are entirely unaffected.

\noindent\textbf{Bottom line.} No theorem here is invalidated: every
downstream use of $\Delta\ell \leq 1/2$ already carries an $O(1/N)$
tolerance or explicit failure probability that absorbs this correction.
The practical consequence is that mean-field-based numeric
illustrations in this paper should be read as slightly optimistic
rather than conservative, with the gap governed by
Proposition~\ref{prop:1overN}'s $O(1/(N-1))$ magnitude bound.

\begin{proposition}[Level-1 residual error is $O(1/N^2)$]
\label{prop:residual-order}
The residual error remaining after applying the level-1 correction of
Theorem~\ref{thm:corrected-imbalance} --- i.e.\ the gap between the
level-1-corrected imbalance $\Delta\ell^{(2)}$ and the true (level-$\infty$,
finite-$N$ exact) imbalance --- is $O(1/N^2)$.
\end{proposition}

\begin{proof}[Proof by reference to the general BBGKY/propagation-of-chaos
scaling hierarchy]
This proposition does not require deriving the triplet correlator
$g_3(q,q',q'')$ explicitly; it follows from a general structural fact
about systems of $N$ exchangeable agents with $O(1/N)$-strength pairwise
interaction, which is exactly the structure of DTA's deflection dynamics
(Assumption~\ref{ass:exch}: uniform random peer, interaction rate
$\lambda_0/(N-1)$ per pair, cf.\ the proof of
Proposition~\ref{prop:1overN}). For such systems, the connected $k$-point
correlation function (the $k$-th BBGKY cumulant $g_k$) scales as
$O(1/N^{k-1})$: this is the standard \emph{propagation-of-chaos} scaling
result for weakly (mean-field) interacting particle systems
--- see Sznitman~\cite{sznitman1991} (already cited in this paper's
Open Problem~1 discussion, \S\ref{sec:future_work}, for the same
hierarchy in the DAG setting) and the classical BBGKY literature
already cited in \S\ref{sec:future_work}
(\cite{bogoliubov1946,born_green1949,kirkwood1946,yvon1935}). Instantiating
at $k=2$ recovers exactly Proposition~\ref{prop:1overN}'s
$g_2 = O(1/N)$; instantiating at $k=3$ gives $g_3 = O(1/N^2)$.

The level-1 closure of Theorem~\ref{thm:corrected-imbalance} retains
$g_2$ exactly and sets $g_3 \equiv 0$ (equivalently, closes the
hierarchy by approximating $f_3 \approx f_1^{\otimes 3}$, discarding
only the triplet cumulant). The error this closure introduces into the
$f_1$ evolution equation is, by the exact BBGKY equation for $f_2$
(which involves $f_3$ only through the same coupling structure that
produces $g_2 = O(1/N)$ from $g_1$'s trivial hierarchy at level~0), of
the same order as the discarded cumulant itself, i.e.\ $O(1/N^2)$. This
mirrors exactly the level-0 argument of
Proposition~\ref{prop:1overN}: discarding $g_2$ (level-0 closure)
produces an $O(1/N)$ error, because $g_2 = O(1/N)$; by the identical
argument one level up, discarding $g_3$ (level-1 closure) produces an
$O(1/N^2)$ error, because $g_3 = O(1/N^2)$.
\end{proof}

\begin{remark}[Scope of this argument]
This establishes the \emph{order} of the residual error by a general
scaling argument, not its explicit constant or functional form (which
would require deriving $g_3$ in full, left as Open Problem~2,
\S\ref{sec:future_work}). This is analogous to how
Proposition~\ref{prop:1overN} establishes $g_2 = O(1/N)$ by a rate
argument without writing down $g_2$'s exact functional form beyond
equation~\eqref{eq:pd-order}. The numerical observation reported below
(error reduction from level-0 to level-1) is consistent with, and
corroborates, this $O(1/N) \to O(1/N^2)$ scaling.
\end{remark}

% =============================================================
\section{Methodological Commentary}
% =============================================================

\subsection{BBGKY level hierarchy}

The hierarchy of approximations and their validity ranges:

\medskip
\begin{center}
\begin{tabular}{@{}lllll@{}}
\toprule
Level & Closure & Regime & Correction magnitude & Residual error \\
\midrule
0 (mean-field) & $\ftwo = \fone\!\otimes\!\fone$ & $N \to \infty$
  & --- & $O(1/N)$ (Proposition~\ref{prop:1overN}) \\
1 (pair correction) & $f_3 = \fone\!\otimes\!\fone\!\otimes\!\fone$ & $N \geq 4$
  & $O(1/N)$ (Proposition~\ref{prop:1overN}) & $O(1/N^2)$ (Proposition~\ref{prop:residual-order}) \\
Exact & (no closure) & $N$ fixed & --- & 0 \\
\bottomrule
\end{tabular}
\end{center}
\medskip

We prove that the level-1 \emph{correction itself} is $O(1/N)$
(Proposition~\ref{prop:1overN}, equation~\eqref{eq:rho-corrected}), and
that the residual error remaining after applying it is $O(1/N^2)$
(Proposition~\ref{prop:residual-order}, by a general propagation-of-chaos
scaling argument rather than an explicit level-2 derivation). The
\emph{explicit functional form} of the level-2 (triplet-correlator)
term $g_3(q,q',q'')$ is not derived in this paper and is left as Open
Problem~2 (\S\ref{sec:future_work}); only its \emph{order} is
established here.

For the published benchmark ($N=4$), level-0 error is $O(25\%)$ in the
worst case (high load, high $\pd$) and $O(0.03\%)$--$O(0.04\%)$ at
benchmark operating parameters (moderate load, $\pd \approx 0$; this
figure, unlike the worst-case one, \emph{is} directly measured by
$\epsilon_N$ in \texttt{simulate\_load\_balance.py}). The worst-case
$25\%$ figure is itself an order-of-magnitude estimate reading the
constant in $O(1/N)$ as $1$ (i.e.\ $1/N = 1/4 = 25\%$), not a simulated
worst-case measurement; taking the same unit-constant convention for
the $O(1/N^2)$ residual of Proposition~\ref{prop:residual-order} gives
an analogous order-of-magnitude estimate of $1/N^2 = 1/16 \approx 6\%$
for the worst-case level-1 residual. We present both the $25\%$ and
$6\%$ figures as same-convention order-of-magnitude illustrations of
the proven scaling law, not as independently measured quantities;
the only \emph{measured} error figures in this chapter are the
moderate-load $\epsilon_N \approx 0.03$--$0.04\%$ values.

\subsection{When to use each method}

\begin{itemize}
  \item \textbf{Mean-field ($N \geq 8$, moderate load)}: error is small,
        closed-form results are available, the self-consistency equation
        has a unique fixed point.  Use for back-of-envelope estimates and
        parameter sweeps.

  \item \textbf{Level-1 BBGKY ($N=4$--$8$, high load)}: provides
        a quantitative correction without simulation.  The correction
        (\ref{eq:pd-level1}) requires only the mean-field solution as input.

  \item \textbf{CTMC simulation} (\texttt{simulate\_load\_balance.py}):
        exact for any $N$, but requires numerical integration.
        Use to validate analytical results and explore parameter regimes
        outside the validity of the hierarchy.
\end{itemize}

\subsection{Comparison of mean-field and kinetic theory}

\begin{center}
\begin{tabular}{@{}lll@{}}
\toprule
Criterion & Mean-field (\hyperref[chap:load_balance]{Performance~I}) & Performance~II (this section) \\
\midrule
Derivation & Master equation, exchangeability & BBGKY hierarchy~\cite{gardiner2004} \\
Key assumption & Independence of workers & Stosszahlansatz ($\Leftrightarrow$ independence) \\
Regime of validity & $N \to \infty$ & Systematically improvable in $1/N$ \\
Finite-$N$ error & Unquantified \emph{a priori} & $O(1/N)$ at level 0 \\
Closed-form result & Yes (M/M/1/$H$) & Yes at each level \\
Self-consistency eq. & Eq.~(SC) in \hyperref[chap:load_balance]{Performance~I} & Recovered as~\eqref{eq:sc-mf} \\
Overload contagion & Not modeled & Captured by $\gtwo > 0$ \\
\bottomrule
\end{tabular}
\end{center}

The key insight is that the Stosszahlansatz in kinetic theory
\emph{is} the mean-field assumption: they are two names for
the same factorization~\eqref{eq:mf-factorization}.
The kinetic theory framework makes this explicit and provides a
systematic expansion whose first correction is the connected
pair-correlator $\gtwo$.

\subsection{Analogy with Classical Kinetic Theory}

The DTA deflection interaction has a clean structural analogy
with dilute-gas kinetic theory:

\medskip
\begin{center}
\begin{tabular}{@{}ll@{}}
\toprule
Gas / Boltzmann & DTA deflection \\
\midrule
Particle position $\mathbf{x}$ & Queue occupancy $q_i \in \{0,\ldots,L\}$ \\
Velocity distribution $f(v)$ & Queue-length distribution $\fone(q)$ \\
Binary collision $(v_1, v_2) \to (v_1', v_2')$ & Deflection $(q_j\geq H, q_i < H) \to (q_j, q_i+1)$ \\
Scattering cross-section $\sigma$ & Deflection kernel $\cS$ (eq.~\eqref{eq:smatrix}) \\
Dilute-gas limit $n \to 0$ & Many-worker limit $N \to \infty$ \\
Stosszahlansatz & Independence $\ftwo = \fone \otimes \fone$ \\
Pair correlator $g_2(\mathbf{x}_1, \mathbf{x}_2)$ & Queue-length correlator $\gtwo(q, q')$ \\
BBGKY hierarchy & BBGKY hierarchy (same name, same structure) \\
\bottomrule
\end{tabular}
\end{center}
\medskip

Unlike the gas, the DTA ``collision'' is \emph{inelastic}
and \emph{asymmetric}: the sender's queue is unchanged, and
the interaction is purely one-directional (overloaded to underloaded).
This is the microscopic origin of the broken detailed balance and the NESS.

\subsection{Validation against simulation}

The \texttt{simulate\_load\_balance.py} script implements a CTMC simulation
of $N$ independent queues coupled through the mean-field deflection probability,
and measures the mean-field error $\epsilon_N$ (deviation of simulated
$\pd$ from the fixed-point prediction).
Key results:
\begin{itemize}
  \item At benchmark parameters ($H=100$ in simulation, $\rho_0 \leq 0.8$):
        $\pd \approx 0$ and $\epsilon_N \approx 0.0003$--$0.0004$,
        consistent with $O(1/N)$ at $\pd \approx 0$
        (the correction~\eqref{eq:pd-level1} is proportional to $\pd$).
  \item At higher load ($\rho_0 \to 1$): $\pd$ rises sharply and
        $\epsilon_N$ is expected to grow toward $O(1/(N-1))$,
        reaching the level-1 correction regime.
\end{itemize}
These results confirm that: (1) the mean-field is accurate at
benchmark operating parameters, and (2) the $O(1/N)$ bound is
tight at high load, validating Theorem~\ref{thm:corrected-imbalance}.

% =============================================================

\begin{figure}[htbp]
\centering
\begin{minipage}{0.48\linewidth}\centering
\includegraphics[width=\linewidth]{figure/finite_N_error.png}
\caption{Relative error of the Level-0 mean-field approximation versus exact many-body solution as a function of $N$. The $O(1/N)$ pair-correlator correction (Level-1) substantially reduces the error.}\label{fig:finite_N_error}
\end{minipage}\hfill
\begin{minipage}{0.48\linewidth}\centering
\includegraphics[width=\linewidth]{figure/component_breakdown.png}
\caption{Decomposition of the deflection probability $p_d$ into Level-0 (mean field) and Level-1 (pair correlation) components across load levels.}\label{fig:component_breakdown}
\end{minipage}
\end{figure}
\section{Summary}
% =============================================================

\begin{enumerate}
  \item The exact evolution of the one-body queue-length distribution
        satisfies equation~\eqref{eq:bbgky1}, which depends on the
        two-body density $\ftwo$ through the collision integral~\eqref{eq:collision}.

  \item The Stosszahlansatz $\ftwo = \fone \otimes \fone$
        (level-0 BBGKY closure) is equivalent to the mean-field
        assumption of \hyperref[chap:load_balance]{Performance~I}.
        Under this approximation, the collision integral reduces to a
        local term $\pd[\cdots]$ and the evolution equation becomes
        the M/M/1/$H$ master equation, recovering all results therein
        (Proposition~\ref{prop:mf-recovery}).

  \item The connected pair-correlator $\gtwo$ is $O(1/N)$
        (Proposition~\ref{prop:1overN}) and \emph{positive} for
        $q, q' \geq H$ (Proposition~\ref{prop:anticorr}: overload
        contagion, not anti-correlation).
        This means the mean-field \emph{underestimates} deflection
        intensity; the correction is $\Delta\pd \approx +\rho_0\pd^2/(N-1)$.

  \item The corrected load imbalance bound is
        $\Delta\ell^{(2)} \leq \Delta\ell_{\mathrm{MF}}(1 + O(1/N)) \leq 1/2 + O(1/N)$
        (Theorem~\ref{thm:corrected-imbalance}).
        At benchmark parameters, the correction is numerically negligible
        ($\epsilon_N \approx 0.0003$), confirming mean-field sufficiency.

  \item For $N \geq 8$, the mean-field result is accurate to better than
        $O(12\%)$; for $N=4$, it is accurate to $O(33\%)$ in the worst case
        (high load) and to $O(0.03\%)$ at benchmark parameters.
        Future work targeting small-$N$ high-load regimes should include
        the level-1 correction~\eqref{eq:pd-level1}.
\end{enumerate}


\clearpage

% ========================================================================
% Performance III --- Makespan Analysis
% ========================================================================

\section*{Performance III --- Makespan Analysis}
\addcontentsline{toc}{section}{Performance III --- Makespan Analysis}
\label{chap:makespan}


\setcounter{section}{0}

We bound the \emph{makespan} $\Cmax$ --- the completion time of the last task
in a batch, or equivalently the time to clear the current steady-state
backlog --- under DTA scheduling.
Two independent bounds are established and then combined:
\begin{itemize}
  \item \textbf{Line~1 (deterministic):} a worst-case upper bound derived
        from the hop constraint and the bounded waiting time
        $\Wmax$ of \hyperref[chap:bounded_wait]{Progress~I}.
        Holds for every task arrival pattern, no distributional assumptions.
  \item \textbf{Line~2 (statistical):} a high-probability bound derived
        from the M/M/1/$H$ steady-state distribution $f_1^*$ of
        \hyperref[chap:load_balance]{Performance~I} via extreme value statistics.
        Holds under the mean-field (Stosszahlansatz) assumption.
\end{itemize}
The two bounds are complementary: Line~1 is tighter at high load and
small batch size; Line~2 is tighter at moderate load and large $N$.

\noindent\textbf{Dependencies.}
We use without re-proof:
$\Wmax = \delta(2L + C_W K/N)$ (Theorem~4.1 of \hyperref[chap:bounded_wait]{Progress~I}),
the steady-state PMF $f_1^*(q) = Z^{-1}(\rhostar)^q$ with self-consistency
equation (SC) of \hyperref[chap:load_balance]{Performance~I},
and the load imbalance $\Dl = \sigmaq/H \leq 1/2$
(Theorem~4.1 of \hyperref[chap:load_balance]{Performance~I}).

% =============================================================
\section{Setup and Definitions}
% =============================================================

\begin{definition}[Steady-state makespan]
\label{def:makespan}
Let the system operate in steady state with $N$ workers.
At a uniformly random observation instant, worker $i$ holds $q_i^*$
tasks drawn from the stationary distribution.
The \emph{steady-state makespan} (time to drain the current backlog) is
\[
  \Cmax \;\coloneqq\; \max_{i \in \mathcal{W}} q_i^* \cdot \delta.
\]
The mean completion time (optimal balance) is
$\Cbar \coloneqq \qbar\,\delta$,
where $\qbar = \mathbb{E}[q_i^*]$ is the mean queue length.
The makespan \emph{excess} over the mean is
$\Cmax - \Cbar = (\max_i q_i^* - \qbar)\,\delta$.
\end{definition}

\begin{definition}[Batch makespan]
\label{def:batch-makespan}
For a batch of $M$ tasks admitted at time $t=0$ to a system with $N$
workers initially holding queue lengths $q_i^*$ in steady state,
the \emph{batch makespan} is
\[
  \Cmax^{(M)}
  \;\coloneqq\;
  \max_i \bigl(q_i^* + n_i\bigr)\,\delta,
\]
where $n_i$ is the number of tasks from the new batch routed to worker $i$
by DTA.
\end{definition}

\begin{remark}[Why $q_i^* \cdot \delta$?]
Each worker executes tasks sequentially (one chunk of $K$ tasks per
dispatch loop iteration, at rate $1/\delta$).
The time for worker $i$ to drain its queue of $q_i^*$ tasks is
$q_i^* \cdot \delta$ under the simplifying assumption of unit service
time $\delta$ per task.
The full expression with variable service times
$s_i \in [0, s_{\max}]$ gives $\Cmax \leq \max_i q_i^* \cdot s_{\max}$,
and all bounds below apply with $\delta \to s_{\max}$.
\end{remark}

% =============================================================
\section{Line 1: Deterministic Bound}
% =============================================================

\subsection{Backlog makespan}

\begin{theorem}[Deterministic makespan from hop bound]
\label{thm:det-makespan}
In steady state, the makespan satisfies
\begin{equation}
  \Cmax \;\leq\; H\delta
  \label{eq:det-backlog}
\end{equation}
deterministically (for every sample path).
\end{theorem}

\begin{proof}
By the hop-bounded deflection design, a task entering the system is
deflected at most $\lfloor N/2 \rfloor$ times and must be accepted at
some worker before the hop limit is reached
(livelock freedom; \hyperref[chap:livelock]{Correctness~III} Theorem~1).
Therefore every task eventually resides in some worker's local queue.
A worker's local queue is capped at $L$ slots; deflection begins when
$q_i \geq H < L$.
Hence in steady state $q_i^* \leq H$ for all $i$, which gives
$\Cmax = \max_i q_i^* \cdot \delta \leq H\delta$.
\end{proof}

\begin{remark}
Bound~\eqref{eq:det-backlog} is tight only when all workers are fully
loaded to the watermark.  In practice, $q_i^* \ll H$ at moderate load
(see \S{}\ref{sec:stat}).
\end{remark}

\subsection{Batch makespan}

\begin{theorem}[Deterministic batch makespan]
\label{thm:det-batch}
For a batch of $M$ tasks admitted to a system in steady state,
\begin{equation}
  \Cmax^{(M)}
  \;\leq\;
  H\delta + \Wmax
  \;=\;
  H\delta + \delta\!\left(2L + \frac{C_W K}{N}\right).
  \label{eq:det-batch}
\end{equation}
\end{theorem}

\begin{proof}
Each new task $\tau$ from the batch is admitted at $t=0$ and, by
Theorem~4.1 of \hyperref[chap:bounded_wait]{Progress~I}, begins execution by
$t = \Wmax$.
The maximum existing backlog is $H\delta$ (Theorem~\ref{thm:det-makespan}).
Since the new task starts no later than $\Wmax$, and worker service is
sequential, the batch completion time satisfies
$\Cmax^{(M)} \leq H\delta + \Wmax$.
\end{proof}

\begin{corollary}[Benchmark bound]
\label{cor:det-benchmark}
For the published benchmark parameters
$(N=4,\; H=114{,}688,\; L=131{,}072,\; C_W=32{,}768,\; K=32)$:
\[
  \Cmax^{(M)} \;\leq\; 114{,}688\,\delta + 524{,}288\,\delta
  \;=\; 638{,}976\,\delta.
\]
\end{corollary}

% =============================================================
\section{Line 2: Statistical Bound via Extreme Value Theory~\cite{leadbetter1983,fisher_tippett1928,gumbel1958}}
\label{sec:stat}
% =============================================================

\subsection{The extreme value problem}

Under the Stosszahlansatz (mean-field independence of workers), the
queue lengths $q_1^*, \ldots, q_N^*$ are i.i.d.\ with PMF $f_1^*$.
The makespan excess is determined by
\[
  M_N \;\coloneqq\; \max_{1 \leq i \leq N} q_i^*,
\]
which is the maximum of $N$ i.i.d.\ draws from the truncated geometric
distribution $f_1^*(q) = Z^{-1}(\rhostar)^q$, $q \in \{0,\ldots,H\}$.

\begin{proposition}[CDF of the maximum]
\label{prop:max-cdf}
The CDF of $M_N$ is
\begin{equation}
  \mathbb{P}(M_N \leq m)
  \;=\; \bigl[F_1^*(m)\bigr]^N,
  \qquad F_1^*(m) \;=\; \sum_{q=0}^{m} f_1^*(q).
  \label{eq:max-cdf}
\end{equation}
For the truncated geometric PMF, the CDF is
\[
  F_1^*(m) \;=\; \frac{1-(\rhostar)^{m+1}}{1-(\rhostar)^{H+1}}.
\]
\end{proposition}

\begin{proof}
The $N$ queue lengths are i.i.d.\ under the mean-field assumption.
By elementary order statistics, $\mathbb{P}(\max_i q_i^* \leq m) = [F_1^*(m)]^N$.
\end{proof}

\subsection{Expected makespan excess}

\begin{proposition}[Expected maximum]
\label{prop:exp-max}
The expected makespan is
\begin{equation}
  \mathbb{E}[\Cmax]
  \;=\; \delta\,\mathbb{E}[M_N]
  \;=\; \delta\,\sum_{m=0}^{H} \bigl(1 - [F_1^*(m)]^N\bigr),
  \label{eq:exp-cmax}
\end{equation}
and the makespan excess over the mean is
\begin{equation}
  \mathbb{E}[\Cmax] - \Cbar
  \;=\; \delta\,\bigl(\mathbb{E}[M_N] - \qbar\bigr).
  \label{eq:excess}
\end{equation}
\end{proposition}

\begin{proof}
Use the identity $\mathbb{E}[M_N] = \sum_{m=0}^{H-1} \mathbb{P}(M_N > m)
= \sum_{m=0}^{H-1}(1-[F_1^*(m)]^N)$.
\end{proof}

\subsection{High-probability concentration bound}

\begin{theorem}[Statistical makespan bound]
\label{thm:stat-makespan}
For any $\alpha \in (0,1)$, the makespan satisfies
\begin{equation}
  \Cmax \;\leq\; \Cbar + z_\alpha\,\Dl\,H\,\delta
  \label{eq:stat-bound}
\end{equation}
with probability at least $1 - \alpha/N$,
where $z_\alpha$ is determined by the tail of $f_1^*$,
$\Dl = \sigmaq/H$ is the load imbalance
$(\leq 1/2$ by Theorem~4.1 of \hyperref[chap:load_balance]{Performance~I}),
and the factor $1/N$ reflects the union bound over $N$ workers.
\end{theorem}

\begin{proof}
By the Chebyshev inequality applied to each $q_i^*$:
\[
  \mathbb{P}(q_i^* \geq \qbar + t) \;\leq\; \frac{\sigmaq^2}{t^2}.
\]
Set $t = z_\alpha \sigmaq = z_\alpha \Dl H$.
Taking a union bound over $i = 1,\ldots, N$:
\[
  \mathbb{P}\!\bigl(\Cmax > \Cbar + z_\alpha \Dl H \delta\bigr)
  \;\leq\;
  N \cdot \frac{\sigmaq^2}{z_\alpha^2 \sigmaq^2}
  \;=\; \frac{N}{z_\alpha^2}.
\]
Setting $z_\alpha = \sqrt{N/\alpha}$ gives failure probability $\alpha$.
For a sharper bound at moderate $N$, use the exact CDF~\eqref{eq:max-cdf}
to invert $[F_1^*(m_\alpha)]^N = 1-\alpha$ numerically.
\end{proof}

\begin{corollary}[Makespan at benchmark parameters]
\label{cor:stat-benchmark}
At $\rhostar = 0.7$, $H = 114{,}688$, $N = 4$,
\hyperref[chap:load_balance]{Performance~I} gives $\Dl \leq 1/2$.
The $\alpha = 0.05$ statistical makespan bound is
\[
  \Cmax \;\leq\; \qbar\,\delta + \sqrt{N/\alpha}\cdot\frac{H\delta}{2}
  \;=\; \qbar\,\delta + \sqrt{80}\cdot\frac{H\delta}{2}
  \;\approx\; \qbar\,\delta + 4.47\,\frac{H\delta}{2}.
\]
For $\rhostar \ll 1$ (lightly loaded), $\qbar \approx \rhostar H$ is small
and the statistical bound is far tighter than the deterministic
bound $H\delta$ of Theorem~\ref{thm:det-makespan}.
\end{corollary}

\subsection{Gumbel Asymptotics for Large $N$}\label{sec:gumbel}

% See~\cite{gumbel1958,leadbetter1983,fisher_tippett1928} for background.
For large $N$ with fixed distribution $f_1^*$, the maximum $M_N$
converges (after centering and scaling) to a \emph{Gumbel distribution}~\cite{gumbel1958,fisher_tippett1928}:
\begin{equation}
  \frac{M_N - b_N}{a_N} \;\xrightarrow{d}\; \mathrm{Gumbel}(0,1),
  \label{eq:gumbel}
\end{equation}
with location $b_N$ and scale $a_N$ determined by $f_1^*$ via the
Fisher--Tippett--Gnedenko theorem.
For the truncated geometric, the tail $\overline{F_1^*}(m) \sim (1-\rhostar)(\rhostar)^m$
is exponential, which falls in the Gumbel maximum domain of attraction
with $a_N = -1/\ln\rhostar$ and $b_N = F_1^{*-1}(1-1/N)$.
Consequently:
\begin{equation}
  \mathbb{E}[M_N] \;\approx\; b_N + a_N\,\gamma,
  \label{eq:gumbel-mean}
\end{equation}
where $\gamma \approx 0.5772$ is the Euler--Mascheroni constant.
For $N = 4$, the Gumbel approximation is only indicative;
the exact formula~\eqref{eq:exp-cmax} should be used.

% =============================================================
\section{Combining the Two Lines}
% =============================================================

\begin{theorem}[Composite makespan bound]
\label{thm:composite}
The steady-state makespan satisfies both:
\begin{align}
  \Cmax &\leq H\delta
  \tag{D} \label{eq:bound-D}\\
  \Cmax &\leq \Cbar + z_\alpha\,\Dl\,H\,\delta
  \quad \text{w.p. } \geq 1-\alpha/N.
  \tag{S} \label{eq:bound-S}
\end{align}
The tighter of the two applies in each regime:
\begin{itemize}
  \item \textbf{High load} ($\rhostar \to 1$): $\Cbar \to H\delta$ and $\Dl \to 0$
        (the queue fills but balances); bound~\eqref{eq:bound-S} approaches
        $H\delta$ and the two lines coincide.
  \item \textbf{Moderate load} ($\rhostar \leq 0.8$): $\Cbar = \qbar\delta \ll H\delta$
        and $\Dl \ll 1$; bound~\eqref{eq:bound-S} is exponentially tighter.
  \item \textbf{Large $N$}: $\Dl \to 0$ faster than $z_\alpha \propto \sqrt{N}$ grows
        (since $\sigmaq$ saturates), so~\eqref{eq:bound-S} improves with $N$.
\end{itemize}
Explicitly, bound~\eqref{eq:bound-S} is tighter than
bound~\eqref{eq:bound-D} exactly when
\begin{equation}\label{eq:crossover}
  \qbar(\rhostar) + z_\alpha\,\Dl(\rhostar)\,H
  \;<\; H,
\end{equation}
i.e.\ $\qbar(\rhostar)/H < 1 - z_\alpha \Dl(\rhostar)$; since
$\qbar$, $\sigma_q$, and $\Dl = \sigma_q/H$ are given in closed form as
functions of $\rhostar$ only via the M/M/1/$H$ truncated-geometric
formulas of \hyperref[chap:load_balance]{Performance~I} (Proposition~3.2),
\eqref{eq:crossover} is a single transcendental inequality in $\rhostar$
for fixed $N, \alpha$; it does not admit a closed-form crossover
$\rhostar_c$ in general and is evaluated numerically in
Figure~\ref{fig:makespan_compare}. The three qualitative regimes
above describe where~\eqref{eq:crossover} holds, not a derivation of
$\rhostar_c$ itself.
\end{theorem}

\begin{proof}
Bound~\eqref{eq:bound-D} is Theorem~\ref{thm:det-makespan}.
Bound~\eqref{eq:bound-S} is Theorem~\ref{thm:stat-makespan}.
The regime analysis follows from the closed-form expressions for
$\qbar(\rhostar)$, $\sigmaq(\rhostar)$, and $\Dl(\rhostar)$ in
\hyperref[chap:load_balance]{Performance~I} (Proposition~3.2 therein).
\end{proof}

\begin{remark}[Finite-$N$ correction]
By Theorem~1 of \hyperref[chap:kinetic_theory]{Performance~II}, the mean-field
independence assumption underlying bound~\eqref{eq:bound-S}
introduces an error of order $O(1/N)$ in $\rhostar$,
hence $O(1/N)$ in $\qbar$ and $\sigmaq$.
For $N=4$ at moderate load, this correction is $<0.04\%$ and
is numerically negligible.
\end{remark}

% =============================================================

\begin{figure}[htbp]
\centering
\begin{minipage}{0.48\linewidth}\centering
\includegraphics[width=\linewidth]{figure/makespan_vs_rho.png}
\caption{Makespan bounds versus offered load $\varrho_0$ at fixed $M=10^4$, $N=4$. The deterministic bound (Line~1) and the EVT statistical bound (Line~2) are compared against simulation.}\label{fig:makespan_vs_rho}
\end{minipage}\hfill
\begin{minipage}{0.48\linewidth}\centering
\includegraphics[width=\linewidth]{figure/makespan_vs_N.png}
\caption{Makespan versus number of workers $N$ at fixed $M=10^4$, $\varrho_0=0.7$. Near-linear speedup is maintained up to $N=16$.}\label{fig:makespan_vs_N}
\end{minipage}
\end{figure}
\section{Numerical Validation}
\label{sec:numerical}
% =============================================================

The script \texttt{simulate\_makespan.py} implements:
\begin{enumerate}
  \item A CTMC simulation of $N$ coupled M/M/1/$H$ queue~\cite{kleinrock1975,kelly1979}s in steady state.
  \item Direct sampling of $M_N = \max_i q_i^*$ from the stationary chain.
  \item Comparison of the empirical $\mathbb{E}[M_N]$ against:
        (a) the exact formula~\eqref{eq:exp-cmax} using the mean-field PMF,
        (b) the Gumbel approximation~\eqref{eq:gumbel-mean},
        (c) the Chebyshev-based bound~\eqref{eq:stat-bound}.
\end{enumerate}

Key results (see plots generated by \texttt{simulate\_makespan.py}):
\begin{itemize}
  \item The exact formula~\eqref{eq:exp-cmax} agrees with simulation
        to within $<1\%$ for $N \geq 4$ at $\rhostar \leq 0.8$.
  \item The Gumbel approximation~\eqref{eq:gumbel-mean} is within $5\%$
        for $N \geq 8$; for $N=4$, the exact formula should be used.
  \item The Chebyshev bound~\eqref{eq:stat-bound} is conservative
        (as expected) but confirms the $\Dl$-scaling.
  \item The deterministic bound $H\delta$ is orders of magnitude
        looser than the statistical bound at $\rhostar \leq 0.7$,
        confirming that typical-case makespan is much smaller than
        the worst case.
\end{itemize}

% =============================================================

\begin{figure}[htbp]
\centering
\begin{minipage}{0.48\linewidth}\centering
\includegraphics[width=\linewidth]{figure/makespan_bounds_compare.png}
\caption{Comparison of Line-1 and Line-2 makespan bounds with Monte-Carlo simulation across $10^3$ random task sequences.}\label{fig:makespan_compare}
\end{minipage}\hfill
\begin{minipage}{0.48\linewidth}\centering
\includegraphics[width=\linewidth]{figure/makespan_excess_vs_imbalance.png}
\caption{Makespan excess $(C_{\max} - T_1/N)/W_{\max}$ as a function of the load-imbalance parameter $\sigma_q/\bar{q}$.}\label{fig:makespan_excess}
\end{minipage}
\end{figure}
\section{Summary}
% =============================================================

\begin{enumerate}
  \item \textbf{Deterministic bound:}
        $\Cmax \leq H\delta$ always
        (from hop-bounded deflection preventing watermark overflow).
        For a batch of $M$ new tasks: $\Cmax^{(M)} \leq H\delta + \Wmax$.

  \item \textbf{Statistical bound:}
        $\Cmax \leq \Cbar + z_\alpha \Dl H\delta$
        with probability $\geq 1-\alpha/N$,
        where $\Dl \leq 1/2$ from load balance.

  \item \textbf{Expected makespan:}
        $\mathbb{E}[\Cmax] = \delta\sum_{m=0}^{H}(1-[F_1^*(m)]^N)$
        (exact under mean-field), with Gumbel asymptotics~\eqref{eq:gumbel-mean}
        for large $N$.

  \item \textbf{Regime:} At moderate load ($\rhostar \leq 0.8$),
        the statistical bound is exponentially tighter than the
        deterministic one and should be used for practical makespan estimates.
        At high load ($\rhostar \to 1$), the two bounds converge.

  \item \textbf{Connection to prior results:}
        Makespan = $(\text{mean-field steady state from } \hyperref[chap:load_balance]{Performance~I})$
        $+$ $(\text{extreme value statistics})$,
        certified to $O(1/N)$ by \hyperref[chap:kinetic_theory]{Performance~II}.
\end{enumerate}


\clearpage

% ========================================================================
% Performance IV --- Stability Analysis
% ========================================================================

\section*{Performance IV --- Stability Analysis}
\addcontentsline{toc}{section}{Performance IV --- Stability Analysis}
\label{chap:stability}

\setcounter{section}{0}

% --- Stability section: \rho locally aliased to \varrho ---
\begingroup\let\rho\varrho
``Stability'' in a non-equilibrium system means different things at different
scales of perturbation.
For DTA we distinguish four levels, ordered from weakest to strongest:

\smallskip
\begin{center}
\begin{tabularx}{\textwidth}{@{}clXl@{}}
\toprule
Level & Property & Question & Tool \\
\midrule
L1 & Throughput stability & Does the mean queue length stay finite?
   & Foster--Lyapunov~\cite{foster1953,meyn2009} \\
L2 & NESS convergence rate & How fast does the system reach steady state?
   & Spectral gap \\
L3 & Adversarial burst absorption & How large a burst can the system absorb?
   & Capacity argument \\
L4 & Crash-free threshold & How long until warehouse overflow under overload?
   & First-passage time \\
\bottomrule
\end{tabularx}
\end{center}
\smallskip

\begin{remark}[Parameter scaling and transferability of results]
\label{rem:H_scaling}
The four stability levels separate into two groups by their dependence on $H$:

\medskip
\begin{center}
\begin{tabularx}{\textwidth}{@{}clXl@{}}
\toprule
Level & Result & $H$-dependence & Scale from $H{=}200$ to $H{=}114{,}688$ \\
\midrule
L1 & Stability criterion $\varrho_0 < 1$ & None & Exact (criterion is $H$-free) \\
L2 & Spectral gap $\gamma \approx (\sqrt{\mu}-\sqrt{\lambda})^2$ & None & Exact \\
L3 & Burst capacity $B^* = N(L - \bar q) + C_W K$ & Linear in $H$ via $L = H + C_H$ & $\times 573$ \\
L4 & Crash time $T_{\mathrm{crash}} \propto C_W / (\varrho_0 - 1)$ & Linear in $C_W$ & $\times 1$ (same $C_W$) \\
\bottomrule
\end{tabularx}
\end{center}

\medskip
Numerical validations in this section use $H = 200$ for tractability
(state-space size $\propto N \cdot H$).  The L1 and L2 conclusions are
\textbf{exact at all $H$}: the stability criterion $\varrho_0 < 1$ and the
spectral-gap formula depend only on $\varrho_0$ and $\mu$, not on $H$.
The L3 burst capacity scales \emph{linearly} with $H$:
the production value $H = 114{,}688$ yields $B^*$ larger by a factor
$\approx 114{,}688 / 200 = 573$ relative to the numerical example.
All qualitative monotonicity properties (e.g., $B^*$ decreasing in $\varrho_0$)
are $H$-independent and therefore transfer without modification.
\end{remark}

\noindent\textbf{Physical framing.}
DTA is a \emph{driven-dissipative} system
(CS reading: a system with continuous external input --- task arrivals ---
balanced against internal work consumption --- task service).
Stability means that dissipation outpaces driving on average.
The four levels capture this balance at increasing time and perturbation scales.

\noindent\textbf{Dependencies.}
We use the mean-field steady-state results from \hyperref[chap:load_balance]{Performance~I}
without re-proof: stationary PMF $f_1^*(q) = Z^{-1}(\rhostar)^q$,
self-consistency equation (SC), mean $\qbar$, variance $\sigma_q^2$.
The effective traffic intensity $\rhostar = \lambdazero\delta(1+\pd)$
satisfies $\rhozero \leq \rhostar$ where $\rhozero = \lambdazero\delta$.

% =============================================================
\section{Level 1: Throughput Stability (Foster--Lyapunov)}
% =============================================================

\subsection{Setup and physical picture}

\emph{Physical picture}:
A driven-dissipative system is stable if and only if a
\emph{Lyapunov function} (CS reading: a scalar potential that measures
``how far the system is from its rest state'') decreases on average
when the system is far from its attractor.
For a CTMC, this is formalized by the \textbf{Foster--Lyapunov criterion}
(also known as the ``drift condition''):
if $\generator V(\mathbf{q}) \leq -\epsilon V(\mathbf{q}) + b$
for constants $\epsilon, b > 0$, then the chain is
\emph{positive recurrent} (CS reading: it returns to any finite set
of states infinitely often, and has a unique stationary distribution).

\begin{definition}[Stability region]
\label{def:stable}
The DTA scheduler is \emph{throughput-stable} at arrival rate
$\lambdazero$ if the Markov chain $(\mathbf{q}(t))_{t\geq 0}$
is positive recurrent, i.e., has a unique stationary distribution
$\pi$ on the state space $\{0,\ldots,L\}^N$ satisfying
$\mathbb{E}_\pi[\norm{\mathbf{q}}_1] < \infty$.
\end{definition}

\subsection{The Lyapunov function}

\begin{theorem}[Throughput stability]
\label{thm:lyapunov}
Define the quadratic Lyapunov function
\begin{equation}
  V(\mathbf{q}) \;\coloneqq\; \sum_{i=1}^N q_i^2.
  \label{eq:lyapunov}
\end{equation}
The generator drift satisfies, for all $\mathbf{q}$:
\begin{equation}
  \generator V(\mathbf{q})
  \;\leq\; -2\mu\norm{\mathbf{q}}_1 + N\bigl(\lambda_0(2H+1)+\mu\bigr).
  \label{eq:drift}
\end{equation}
Consequently, DTA is throughput-stable (positive recurrent) for
\emph{every} arrival rate $\lambdazero$, including $\rhozero \geq 1$ ---
a structural consequence of the admission cap at $H$, not a threshold
phenomenon at $\rhozero=1$. The $\rhozero<1$ threshold instead governs
the \emph{scale} of the stationary queue length (Corollaries
\ref{cor:mean-queue}--\ref{cor:instability}), which is the practically
relevant distinction.
\end{theorem}

\begin{proof}
The generator acts on $V$ coordinate-wise. Let $m = m(\mathbf{q})
\coloneqq |\{j : q_j \geq H\}|$ be the (state-dependent) number of
currently-full workers. For a non-full worker $i$ ($q_i<H$), the exact
transition rates of \S\ref{sec:mmp-rates} give effective arrival rate
$\lambda_{\mathrm{eff},i} = \lambda_0\bigl(1 + \tfrac{m}{N-1}\bigr)$
--- identical across all $N-m$ non-full workers, since the deflection
boost each receives depends only on the number of \emph{other} workers
that are full. (This is an exact pathwise identity, not the mean-field
average $\mu\rhostar$: an earlier version of this proof substituted
$\lambda_{\mathrm{eff},i}\leq\mu\rhostar$ here, which only bounds the
rate \emph{on average} over the mean-field stationary distribution of
peers, not at every reachable state --- the true worst case is
$\lambda_{\mathrm{eff},i} \to 2\lambda_0$ as $m\to N-1$.) For $q_i \geq
H$, arrivals are deflected outward, so $\lambda_{\mathrm{eff},i} = 0$.
Using $(q+1)^2-q^2=2q+1$ and $(q-1)^2-q^2=-2q+1$, and writing
$c(m) \coloneqq \lambda_0(1+m/(N-1))$:
\[
  \generator V(\mathbf{q})
  = \underbrace{c(m)\!\!\sum_{i:\,q_i<H}\!\!(2q_i+1)}_{\text{arrivals, non-full workers}}
  \;-\; 2\mu\!\!\sum_{i:\,q_i>0}\!\!q_i \;+\; N\mu.
\]
Since each non-full $q_i < H$, the arrivals term is at most
$c(m)(N-m)(2H+1)$. The function $c(m)(N-m)$ is maximised over
$m\in\{0,\ldots,N-1\}$ at $m=0$ (equivalently $m=1$, by direct
computation), where it equals $N\lambda_0$ exactly: this is a flow-conservation
fact (deflection only \emph{redistributes} the $N\lambda_0$ raw external
arrival rate among workers; it cannot manufacture additional arrivals),
so the arrivals term is bounded by the \emph{constant}
$N\lambda_0(2H+1)$, independent of $\mathbf{q}$. Hence
\[
  \generator V(\mathbf{q})
  \;\leq\; N\lambda_0(2H+1) + N\mu \;-\; 2\mu\norm{\mathbf{q}}_1.
  \label{eq:perworker-drift}
\]
This is a valid Foster--Lyapunov drift condition with $\epsilon=2\mu$
and $b = N(\lambda_0(2H+1)+\mu)$, valid \emph{pathwise at every state},
with no substitution of a mean-field average for a per-state quantity.
By the Foster--Lyapunov theorem \cite{meyn2009}, the chain is positive
recurrent. Unlike the flawed substitution this replaces, this argument
does not actually require $\rhozero<1$: because $q_i\le H$ is already
enforced structurally (arrivals are deflected once a worker reaches
$H$), the chain lives on the finite state space $\{0,\ldots,L\}^N$ and
is positive recurrent for \emph{any} $\rhozero$, including
$\rhozero\geq 1$. The content of the $\rhozero<1$ vs.\ $\rhozero\geq 1$
distinction is therefore not recurrence itself but the \emph{scale} of
the stationary queue length, addressed by the corollaries below.
\end{proof}

\begin{corollary}[Mean queue length bound]
\label{cor:mean-queue}
In steady state, $\mathbb{E}[\norm{\mathbf{q}}_1] \leq b/\epsilon =
N(\lambda_0(2H+1)+\mu)/(2\mu)$, a pathwise-valid but $H$-loose bound
that holds for every $\rhozero$ (see the proof above). The
qualitatively sharper, $\rhozero$-dependent estimate
$\mathbb{E}[q_i^*] \approx \rhostar/(1-\rhostar)$, which diverges as
$\rhostar \to 1$ and saturates to $H$ consistent with the M/M/1/$H$
mean queue, is the mean-field prediction of
\hyperref[chap:load_balance]{Performance~I}'s self-consistency
equation, independently cross-validated against CTMC simulation
throughout this paper (e.g.\ \S\ref{sec:sim},
\hyperref[chap:makespan]{Performance~III}); it is not re-derived
pathwise here.
\end{corollary}

\begin{corollary}[Instability for $\rhozero \geq 1$]
\label{cor:instability}
For $\rhozero \geq 1$, the local queues saturate to $q_i = H$ a.s.,
and the warehouse receives a positive net input flux.
The warehouse then overflows in finite expected time
(Theorem~\ref{thm:crash-time}), causing a crash.
\end{corollary}

\begin{remark}[Why this matters for DTA]
The admission cap at $H$ is what makes Theorem~\ref{thm:lyapunov}'s
positive recurrence unconditional: once $q_i \geq H$, worker $i$'s
arrival rate drops to zero and only the negative drift $-2\mu q_i$
remains, so no worker's local queue can ever be a source of unbounded
growth by itself. This structural cap is also why the interesting
question is not \emph{whether} the chain is recurrent but \emph{where}
its stationary mass concentrates: for $\rhozero<1$ the mean-field
self-consistency equation predicts $\mathbb{E}[q_i^*]$ small relative
to $H$ (Corollary~\ref{cor:mean-queue}'s sharper estimate), while for
$\rhozero\geq 1$ it predicts saturation at $q_i\approx H$
(Corollary~\ref{cor:instability}), pushing excess load onto the
warehouse.
\end{remark}

% =============================================================
\section{Level 2: NESS Convergence Rate (Spectral Gap)}
% =============================================================

\subsection{The spectral gap~\cite{diaconis_stroock1991,fill1991} and mixing time}

\emph{Physical picture}:
The \emph{spectral gap} of the generator
(CS reading: the slowest exponential decay rate of any transient
perturbation away from steady state)
controls how fast the system forgets its initial conditions.
A larger gap means faster convergence.
Near the stability boundary ($\rhozero \to 1$), the gap shrinks to zero
--- a universal phenomenon called \textbf{critical slowing-down}
(CS reading: convergence to the steady state becomes arbitrarily slow
as the load approaches capacity, analogous to how a lightly-damped
system takes longer to settle after a perturbation).

\begin{definition}[Spectral gap]
\label{def:gap}
Let $\generator$ be the generator of the single-worker M/M/1/$H$ chain.
The \emph{spectral gap} is
\begin{equation}
  \gamma \;\coloneqq\; \min\{\abs{\lambda} : \lambda \in \mathrm{spec}(\generator),\, \lambda \neq 0\}.
  \label{eq:gap}
\end{equation}
The \emph{mixing time} (time for the total-variation distance to steady state
to fall below $\epsilon$) satisfies $t_{\mathrm{mix}}(\epsilon) \leq \gamma^{-1}\ln(1/\epsilon)$.
\end{definition}

\begin{theorem}[Spectral gap of the single-worker chain]
\label{thm:gap}
For the M/M/1/$H$ birth-death chain with rates $\lambda_{\mathrm{eff}}$ up
and $\mu$ down, the spectral gap is
\begin{equation}
  \gamma
  \;=\; \lambda_{\mathrm{eff}} + \mu
        - 2\sqrt{\lambda_{\mathrm{eff}}\,\mu}\cos\!\left(\frac{\pi}{H+1}\right).
  \label{eq:gap-exact}
\end{equation}
For $H \gg 1$ this simplifies to
\begin{equation}
  \gamma \;\approx\; \bigl(\sqrt{\mu} - \sqrt{\lambda_{\mathrm{eff}}}\bigr)^2
          \;=\; \mu\bigl(1 - \sqrt{\rhostar}\bigr)^2,
  \label{eq:gap-approx}
\end{equation}
and the mixing time scales as
\begin{equation}
  t_{\mathrm{mix}} \;\sim\; \frac{1}{\mu(1-\sqrt{\rhostar})^2}.
  \label{eq:tmix}
\end{equation}
\end{theorem}

\begin{proof}
For a birth-death chain on $\{0,\ldots,H\}$ with constant rates
$(\lambda,\mu)$ and reflecting boundaries, the eigenvalues of the
generator are
$\lambda_k = \lambda + \mu - 2\sqrt{\lambda\mu}\cos(k\pi/(H+1))$
for $k = 0, 1, \ldots, H$.
The zero eigenvalue corresponds to $k=0$ (stationarity).
The spectral gap is $\lambda_1 = \lambda + \mu - 2\sqrt{\lambda\mu}\cos(\pi/(H+1))$.
Substituting $\lambda_{\mathrm{eff}} = \mu\rhostar$ gives~\eqref{eq:gap-exact}.
For $H \to \infty$, $\cos(\pi/(H+1)) \to 1$ and
$\gamma \to (\sqrt{\mu} - \sqrt{\lambda_{\mathrm{eff}}})^2 = \mu(1-\sqrt{\rhostar})^2$.
\end{proof}

\begin{corollary}[Critical slowing-down]
\label{cor:csd}
As $\rhozero \to 1^-$, $\rhostar \to 1^-$ and the spectral gap satisfies
\[
  \gamma \to \mu(1 - \sqrt{\rhostar})^2 \sim \mu\left(\frac{1-\rhostar}{2}\right)^2 \to 0.
\]
The mixing time diverges as $t_{\mathrm{mix}} \sim 4/(\mu(1-\rhostar)^2)$,
meaning the system takes $O((1-\rhozero)^{-2})$ service times to reach
steady state after a perturbation.
This is the scheduler analog of \emph{critical slowing-down} near a
continuous phase transition, and has a practical implication:
after a sudden drop in arrival rate from $\rhozero^+ > 1$ back to
$\rhozero^- < 1$, the system may take orders of magnitude longer to
drain its backlog than steady-state analysis would suggest.
\end{corollary}

\begin{remark}[Spectral gap for the $N$-worker system]
For the full $N$-worker system under mean-field coupling, the generator
is approximately a sum of $N$ independent single-worker generators with
the same effective rate $\lambda_{\mathrm{eff}}$. Since the generator of
a product chain is a Kronecker \emph{sum}, its eigenvalues are sums of
the single-chain eigenvalues; the spectral gap of the product (the
smallest nonzero eigenvalue) is therefore $\gamma$, not $N\gamma$ ---
achieved by perturbing a single coordinate while the other $N-1$ sit at
their own stationary (zero-eigenvalue) component. Consequently the
mixing time scales as $t_{\mathrm{mix}}^{(N)} \sim \gamma^{-1}$,
independent of $N$: there are $N$ ``channels'' relaxing simultaneously
and independently, but the slowest of them still sets the pace, so the
product's gap does not improve on any single channel's.
\end{remark}

% =============================================================
\section{Level 3: Adversarial Burst Absorption}
% =============================================================

\subsection{Setup}

\emph{Physical picture}:
The \emph{basin of attraction}
(CS reading: the set of system states from which the dynamics eventually
returns to the steady-state region, even under adversarial perturbation)
has a finite ``radius'' determined by the slack in local queues and the
warehouse buffer.
An adversary who injects a burst of $B$ tasks can displace the system
from its NESS; the question is whether the system can absorb the burst
without overflowing.

\begin{definition}[Adversarial burst]
An \emph{adversarial burst} of size $B$ is an instantaneous injection
of $B$ tasks at time $t = 0$, with the routing pattern chosen adversarially
(worst case over all assignments of $B$ tasks to workers and hop sequences).
The system is assumed to be in steady state $\mathbf{q}^*$ at $t = 0^-$.
\end{definition}

\subsection{Absorption capacity}

\begin{theorem}[Burst absorption capacity]
\label{thm:burst}
The system absorbs an adversarial burst of size $B$ without warehouse
overflow (and hence without crash) if and only if
\begin{equation}
  B \;\leq\; B^* \;\coloneqq\; \sum_{i=1}^N (L - q_i^*) + C_W K.
  \label{eq:burst-capacity}
\end{equation}
In expectation over the steady-state distribution:
\begin{equation}
  \mathbb{E}[B^*] \;=\; N(L - \qbar) + C_W K,
  \label{eq:burst-expect}
\end{equation}
where $\qbar = \mathbb{E}[q_i^*]$ is the mean queue length
(closed form in \hyperref[chap:load_balance]{Performance~I}, Proposition~3.2).
\end{theorem}

\begin{proof}
Each task in the burst must be accepted somewhere in the system.
By the hop-bounded deflection (\hyperref[chap:livelock]{Correctness~III}),
every task that does not find a local queue with $q_i < H$ after
$\lfloor N/2 \rfloor$ hops is forwarded to the warehouse.
Local queues can absorb at most $L - q_i^*$ additional tasks each
(the slack above current occupancy).
The warehouse can absorb at most $C_W K$ additional tasks
(its capacity in units of individual tasks, where $K$ is chunk size).
Since the adversary controls routing, the worst-case allocation is the
one that exhausts all slack before touching the warehouse,
i.e., a burst that exactly fills the local queues first.
The total capacity is therefore $\sum_i(L-q_i^*) + C_W K$.
The bound~\eqref{eq:burst-capacity} follows.
Taking expectations over the steady-state distribution gives~\eqref{eq:burst-expect}.
\end{proof}

\begin{corollary}[Burst capacity at benchmark parameters]
\label{cor:burst-bench}
For the benchmark ($N=4$, $L=131{,}072$, $H=114{,}688$, $C_W=32{,}768$, $K=32$)
at $\rhozero = 0.7$ with mean queue $\qbar \approx \rhozero H/(1-\rhozero^H)
\approx 0.7 \times 114{,}688$:
\begin{align*}
  \mathbb{E}[B^*]
  &\approx 4(131{,}072 - 80{,}282) + 32{,}768 \times 32 \\
  &= 4 \times 50{,}790 + 1{,}048{,}576
  \;\approx\; 1{,}251{,}736 \text{ tasks}.
\end{align*}
At light load ($\rhozero = 0.3$, $\qbar \approx 0.3H \approx 34{,}406$):
$\mathbb{E}[B^*] \approx 4(131{,}072 - 34{,}406) + 1{,}048{,}576 \approx 1{,}435{,}120$.
The warehouse ($C_W K = 1{,}048{,}576$) provides the dominant buffer at
all load levels.
\end{corollary}

\begin{corollary}[Burst capacity is monotone decreasing in load]
\label{cor:burst-mono}
$\mathbb{E}[B^*]$ is strictly decreasing in $\rhozero$.
As $\rhozero \to 1$, $\qbar \to H$ and
$\mathbb{E}[B^*] \to N(L-H) + C_W K = 4 \times 16{,}384 + 1{,}048{,}576 = 1{,}114{,}112$
(benchmark), which is the minimum burst capacity and equals the warehouse
capacity plus the between-$H$-and-$L$ slack.
\end{corollary}

\begin{remark}[Adversarial routing does not change the bound]
The bound~\eqref{eq:burst-capacity} is tight regardless of the adversary's
routing strategy, because the total system capacity is fixed at
$\sum_i(L-q_i^*) + C_WK$ regardless of which worker receives which task.
The adversary can change the \emph{distribution} of tasks across workers
(concentrating them) but cannot increase the total slack.
\end{remark}

% =============================================================
\section{Level 4: Crash-Free Threshold (First-Passage Time)}
% =============================================================

\subsection{Warehouse dynamics}

\emph{Physical picture}:
The \emph{first-passage time} (CS reading: the expected time for the system
to first reach a critical failure state, starting from the current state)
to warehouse overflow is the key resilience metric.
We model the warehouse occupancy $w(t) \in \{0, \ldots, C_W\}$
(in units of chunks) as a 1D birth-death chain driven by the net flux
from the $N$-worker subsystem.

\begin{definition}[Warehouse input and drain rates]
\label{def:warehouse-rates}
In mean-field steady state:
\begin{itemize}
  \item \textbf{Input rate} $\Lambda_W$: the rate at which tasks arrive
        at the warehouse when all deflection paths are saturated. The
        deflection protocol tries at most $h_{\max} = \lfloor N/2\rfloor$
        peers (\S\ref{sec:dag_kinetic}, Definition of the hop bound) before
        parking a chunk in the warehouse --- \emph{not} all $N-1$ peers.
        In mean-field, treating the $h_{\max}$ sampled peers as an
        independent sample (Stosszahlansatz), a worker deflects to the
        warehouse when all $h_{\max}$ tried peers have $q_j \geq H$
        simultaneously:
        \begin{equation}
          \Lambda_W \;=\; N\lambdazero\, \pd^{\lfloor N/2\rfloor},
          \label{eq:wh-input}
        \end{equation}
        where $\pd = \pi(H;\rhostar)$ is the per-worker probability of
        being above threshold; using $\lfloor N/2\rfloor \le N-1$ gives a
        strictly larger (more pessimistic, and more accurate) $\Lambda_W$
        than an all-$N-1$-peers model for $\pd<1,N>3$.
  \item \textbf{Drain rate} $\Lambda_D$: the rate at which workers pull
        tasks from the warehouse (during idle or below-threshold periods):
        \begin{equation}
          \Lambda_D \;=\; N\mu(1 - \pd).
          \label{eq:wh-drain}
        \end{equation}
\end{itemize}
\end{definition}

\subsection{Warehouse stability condition}

\begin{theorem}[Warehouse stability]
\label{thm:wh-stable}
The warehouse subsystem is stable (non-overflowing in expectation) if and
only if the net warehouse flux is non-positive:
\begin{equation}
  \Lambda_W < \Lambda_D
  \quad\Longleftrightarrow\quad
  \rhozero\,\pd^{\lfloor N/2\rfloor} < 1 - \pd.
  \label{eq:wh-stable}
\end{equation}
\end{theorem}

\begin{proof}
The warehouse occupancy $w(t)$ is a birth-death chain with input rate
$\Lambda_W$ and drain rate $\Lambda_D$.
It is stable (positive recurrent) iff $\Lambda_W < \Lambda_D$.
Dividing~\eqref{eq:wh-input} by $N\mu$ and using $\rhozero = \lambdazero/\mu$
gives~\eqref{eq:wh-stable}.
\end{proof}

\begin{proposition}[Relation to throughput stability]
\label{prop:l1-implies-wh}
For $\rhozero < \rhozero^{\dagger}$ where
\begin{equation}
  \rhozero^{\dagger}
  \;\coloneqq\;
  \frac{1-\pd^*(\rhozero^{\dagger})}{(\pd^*(\rhozero^{\dagger}))^{\lfloor N/2\rfloor}},
  \label{eq:rho-dagger}
\end{equation}
the warehouse stability condition~\eqref{eq:wh-stable} is satisfied.
For moderate load ($\pd \ll 1$), $\rhozero^{\dagger} \approx 1/\pd^{\lfloor N/2\rfloor} \gg 1$,
so the warehouse is stable for all $\rhozero < 1$ (i.e., Level-1 stability
implies warehouse stability).
At high load ($\pd \to 1$), $\rhozero^{\dagger} \to 0$, and the warehouse
becomes the bottleneck before the local queues saturate.
\end{proposition}

\begin{proof}
Substitute $\lambdazero = \rhozero\mu$ into~\eqref{eq:wh-stable} and
solve for $\rhozero$.  The self-consistency $\pd = \pd(\rhozero)$ is an
increasing function of $\rhozero$.  At $\rhozero = 0$, $\pd = 0$ and
the left side of~\eqref{eq:wh-stable} is 0, so the condition holds.
By continuity, it holds for all $\rhozero < \rhozero^\dagger$.
\end{proof}

\subsection{Mean time to crash}

\begin{theorem}[Mean time to warehouse overflow]
\label{thm:crash-time}
Under sustained overload ($\Lambda_W > \Lambda_D$), the warehouse
occupancy performs a biased random walk with net drift
$\Delta = \Lambda_W - \Lambda_D > 0$ (chunks per unit time).
Starting from occupancy $w_0 = 0$ (empty warehouse),
the mean first-passage time to overflow is
\begin{equation}
  T_{\mathrm{crash}}
  \;\approx\; \frac{C_W}{\Lambda_W - \Lambda_D}
  \;=\; \frac{C_W}{N(\rhozero\pd^{\lfloor N/2\rfloor}\mu - \mu(1-\pd))}.
  \label{eq:crash-time}
\end{equation}
Under stable operation ($\Lambda_W < \Lambda_D$), the warehouse is
positive recurrent and $T_{\mathrm{crash}} = \infty$.
\end{theorem}

\begin{proof}
For a birth-death chain with constant drift $\Delta > 0$ on
$\{0,\ldots,C_W\}$ starting at 0, the mean first-passage time
to $C_W$ is $T \approx C_W/\Delta$ by the optional stopping theorem
applied to the martingale $w(t) - \Delta t$.
Substituting $\Delta = \Lambda_W - \Lambda_D$ gives~\eqref{eq:crash-time}.
\end{proof}

\begin{corollary}[Benchmark crash time at near-saturation]
At $\rhozero = 0.95$ with $N=4$ ($\lfloor N/2\rfloor = 2$) and
$\pd = \pi(H;\rhostar)$ evaluated at the self-consistent $\rhostar$:
\[
  T_{\mathrm{crash}} \approx \frac{C_W}{N\mu(\rhozero\pd^2 - (1-\pd))}.
\]
For $\pd \ll 1$ (even at $\rhozero = 0.95$, $\pd$ is small),
the denominator $\approx N\mu\pd^2 \approx 0$, giving
$T_{\mathrm{crash}} \to \infty$ --- the warehouse essentially never overflows
at moderate per-worker deflection probability.
Overflow only becomes likely when $\pd$ is substantial, i.e., the entire
system is near saturation. (Using $\lfloor N/2\rfloor=2$ rather than $N-1=3$
makes this threshold easier to cross than a naive all-peers model would
suggest, though still requiring $\pd$ not small.)
\end{corollary}

\begin{remark}[Finite $T_{\mathrm{crash}}$ is not a failure mode at nominal load]
Equation~\eqref{eq:crash-time} gives a finite crash time only when
$\Lambda_W > \Lambda_D$, which requires $\pd^{\lfloor N/2\rfloor} > (1-\pd)/\rhozero$.
At benchmark parameters with $\rhozero \leq 0.95$, $\pd^2 \ll (1-\pd)$
for all $\pd < 0.5$, so the system is stable and $T_{\mathrm{crash}} = \infty$
in steady state.
The crash time formula is relevant only for diagnosing behavior under
deliberate overload (e.g., load testing above capacity).
\end{remark}

% =============================================================
\section{Synthesis: The Stability Ladder}
% =============================================================

\begin{theorem}[Stability hierarchy]
\label{thm:hierarchy}
The four stability properties are ordered:
\[
  \underbrace{\rhozero < 1}_{\text{L1: throughput stable}}
  \;\Longrightarrow\;
  \underbrace{\gamma > 0}_{\text{L2: exponential mixing}}
  \;\Longrightarrow\;
  \underbrace{B^* > 0}_{\text{L3: burst absorbing}}
  \;\Longrightarrow\;
  \underbrace{T_{\mathrm{crash}} = \infty}_{\text{L4: crash-free}}.
\]
\end{theorem}

\begin{proof}
$\text{L1} \Rightarrow \text{L2}$:
Positive recurrence and irreducibility of a finite CTMC imply a
unique stationary distribution.
By the Perron--Frobenius theorem, the generator has a spectral gap
$\gamma > 0$ (the second eigenvalue is strictly negative).
Theorem~\ref{thm:gap} gives the explicit value.

$\text{L2} \Rightarrow \text{L3}$:
Exponential mixing with rate $\gamma$ implies the stationary distribution
$f_1^*$ exists and has finite support.
The burst capacity $B^* = N(L-\qbar) + C_WK > 0$ follows from
$\qbar < L$ (guaranteed by $\rhozero < 1$ and the truncation at $L$).

$\text{L3} \Rightarrow \text{L4}$:
Under $\rhozero < 1$ and Proposition~\ref{prop:l1-implies-wh},
$\Lambda_W < \Lambda_D$, so the warehouse is positive recurrent
and $T_{\mathrm{crash}} = \infty$.
\end{proof}

\medskip
\noindent The stability properties and their quantitative values are
summarized in Table~\ref{tab:summary}.

\begin{table}[h]
\centering
\caption{Stability properties of DTA at benchmark parameters
($N=4$, $L=131072$, $H=114688$, $C_W=32768$, $K=32$, $\delta=1$).}
\label{tab:summary}
\begin{tabular}{@{}lll@{}}
\toprule
Level & Condition & Quantitative value \\
\midrule
L1: Throughput stable    & $\rhozero < 1$
  & Stability region: $\lambdazero \in [0, \mu)$ \\
L2: Exponential mixing   & $\gamma = \mu(1-\sqrt{\rhostar})^2$
  & At $\rhozero=0.7$: $\gamma \approx 0.025\mu$ \\
L3: Burst absorption     & $B^* = N(L-\qbar) + C_WK$
  & At $\rhozero=0.7$: $B^* \approx 1{,}252{,}000$ tasks \\
L4: Crash-free threshold & $T_{\mathrm{crash}} = \infty$ iff $\Lambda_W < \Lambda_D$
  & Satisfied for all $\rhozero < \rhozero^\dagger \approx 1$ \\
\bottomrule
\end{tabular}
\end{table}

% =============================================================

\begin{figure}[htbp]
\centering
\begin{minipage}{0.48\linewidth}\centering
\includegraphics[width=\linewidth]{figure/stability_L1_mean_queue.png}
\caption{Level-1: Mean queue length $\bar{q}$ versus offered load $\varrho_0$. Divergence at $\varrho_0=1$ confirms Foster--Lyapunov stability boundary.}\label{fig:stab_L1}
\end{minipage}\hfill
\begin{minipage}{0.48\linewidth}\centering
\includegraphics[width=\linewidth]{figure/stability_L2_spectral_gap.png}
\caption{Level-2: Spectral gap $\gamma$ of the M/M/1/$H$ generator versus $\varrho_0$. Critical slowing-down $\gamma\to 0$ as $\varrho_0\to 1^-$ is visible; analytical formula overlaid.}\label{fig:stab_L2}
\end{minipage}
\end{figure}

\begin{figure}[htbp]
\centering
\begin{minipage}{0.48\linewidth}\centering
\includegraphics[width=\linewidth]{figure/stability_L3_burst_capacity.png}
\caption{Level-3: Adversarial burst capacity $B^*$ versus $\varrho_0$. Monotone decrease and the warehouse bonus $C_W K$ are both visible.}\label{fig:stab_L3}
\end{minipage}\hfill
\begin{minipage}{0.48\linewidth}\centering
\includegraphics[width=\linewidth]{figure/stability_L4_crash_time.png}
\caption{Level-4: First-passage crash time $T_{\mathrm{crash}}$ versus warehouse overload $\varrho_0 > 1$. The $1/\Delta$ scaling predicted by Theorem~4.1 is confirmed.}\label{fig:stab_L4}
\end{minipage}
\end{figure}
\section{Numerical Validation}
% =============================================================

The script \texttt{simulate\_stability.py} validates all four levels:
\begin{enumerate}
  \item \textbf{L1}: Sweeps $\rhozero \in [0, 1.1]$, confirms mean queue
        diverges at $\rhozero = 1$ and is bounded below.
  \item \textbf{L2}: Computes eigenvalues of the M/M/1/$H$ generator
        numerically; plots spectral gap $\gamma$~\cite{diaconis_stroock1991,fill1991} vs $\rhozero$ and
        confirms the critical slowing-down $\gamma \to 0$ as $\rhozero \to 1$.
  \item \textbf{L3}: Computes $B^*(\rhozero)$ analytically and plots
        versus $\rhozero$; checks monotone decrease.
  \item \textbf{L4}: Simulates the warehouse as a 1D birth-death chain;
        measures $T_{\mathrm{crash}}$ at various overload levels above $\rhozero = 1$;
        confirms the $1/\Delta$ scaling of Theorem~\ref{thm:crash-time}.
\end{enumerate}
\endgroup

\bigskip
\noindent\textbf{Summary.}
The five performance chapters collectively characterise DTA quantitatively.
The mean-field load-balance analysis shows deflection drives imbalance to
$O(1/\sqrt{N})$ under Poisson arrivals; the BBGKY correction quantifies
the $O(1/N)$ error from finite-$N$ pair correlations.
The makespan bound decomposes runtime into a deterministic work-conservation
term and a Gumbel extreme-value fluctuation term.
The stability ladder establishes four nested regimes from throughput
stability to first-passage crash safety.
The DAG kinetic-theory section extends these results to fork-join graphs,
deriving mean-field makespan bounds and identifying a directed-percolation
phase transition as an open problem.
Across all five chapters, every quantitative bound is derived from
routing decisions that use only local queue occupancy and single-hop
deflection; the recurring $O(1/\sqrt{N})$-with-$O(1/N)$-correction
pattern is the quantitative signature of bounded-information scheduling
losing nothing at leading order.


\clearpage

\partbreak
\part{Competitive and Empirical Analysis}
\clearpage

\bigskip
\noindent\textbf{Overview.}
This part places DTA in its competitive and empirical context.
The competitive analysis~\cite{borodin_elyaniv1998,borodin1996} derives online competitive ratios for Bag-of-Tasks
workloads relative to work-stealing, including an adversarial worst-case
bound whose \emph{proof} requires no mean-field assumption on the task
sequence itself (Theorem~\ref{thm:cr_bot_worst} holds pathwise for every
arrival pattern); its precondition $\varrho_0 < 1$ (system stability)
is, however, established via the mean-field stability theorem of
\hyperref[chap:stability]{Performance~IV}, so the overall result is not
entirely assumption-free.
The comparison chapter benchmarks DTA against a centralised queue baseline
and work-stealing across the four stability metrics.
The empirical chapter validates every theoretical bound against
wall-clock Criterion measurements from the public GitHub Actions CI pipeline,
and extrapolates to production-scale parameter regimes.


% ========================================================================
% Competitive Ratio Analysis
% ========================================================================

\section*{Competitive Ratio Analysis}
\addcontentsline{toc}{section}{Competitive Ratio Analysis}
\label{chap:competitive_ratio}
\setcounter{section}{0}

\textbf{Notation.}
Throughout, $N$ denotes the number of workers, $\mu$ the service rate,
$\varrho_0 = \lambda_0/\mu$ the offered load, $\varrho^*$ the self-consistent
load (see \hyperref[chap:formal_model]{the Formal Model}), $H$ the per-worker queue capacity,
$p_d = \mathbb{P}(q^* \ge H)$ the mean-field deflection probability,
$\delta_{\mathrm{hop}}$ the latency of one SPSC mailbox message, and
$\delta = 1/\mu$ the mean task processing time.

% ===========================================================================
\section{Online Scheduling Model and Competitive Ratio}
% ===========================================================================

\subsection{Task model}

Tasks arrive as a Poisson process of rate $\lambda_0$ and carry i.i.d.\
service demands $s_i \sim \mathrm{Exp}(\mu)$ (unless stated otherwise).
A \emph{task topology} specifies possible dependency structures:

\begin{definition}[Task topologies]
\begin{enumerate}
  \item \textbf{Bag-of-tasks (BoT):} $M$ mutually independent tasks,
        no ordering constraints.  This is DTA's primary design target.
  \item \textbf{Heavy-tail (HT):} tasks i.i.d.\ $\mathrm{Pareto}(\alpha, s_{\min})$
        with $\alpha > 1$ (finite mean) but possibly $\alpha \le 2$ (infinite variance).
  \item \textbf{Fork-join DAG:} task graph $G=(V,E)$; a task $v$ may not
        start until all predecessors complete.  Critical-path length
        $T_\infty = $ longest path in $G$; total work $T_1 = \sum_{v} s_v$.
\end{enumerate}
\end{definition}

\subsection{Offline optimum and competitive ratio}

\begin{definition}[Competitive ratio]
Let $\sigma = (s_1, \ldots, s_M)$ be a task sequence and
$\mathcal{A}$ an online scheduler.  Define
\[
  \varrho_c(\mathcal{A})
  \;=\;
  \sup_{\sigma}\;
  \frac{C_{\max}^{\mathcal{A}}(\sigma)}{C_{\max}^{\mathrm{OPT}}(\sigma)},
\]
where $C_{\max}^{\mathrm{OPT}}(\sigma)$ is the makespan of the \emph{clairvoyant}
offline optimum (knows all $s_i$ before scheduling begins).
A scheduler is \emph{$c$-competitive} if $\varrho_c \le c$.
\end{definition}

\begin{remark}[Physical reading]
The competitive ratio measures the \physics{cost of irreversibility}{price
of online: not knowing the future}.  The OPT plays the role of a
\physics{reversible (quasi-static) process}{Carnot engine}: it achieves
maximum ``thermodynamic'' efficiency with full information.
DTA is the corresponding irreversible process; the gap $\varrho_c - 1$
is the \physics{entropy production overhead}{scheduling inefficiency due
to causal constraints}.
\end{remark}

% ===========================================================================
\section{Path-Length Analysis}
% ===========================================================================

The central physical tool is the \physics{mean free path}{expected number of
deflection hops per task}~\cite{lifshitz_kinetics,chapman_cowling}.  In kinetic theory (see
\hyperref[chap:kinetic_theory]{Performance~II}), a particle travels distance
$\lambda_{\mathrm{mfp}} = v/\nu_{\mathrm{coll}}$ between collisions
($v$ = speed, $\nu_{\mathrm{coll}}$ = collision frequency).  In DTA,
a task ``scatters'' off a full worker and hops to the next peer on the ring.

\begin{lemma}[Path-length distribution]\label{lem:path_length}
Under the mean-field approximation, successive workers visited by a task
are i.i.d.\ Bernoulli: each is full (deflects) with probability $p_d$
independently.  The number of hops $H_{\mathrm{task}}$ before acceptance
follows a geometric distribution:
\[
  \mathbb{P}(H_{\mathrm{task}} = k) \;=\; (1 - p_d)\, p_d^k,
  \quad k = 0, 1, \ldots, \lfloor N/2 \rfloor.
\]
Consequently,
\[
  \mathbb{E}[H_{\mathrm{task}}] \;=\; \frac{p_d}{1 - p_d},
  \qquad
  \mathrm{Var}[H_{\mathrm{task}}] \;=\; \frac{p_d}{(1-p_d)^2}.
\]
\end{lemma}

\begin{proof}
Each hop independently tests an i.i.d.\ draw from $f_1^*(q)$ (mean-field
stationary density from \hyperref[chap:kinetic_theory]{Performance~II}, Level-0 closure).
Under Stosszahlansatz ($f_2 = f_1 \otimes f_1$), successive workers are
statistically independent, and each accepts the task with probability
$1 - p_d$.  The stopping time is therefore geometric.
\end{proof}

\begin{remark}[Physical analogy]
$\mathbb{E}[H_{\mathrm{task}}]$ is the \physics{mean free path in hop units}
{expected routing hops}.  The analogy table is:
\medskip

\noindent
\begin{tabularx}{\textwidth}{@{}clXl@{}}
\toprule
Kinetic theory & DTA \\
\midrule
Mean free path $\lambda_{\mathrm{mfp}}$ & $\mathbb{E}[H_{\mathrm{task}}]\cdot d_{\mathrm{ring}}$ \\
Scattering cross-section $\sigma_{\mathrm{scat}}$ & Deflection probability $p_d$ \\
Ballistic transport (OPT) & Direct assignment, 0 hops \\
Diffusive transport (DTA) & Hop-bounded random walk \\
Mean-field closure (Boltzmann) & Stosszahlansatz $f_2 = f_1 \otimes f_1$ \\
\bottomrule
\end{tabularx}
\end{remark}

\begin{proposition}[Routing overhead]\label{prop:routing_overhead}
For a batch of $M$ tasks, the total latency added by deflection hops is
\[
  \Delta C_{\max}^{\mathrm{routing}}
  \;\le\;
  \frac{M}{N} \cdot \frac{p_d}{1 - p_d} \cdot \delta_{\mathrm{hop}},
\]
where $\delta_{\mathrm{hop}}$ is the latency of one SPSC mailbox write.
\end{proposition}

\begin{proof}
Total hop messages = $M \cdot \mathbb{E}[H_{\mathrm{task}}] = M p_d/(1-p_d)$.
These are distributed across $N$ workers; the bottleneck worker sees at most
$M/N$ tasks and $M p_d / (N(1-p_d))$ extra hops, each costing $\delta_{\mathrm{hop}}$.
The bound follows from linearity and the mean-field i.i.d.\ assumption.
\end{proof}

% ===========================================================================
\section{Bag-of-Tasks: Competitive Ratio}
% ===========================================================================

\begin{theorem}[DTA competitive ratio, BoT]\label{thm:cr_bot}
For i.i.d.\ $\mathrm{Exp}(\mu)$ tasks on $N$ workers with queue capacity $H$,
DTA achieves competitive ratio
\begin{equation}\label{eq:cr_main}
  \varrho_c^{\mathrm{DTA}}
  \;\le\;
  \underbrace{\left(1 + \frac{p_d}{1-p_d}\cdot \delta_{\mathrm{hop}}\mu\right)}_{\text{routing factor}}
  \;\times\;
  \underbrace{\left(1 + \frac{\mathbb{E}[M_N] - \bar q}{\bar q}\right)}_{\text{imbalance factor}},
\end{equation}
where $\bar q = \varrho^*/(1-\varrho^*)$ is the mean-field queue mean,
$\mathbb{E}[M_N]$ the expected maximum queue (see
\hyperref[chap:makespan]{Performance~III}, Proposition~3.2), and
$C_{\max}^{\mathrm{OPT}} = \bar q / \mu$ for the same offered load.
\end{theorem}

\begin{proof}
Decompose $C_{\max}^{\mathrm{DTA}} = C_{\max}^{\mathrm{balance}} + \Delta C_{\max}^{\mathrm{routing}}$.
The balance term gives
$C_{\max}^{\mathrm{balance}} / C_{\max}^{\mathrm{OPT}} \le \mathbb{E}[M_N]/\bar q$,
yielding the imbalance factor.
The routing overhead is bounded by Proposition~\ref{prop:routing_overhead},
giving the routing factor.  Multiplying the two factors
(both $\ge 1$) and noting $C_{\max}^{\mathrm{OPT}} \ge \bar q/\mu$
completes the proof.
\end{proof}

\begin{corollary}[Large-$H$ limit]\label{cor:large_H}
For M/M/1/H with $H \to \infty$ at fixed $\varrho_0 < 1$,
the mean-field stationary mass at $q = H$ satisfies
$p_d = \mathcal{O}(\varrho^{*H}) \to 0$ exponentially.  Hence
\[
  \varrho_c^{\mathrm{DTA}} \;\xrightarrow{H\to\infty}\; 1
  \;+\; O\!\left(\frac{\mathbb{E}[M_N] - \bar q}{\bar q}\right).
\]
The imbalance factor alone determines the competitive ratio. Unlike a
Gaussian-tailed maximum (which would grow as $O(\sqrt{\ln N})$), the
truncated-geometric tail $\overline{F_1^*}(m) \sim (1-\varrho^*)\varrho^{*m}$
is in the \emph{Gumbel} (exponential-tail) domain of attraction, so by
the paper's own Gumbel asymptotics
(\hyperref[chap:makespan]{Performance~III}, \S\ref{sec:gumbel}),
$\mathbb{E}[M_N] = b_N + a_N\gamma$ with $a_N = -1/\ln\varrho^*$ and
$b_N = \bigl[\ln N + \ln(1-\varrho^*)\bigr]/(-\ln\varrho^*)$, i.e.
\begin{equation}\label{eq:cr_large_H}
  \mathbb{E}[M_N] - \bar q
  \;=\;
  \frac{\ln N + \ln(1-\varrho^*) + \gamma}{-\ln\varrho^*} - \frac{\varrho^*}{1-\varrho^*}
  \;=\; O\!\left(\frac{\ln N}{-\ln\varrho^*}\right)
  \quad\text{as } N \to \infty,
\end{equation}
so that
\[
  \varrho_c^{\mathrm{DTA}} \;\le\; 1 + O\!\left(\frac{\ln N}{-\bar q\,\ln\varrho^*}\right).
\]
For heavy-traffic $\varrho^* \to 1$: $\bar q = \varrho^*/(1-\varrho^*) \to
\infty$ while $-\ln\varrho^* \to 0$ (so the denominator grows like
$\bar q \cdot (1-\varrho^*)^{-1}$, i.e.\ like $\bar q^2$, which dominates
the numerator's logarithmic growth in $N$ for any fixed $N$), giving
$\varrho_c^{\mathrm{DTA}} \to 1$.  DTA is \textbf{asymptotically optimal}
under heavy load, i.e.\ as $\varrho^* \to 1$ \emph{at fixed, or even slowly
growing, $N$} --- this bound is \emph{not}, by itself, informative at
moderate load ($\varrho^*$ bounded away from $1$) and large $N$, where
$\bar q$ stays small while $\ln N$ grows; see
Corollary~\ref{cor:benchmark} for a worked numerical example of this
distinction.
\end{corollary}

\begin{corollary}[Target deployment scale]\label{cor:benchmark}
With DTA's \emph{target deployment} parameters ($N = 256$, $H = 114{,}688$,
$\delta_{\mathrm{hop}}/\delta \le 10^{-3}$, $\varrho_0 = 0.7$) --- as
opposed to the $N=4$ scale actually exercised by the published Criterion
CI benchmarks (\S\ref{sec:validity}) --- we obtain:
\[
  p_d \;\le\; \frac{1}{H+1} \;=\; 8.7 \times 10^{-6},
  \qquad
  \text{routing factor} \;\le\; 1 + 10^{-3} \cdot p_d / (1-p_d)
  \;\approx\; 1 + 8.7\times 10^{-9}.
\]
The \emph{imbalance} factor, however, is governed by
Equation~\eqref{eq:cr_large_H}, and at $\varrho^* \approx 0.7$ the system
is not yet in the heavy-traffic regime that corollary requires: with
$\bar q = \varrho^*/(1-\varrho^*) \approx 2.33$,
$-\ln\varrho^* \approx 0.357$, and $N=256$,
\[
  \mathbb{E}[M_N] - \bar q
  \;\approx\;
  \frac{\ln 256 + \ln(0.3) + \gamma}{0.357} - 2.33
  \;\approx\; 13.8 - 2.33 \;\approx\; 11.5,
\]
so the imbalance factor is $1 + 11.5/2.33 \approx 5.9$, giving
$\varrho_c^{\mathrm{DTA}} \lesssim 5.9$ at $\varrho_0 = 0.7$, $N=256$ ---
\emph{not} close to $1$. This is not a contradiction of
Corollary~\ref{cor:large_H}: that corollary's asymptotic requires
$\varrho^* \to 1$ specifically (equivalently $\bar q \to \infty$), and at
$\varrho_0 = 0.7$, $\bar q$ is still only $2.33$, far short of the
regime where the $\ln N$ growth of the Gumbel excess is dominated by
$\bar q$. The same calculation at higher load ($\varrho_0=0.95$,
$\bar q \approx 19$) gives imbalance factor $\approx 3.2$ --- still not
$\lesssim 1.12$, though substantially improved, illustrating the slow
(logarithmic in $N$, linear in $\bar q$) approach to the heavy-traffic
limit. At
the $N=256$ target deployment scale and moderate-to-high load, the
informative bound on DTA's competitive ratio is therefore not this
large-$H$ asymptotic but the load-independent, distribution-free
worst-case bound of Theorem~\ref{thm:cr_bot_worst}
($\varrho_c^{\mathrm{DTA,worst}} \leq 1 + N\lfloor N/2\rfloor
\delta_{\mathrm{hop}}/W \to 1$ as total work $W \to \infty$, independent
of $\varrho^*$), which does not suffer from this small-$\bar q$
limitation.
\end{corollary}

% ===========================================================================
\section{Topology-Specific Bounds}
% ===========================================================================

\subsection{Heavy-tail task sizes}

\begin{proposition}[Heavy-tail degradation]\label{prop:heavy_tail}
Let task sizes be i.i.d.\ $\mathrm{Pareto}(\alpha, s_{\min})$ with
$\alpha > 2$ (finite variance $\sigma_s^2$) and coefficient of variation
$\mathrm{CV} = \sigma_s / \mathbb{E}[s]$.  Then
\[
  \varrho_c^{\mathrm{DTA,HT}}
  \;\le\;
  \varrho_c^{\mathrm{DTA,BoT}} \;+\; \frac{1 + \mathrm{CV}^2}{N},
\]
where the additive term is the \emph{variance penalty}:
even OPT cannot perfectly balance when task sizes are random,
and DTA incurs the same fundamental variance lower bound plus
an additional $\mathrm{CV}^2/N$ load-imbalance term.
\end{proposition}

\begin{proof}[Proof sketch]
By the analysis of \cite{graham1969,coffman_graham1972}, the \emph{best} online algorithm
for $N$ machines has competitive ratio $\ge 1 + (1-1/N)\cdot\sigma_s^2 / \mathbb{E}[s]^2 / N$
for i.i.d.\ random job sizes.  DTA matches this lower bound up to
$O(p_d)$ corrections.  For $\alpha \le 2$ (infinite variance), $\mathrm{CV} \to \infty$
and the bound becomes vacuous; the competitive ratio is determined by
the largest task (\emph{whale}), which any online algorithm must accept greedily.
\end{proof}

\begin{remark}[Physical analogy: disorder]
Heavy-tail tasks correspond to \physics{disordered systems with Griffiths singularities}
{rare but catastrophically large tasks dominating makespan}.
The ``whale'' task is the analogue of a rare large cluster in a random ferromagnet
that creates anomalously slow dynamics near the phase transition.
No local load-balancing strategy (DTA or WS) can avoid this fundamental limit.
\end{remark}

\subsection{Fork-join DAG tasks}

DTA has no native dependency-tracking mechanism: it routes tasks
as independent units and cannot enforce that a task's predecessors
have completed.  For DAG topologies, the relevant scheduler is
Work-Stealing (see Section~\ref{sec:ws_compare}).

For reference, the information-theoretic lower bound on makespan for
any scheduler on a fork-join DAG is Graham's bound:
\begin{equation}\label{eq:graham}
  C_{\max}^{\mathrm{OPT}} \;\ge\; \max\!\left(\frac{T_1}{N},\; T_\infty\right),
\end{equation}
where $T_1$ is total work and $T_\infty$ the critical-path length.
Work-Stealing achieves $\mathbb{E}[C_{\max}^{\mathrm{WS}}] = T_1/N + O(T_\infty)$
\cite{blumofe1999}, which is asymptotically optimal when $T_1 \gg N T_\infty$
(high parallelism).

% ===========================================================================
\section{Work-Stealing Comparison: Push/Pull Duality}
\label{sec:ws_compare}
% ===========================================================================

\subsection{Structural duality}

DTA and Work-Stealing (WS) are \emph{mechanistic duals}:

\medskip
\noindent
\begin{tabularx}{\textwidth}{@{}clXl@{}}
\toprule
Property & DTA & Work-Stealing \\
\midrule
Direction & \textbf{Push}: task deflected at arrival & \textbf{Pull}: idle worker steals \\
Trigger & Arrival event ($\lambda_0$-rate) & Worker becomes idle \\
Action range & Bounded ($\le \lfloor N/2 \rfloor$ hops) & Unbounded (any victim) \\
Communication & SPSC per-pair mailbox & Shared Chase--Lev deque + CAS \\
Mean free path & $\mathbb{E}[H_{\mathrm{task}}] = p_d/(1-p_d)$ & \textbf{None}: theft is a global event \\
Design topology & Bag-of-tasks, $\varrho_0 \to 1$ & Fork-join DAG, $\varrho_0 \to 0$ \\
Formal result & Theorem~\ref{thm:cr_bot} (\hyperref[chap:kinetic_theory]{this section}) & \cite{blumofe1999} \\
\bottomrule
\end{tabularx}

\medskip

The mean-free-path concept does \emph{not} apply to WS because theft
is a global operation (any idle worker can steal from any busy worker):
it corresponds to \physics{infinite-range interactions}{global coordination},
analogous to mean-field theory \emph{globally} rather than locally.
DTA's hop bound creates \physics{short-range interactions}{locality
constraint}, analogous to a lattice model with finite interaction radius.

\subsection{Competitive ratio in the overlap regime}

The overlap regime is \textbf{i.i.d.\ BoT tasks at intermediate load},
where both schedulers are applicable.

\begin{theorem}[WS competitive ratio, BoT]\label{thm:cr_ws_bot}
For i.i.d.\ $\mathrm{Exp}(\mu)$ tasks on $N$ workers with WS
(Chase--Lev deque, CAS cost $\delta_{\mathrm{steal}}$),
the expected number of steals for $M$ tasks is
$S_M = O(N \cdot (1 - \varrho^*) \cdot M / N) = O(M(1-\varrho^*))$,
and the competitive ratio satisfies
\begin{equation}\label{eq:cr_ws}
  \varrho_c^{\mathrm{WS}}
  \;\approx\;
  1 \;+\; (1 - \varrho^*)\,\delta_{\mathrm{steal}}\,\mu.
\end{equation}
\end{theorem}

\begin{proof}[Proof sketch]
An idle worker triggers one steal.  In steady state, the fraction of
idle time is $\approx (1-\varrho^*)$, giving steal rate $\approx \mu(1-\varrho^*)$
per worker, total $N\mu(1-\varrho^*)$.  Each steal adds latency $\delta_{\mathrm{steal}}$
(one CAS on the victim's deque).  This overhead adds
$N \cdot (1-\varrho^*)\delta_{\mathrm{steal}}$ per unit time, giving
the stated competitive ratio.
\end{proof}

\begin{theorem}[Regime comparison]\label{thm:regime_compare}
Let $r = \delta_{\mathrm{hop}}/\delta_{\mathrm{steal}} < 1$ (SPSC write cheaper
than CAS).  Then:
\begin{enumerate}
  \item \textbf{High load} ($\varrho_0 \to 1$):
        $\varrho_c^{\mathrm{DTA}} \to 1$ (since $p_d \to p_d^{\infty} \ll 1$)
        and $\varrho_c^{\mathrm{WS}} \to 1$ simultaneously.
        Both are asymptotically optimal; DTA has lower overhead since
        $p_d \cdot \delta_{\mathrm{hop}} < (1-\varrho^*)\delta_{\mathrm{steal}}$
        near saturation.

  \item \textbf{Low load} ($\varrho_0 \to 0$):
        $p_d \to \varrho^{*H} \to 0$, so $\varrho_c^{\mathrm{DTA}} \to 1$.
        WS overhead $(1-\varrho^*)\delta_{\mathrm{steal}}\mu \to \delta_{\mathrm{steal}}\mu$,
        constant and positive.  DTA is better for BoT at low load.

  \item \textbf{DAG tasks}:
        DTA's competitive ratio is undefined (no dependency enforcement).
        WS achieves $T_1/N + O(T_\infty)$ \cite{blumofe1999}.
        WS is the only viable choice for DAG topologies.
\end{enumerate}
\end{theorem}

\begin{remark}[Physical interpretation]
DTA's hop-bounded push is analogous to
\physics{pressure-driven diffusion}{task routing driven by queue pressure}.
WS's pull mechanism is analogous to
\physics{convection}{suction: idle workers actively drawing in tasks}.
For independent tasks at high load, diffusion matches or exceeds convection
efficiency because the ``pressure gradient'' (queue imbalance) is the natural
driving force; convective suction adds overhead without benefit.
For DAGs, convection (WS) is necessary because the ordering constraints
create topological ``channels'' that diffusion cannot exploit.
\end{remark}

\begin{remark}[Scope: this theorem compares overhead terms, not the full ratio]
\label{rem:regime_compare_scope}
Theorem~\ref{thm:regime_compare} compares only the \emph{routing/stealing-overhead}
components of $\varrho_c^{\mathrm{DTA}}$ and $\varrho_c^{\mathrm{WS}}$
($p_d\delta_{\mathrm{hop}}$ vs.\ $(1-\varrho^*)\delta_{\mathrm{steal}}$); it does
not claim the \emph{full} competitive ratio of Theorem~\ref{thm:cr_bot}
(routing $\times$ imbalance) converges to $1$ at fixed $H$ --- that requires
$H\to\infty$ too (Corollary~\ref{cor:large_H}), and the imbalance factor alone
can remain large ($\gtrsim3$--$6$, Corollary~\ref{cor:benchmark}). Accordingly,
Figure~\ref{fig:cr_vs_rho}'s full empirical ratio tracks WS/CQ rather than
beating them outright: ``DTA dominates at all loads'' in item~1 above is an
overhead-only statement.
\end{remark}

\subsection{The CAS-Storm Paradox at Low Load}
\label{subsec:cas_storm}

Theorem~\ref{thm:regime_compare} states that at low load ($\varrho_0 \to 0$),
DTA's competitive ratio approaches 1 ($p_d \to 0$, no deflection overhead),
while WS's ratio approaches $1 + \delta_{\mathrm{steal}}\mu > 1$.
A subtler practical phenomenon reinforces this advantage: the
\emph{CAS-storm paradox}.

\begin{proposition}[CAS-storm overhead at low load]\label{prop:cas_storm}
Let $N$ workers run WS with idle-loop stealing (each idle worker repeatedly
calls \texttt{steal()} on a randomly chosen victim).
At steady-state load $\varrho_0$, the expected number of CAS operations
per task is
\begin{equation}\label{eq:cas_per_task}
  \Gamma_{\mathrm{WS}}(\varrho_0)
  \;=\;
  \frac{\mu\,(1-\varrho_0)\,N}{\lambda_0}
  \;=\;
  \frac{N(1-\varrho_0)}{\varrho_0},
\end{equation}
which \textbf{diverges as} $\varrho_0 \to 0$.  Each CAS carries cost
$\delta_{\mathrm{CAS}}$ (a failed CAS on a shared cache line also
propagates an invalidation to all $N-1$ other caches, paying a
\physics{coherence penalty}{cache-coherence broadcast}).
The total per-task WS overhead at low load is
\begin{equation}\label{eq:ws_low_load_overhead}
  \Delta_{\mathrm{WS}}(\varrho_0)
  \;\approx\;
  \frac{N(1-\varrho_0)}{\varrho_0} \cdot (\delta_{\mathrm{CAS}} + \delta_{\mathrm{coh}}),
\end{equation}
where $\delta_{\mathrm{coh}} \approx (N-1)\cdot\delta_{\mathrm{MESI}}$ is
the cache-coherence round-trip cost of a single failed CAS (MESI protocol
invalidation broadcast to the other $N-1$ caches); the $(N-1)$ factor
enters only here, per-failed-CAS, not as a multiplier on the number of
failed CAS operations itself.
\end{proposition}

\begin{proof}
At load $\varrho_0$, the fraction of time each worker is idle $= 1-\varrho_0$.
An idle worker calls \texttt{steal()} at rate $\approx 1/\delta_{\mathrm{CAS}}$.
Total steal attempts per unit time across $N$ idle workers $\approx N(1-\varrho_0)/\delta_{\mathrm{CAS}}$.
Successful steals per unit time $= \lambda_0$.
Failed steals per unit time $\approx N(1-\varrho_0)/\delta_{\mathrm{CAS}} - \lambda_0$;
per task: $\Gamma_{\mathrm{WS}} \approx N(1-\varrho_0)/\varrho_0$.
\end{proof}

\begin{remark}[The paradox]
This is the \emph{CAS-storm paradox}: WS's theoretical competitive ratio
approaches 1 at $\varrho_0 \to 0$ (Theorem~\ref{thm:cr_ws_bot}), yet its
\emph{practical} per-task overhead \emph{diverges} in the same limit due to
idle-loop contention.  The theoretical analysis accounts only for steals that
\emph{succeed} in rebalancing load; it ignores the $O(N/\varrho_0)$
failed CAS operations and their coherence broadcasts.

DTA has no such pathology: at $\varrho_0 \to 0$, $p_d \to 0$, the SPSC
mailboxes carry zero traffic, and the scheduling overhead vanishes.
This \physics{false vacuum of efficiency}{theoretical optimality masking
practical overhead} makes WS appear superior in theory at low load while
DTA is faster in practice --- as confirmed by the Criterion benchmarks
in \hyperref[chap:empirical]{the Empirical Analysis} (Spawn Efficiency, 1k tasks:
DTA $161\,\mu\mathrm{s}$ vs.\ Tokio $699\,\mu\mathrm{s}$, a \textbf{4.3$\times$}
practical advantage at low load).
\end{remark}

% ===========================================================================

\begin{figure}[htbp]
\centering
\includegraphics[width=0.85\linewidth]{figure/competitive_ratio_vs_rho.png}
\caption{Empirical competitive ratio $\varrho_c$ (batch LPT-vs-\{DTA, Work-Stealing,
  Central Queue, Random\}, $N=8$, $\mathrm{Exp}(\mu)$ and $\mathrm{Pareto}(2.5)$ task
  sizes, 60 Monte Carlo batches per point; the deflection watermark $H$ and batch
  size $M$ scale with $\varrho_0$ so that deflection is non-trivial), versus the
  routing/stealing overhead-only terms of Theorem~\ref{thm:cr_bot} and
  Theorem~\ref{thm:cr_ws_bot} (dotted). DTA tracks Work-Stealing and Central Queue
  closely and stays well below Random; see the discussion below and
  Remark~\ref{rem:regime_compare_scope} for why the solid curves sit above the
  overhead-only dotted lines and for the scope of the ``DTA dominates'' claim.}
\label{fig:cr_vs_rho}
\end{figure}
\section{Physical Synthesis: Transport Efficiency}
% ===========================================================================

Define the \physics{transport efficiency}{scheduling efficiency relative to OPT}:
\[
  \eta \;=\; \frac{1}{\varrho_c} \;\in\; (0, 1].
\]
Figure~\ref{fig:cr_vs_rho} (see \texttt{simulate\_competitive.py}) plots the
empirical competitive ratio $\varrho_c = 1/\eta$ as a function of $\varrho_0$
for DTA, Work-Stealing, Central Queue, and Random on BoT tasks, together with
the overhead-only terms of Theorems~\ref{thm:cr_bot} and~\ref{thm:cr_ws_bot}.

The key qualitative features from Theorems \ref{thm:cr_bot}--\ref{thm:regime_compare}
concern the \emph{overhead-only} terms specifically (Remark~\ref{rem:regime_compare_scope}):
\begin{enumerate}
  \item DTA efficiency $\eta^{\mathrm{DTA}}$ is \textbf{monotone increasing} in $\varrho_0$
        (heavy-traffic optimality, Corollary~\ref{cor:large_H}).
  \item WS efficiency $\eta^{\mathrm{WS}}$ is also increasing in $\varrho_0$
        but has a \textbf{fixed overhead floor} $\delta_{\mathrm{steal}}\mu$ at $\varrho_0 = 0$.
  \item The crossover $\eta^{\mathrm{DTA}} = \eta^{\mathrm{WS}}$ occurs at
        $p_d \cdot \delta_{\mathrm{hop}} = (1-\varrho^*)\delta_{\mathrm{steal}}$,
        i.e., roughly $\varrho_0 \approx 1 - r\cdot p_d/(1-p_d)$.
        For $H$ large ($p_d \approx 0$), DTA's routing-overhead term is smaller
        than WS's stealing-overhead term at \emph{all} loads (Remark~\ref{rem:regime_compare_scope}
        --- this is an overhead-only statement, not a claim that DTA's full
        competitive ratio is uniformly below WS's; see Figure~\ref{fig:cr_vs_rho}).
  \item For heavy-tail tasks, both efficiencies degrade by the variance penalty $1/(1+\mathrm{CV}^2/N)$.
\end{enumerate}


% ===========================================================================
\section{Adversarial Bag-of-Tasks: Worst-Case Competitive Ratio}
\label{sec:cr_bot_worst}
% ===========================================================================

\begin{remark}[Scope of previous results]
Theorem~\ref{thm:cr_bot} and Corollaries~\ref{cor:large_H}--\ref{cor:benchmark}
are \emph{average-case} bounds: they rely on the mean-field stationary
distribution of i.i.d.\ Poisson arrivals and exponential service times.
The following theorem requires \emph{none} of these assumptions: task sizes,
arrival times, and routing targets may all be chosen adversarially.
\end{remark}

\begin{theorem}[Worst-case BoT competitive ratio]\label{thm:cr_bot_worst}
Let $\sigma = (s_1, \ldots, s_M)$ be any bag-of-tasks sequence with total
work $W = \sum_{i=1}^M s_i$, offered load $\varrho_0 < 1$
(so the system is stable by Theorem~\ref{thm:lyapunov}).
DTA's makespan satisfies
\begin{equation}\label{eq:cr_worst}
  C_{\max}^{\mathrm{DTA}}(\sigma)
  \;\le\;
  \frac{W}{N}
  \;+\;
  M \cdot \left\lfloor \frac{N}{2} \right\rfloor \cdot \delta_{\mathrm{hop}},
\end{equation}
and therefore the worst-case competitive ratio is
\begin{equation}\label{eq:cr_worst_ratio}
  \varrho_c^{\mathrm{DTA,worst}}(\sigma)
  \;\le\;
  1 + \frac{N\,\lfloor N/2 \rfloor\,\delta_{\mathrm{hop}}}{W}.
\end{equation}
In particular, $\varrho_c^{\mathrm{DTA,worst}} \to 1$ as $W \to \infty$ (large-workload limit).
\end{theorem}

\begin{proof}
We establish three independent facts.

\textbf{(i) Work conservation.}
By Theorem~\ref{thm:lyapunov}, DTA is stable under $\varrho_0 < 1$.
Combined with no-task-loss (\hyperref[chap:no_task_loss]{Correctness~I}),
DTA is \emph{work-conserving}: every task in $\sigma$ is eventually completed.
Hence $C_{\max}^{\mathrm{DTA}} \ge W/N$ (work lower bound, matching the offline optimum).

\textbf{(ii) Hop-count bound.}
By construction, each task is deflected at most $\lfloor N/2 \rfloor$ times
before being accepted into a worker queue or the warehouse
(Algorithm~\hyperref[chap:formal_model]{DTA-V3}, hop-counter invariant).
Each hop incurs latency $\delta_{\mathrm{hop}}$ (one SPSC mailbox write).
Over $M$ tasks, the total deflection overhead is at most
$M \lfloor N/2 \rfloor \delta_{\mathrm{hop}}$.
This bound is deterministic and holds for any arrival sequence $\sigma$.

\textbf{(iii) Makespan decomposition.}
Partition $C_{\max}^{\mathrm{DTA}}$ into work time and routing overhead:
\[
  C_{\max}^{\mathrm{DTA}}
  \;=\;
  \underbrace{\frac{W}{N}}_{\text{work time}}
  \;+\;
  \underbrace{\Delta_{\mathrm{route}}}_{\text{deflection overhead}}.
\]
By (ii), $\Delta_{\mathrm{route}} \le M \lfloor N/2 \rfloor \delta_{\mathrm{hop}}$.
Since $C_{\max}^{\mathrm{OPT}} \ge W/N$, dividing by $C_{\max}^{\mathrm{OPT}}$
and using $C_{\max}^{\mathrm{OPT}} \ge W/N$ gives \eqref{eq:cr_worst_ratio}.
\end{proof}

\begin{remark}[Comparison with Graham's list-scheduling bound]
Graham's (1969) bound for online list scheduling on $N$ identical machines
is $\varrho_c^{\mathrm{LIST}} \le 2 - 1/N$, a constant independent of workload.
DTA achieves $\varrho_c^{\mathrm{DTA,worst}} \le 1 + N\lfloor N/2\rfloor\delta_{\mathrm{hop}}/W$,
which converges to $1$ as $W$ grows.
The improvement stems from DTA's deflection mechanism providing implicit
load balancing even under adversarial arrivals:
the bounded hop count ensures that concentrated bursts are redistributed
within at most $\lfloor N/2 \rfloor$ ring hops, analogous to a
\physics{bounded-range scattering potential}{finite interaction radius on the
ring topology}.
For the DTA benchmark ($N=4$, $\delta_{\mathrm{hop}} \approx 10\,\mathrm{ns}$,
$\delta = 1\,\mu\mathrm{s}$, $W/M = \delta$):
\[
  \varrho_c^{\mathrm{DTA,worst}} \;\le\; 1 + \frac{4 \cdot 2 \cdot 10^{-8}}{M \cdot 10^{-6}}
  \;=\; 1 + \frac{0.08}{M},
\]
which is $< 1.001$ for any $M \ge 80$ tasks.
\end{remark}

\begin{remark}[DAG tasks are out of scope]
Theorem~\ref{thm:cr_bot_worst} applies to \emph{independent} tasks only.
For tasks with dependency edges (DAG topology), DTA routes tasks as
independent units and cannot enforce predecessor-completion ordering.
The worst-case competitive ratio for DAG tasks under DTA is
\emph{unbounded}: an adversary constructs a two-task DAG
$A \to B$ (B depends on A) such that DTA schedules $B$ before $A$ completes,
producing an incorrect result.
Work-Stealing (WS)~\cite{blumofe1999,acar2002} handles DAGs correctly,
achieving $\mathbb{E}[C_{\max}^{\mathrm{WS}}] = T_1/N + O(T_\infty)$~\cite{arora2001}.
These two schedulers have \textbf{formally incomparable} guarantees:
DTA is worst-case optimal for BoT (competitive ratio $\to 1$);
WS is worst-case optimal for DAGs (Graham-optimal);
neither dominates the other across both workload classes.
\end{remark}

\section*{Summary of Main Results}

\begin{tcolorbox}[colback=yellow!5!white,colframe=orange!60!black,title=Scope Warning]
The DTA results in Table~A below are \textbf{average-case} bounds (i.i.d.\ Poisson
arrivals, mean-field) or \textbf{worst-case} BoT bounds (Theorem~\ref{thm:cr_bot_worst}).
The WS result in Table~B is a \textbf{worst-case} bound over adversarial DAG sequences.
\textbf{These two frameworks are formally incomparable}: DTA and WS optimise
for different workload classes.  Placing them in the same table would conflate
average-case BoT with adversarial DAG worst-case.
\end{tcolorbox}

\medskip
\noindent\textbf{Table A: Bag-of-Tasks (BoT) competitive ratio bounds}

\begin{center}
\begin{tabularx}{\textwidth}{@{}clXl@{}}
\toprule
Result & Model & Bound & Reference \\
\midrule
DTA (BoT, average case) & i.i.d.\ Poisson, mean-field
  & $1 + p_d\delta_{\mathrm{hop}}\mu/(1-p_d) + O(\ln N/\bar q)$
  & Theorem~\ref{thm:cr_bot} \\
DTA (BoT, worst case) & any arrival seq., $\varrho_0 < 1$
  & $1 + N\lfloor N/2\rfloor\delta_{\mathrm{hop}}/W \to 1$
  & Theorem~\ref{thm:cr_bot_worst} \\
DTA large-$H$ limit & mean-field, $H\to\infty$
  & $\to 1$
  & Corollary~\ref{cor:large_H} \\
DTA target deployment & $N=256$, $H=114{,}688$, $\varrho_0{=}0.7$
  & $\varrho_c \lesssim 5.9$ (large-$H$ bound is loose here; use worst-case row above)
  & Corollary~\ref{cor:benchmark} \\
Heavy-tail penalty & Pareto$(\alpha)$ sizes
  & $+\,(1+\mathrm{CV}^2)/N$
  & Proposition~\ref{prop:heavy_tail} \\
WS (BoT, average case) & i.i.d.\ Poisson
  & $1 + (1-\varrho^*)\delta_{\mathrm{steal}}\mu$
  & Theorem~\ref{thm:cr_ws_bot} \\
\bottomrule
\end{tabularx}
\end{center}

\medskip
\noindent\textbf{Table B: Fork-join DAG makespan bounds (reference only; DTA does not apply)}

\begin{center}
\begin{tabularx}{\textwidth}{@{}clXl@{}}
\toprule
Scheduler & Workload & Bound & Reference \\
\midrule
Work-Stealing & Fork-join DAG, adversarial & $T_1/N + O(T_\infty)$ & \cite{blumofe1999} \\
DTA & Fork-join DAG & \textit{N/A} (no dep.\ tracking; ratio unbounded) & Section~\ref{sec:cr_bot_worst} \\
\bottomrule
\end{tabularx}
\end{center}


% ============================================================
\clearpage
\section{DAG Scheduling via Sequential-Nucleation Kinetic Theory}
\label{sec:dag_kinetic}
% ============================================================

DTA routes tasks as independent units; tasks with dependency constraints
(DAG workloads) require an additional \emph{application-layer synchronisation
protocol}.  This section derives rigorous makespan bounds for restricted DAG
topologies by extending the kinetic theory of
\hyperref[chap:kinetic_theory]{Performance~II} to multi-level systems.
Two-phase protocol is used throughout: workers schedule each DAG level as
an independent BoT job, and atomic counters enforce the join barrier between
levels.

% ------------------------------------------------------------
\subsection{Setup and Notation}
% ------------------------------------------------------------

Consider a DAG $G=(V,E)$ with $d$ levels, each level $\ell \in \{0,\ldots,d-1\}$
containing $M_\ell$ tasks.  A task $v$ at level $\ell$ may not start until
all $K_\ell$ of its predecessors at level $\ell-1$ have completed.
\emph{Notational note:} in this section $K$ (and $K_\ell$) denotes DAG
\textbf{fan-in} (a per-task predecessor count), which is unrelated to
the fixed \code{CHUNK\_SIZE} $K=32$ used for $K$ elsewhere in this paper
(Table~\ref{tab:params} and passim); the two meanings are
distinguished purely by context (fan-in $K$ always carries a subscript
$\ell$ or appears in DAG-specific expressions such as $H_K$), and DTA's
chunking mechanism plays no role in this section's analysis.

\begin{definition}[Two-phase protocol]
\label{def:two_phase}
For each level $\ell$:
\begin{enumerate}
  \item \textbf{Schedule phase.} All $M_\ell$ tasks are submitted to DTA
        as an independent BoT batch and processed via the standard
        deflection protocol.
  \item \textbf{Synchronisation phase.} Each level-$\ell$ task that
        completes atomically decrements a shared counter initialised
        to $K_{\ell+1}$.  When the counter reaches zero the single
        decrementing worker submits all level-$(\ell{+}1)$ tasks to DTA.
        No modification to the DTA scheduler is required.
\end{enumerate}
\end{definition}

\textbf{Notation.}
$f_\ell(t)$ denotes the fraction of level-$\ell$ tasks completed by time $t$,
with $f_{-1}(t) \equiv 1$ (source).
$H_K = \sum_{j=1}^K 1/j$ is the $K$-th harmonic number.
$\gamma$ is the spectral gap of the per-worker birth--death chain
(\hyperref[chap:stability]{Performance~IV}, \S{}2).
$T_1 = \sum_\ell M_\ell/\mu$ is the total work.

% ------------------------------------------------------------
\subsection{Mean-Field Kinetic Equation}
% ------------------------------------------------------------

\begin{theorem}[Mean-field DAG kinetics]\label{thm:dag_ode}
In the random DAG ensemble (each of the $K_\ell$ predecessors of a level-$\ell$
task is chosen uniformly and independently from the $M_{\ell-1}$ tasks at
level $\ell{-}1$), the empirical completion process $f_\ell^{(N)}(t)$
converges in probability, as $N \to \infty$, to the unique solution of the ODE
system
\begin{equation}\label{eq:dag_ode}
  \frac{d f_\ell}{dt}
  \;=\; \mu\,(1 - f_\ell)\,\bigl[f_{\ell-1}(t)\bigr]^{K_\ell},
  \qquad \ell = 0, \ldots, d-1,
\end{equation}
with $f_{-1}(t) \equiv 1$ and $f_\ell(0) = 0$.
The convergence rate is $O(1/\sqrt{N})$ uniformly in $t$.
\end{theorem}

\begin{proof}
Apply the propagation-of-chaos theorem~\cite{sznitman1991} to the following
$N$-particle system.
Each particle (task) at level $\ell$ carries a binary state
$\sigma_v \in \{0,1\}$ (incomplete/complete) and transitions $0 \to 1$
at rate $\mu$ once all $K_\ell$ of its designated predecessors satisfy
$\sigma_u = 1$.
Under the random-predecessor model, the overlap probability that two distinct
level-$\ell$ tasks share a predecessor is $O(K_\ell^2 / M_{\ell-1})$.
For $M_{\ell-1} \to \infty$ the predecessors of distinct tasks are
asymptotically independent (\emph{propagation of chaos}):
the joint law of any fixed set of predecessors factorises in the limit.
Conditional on $f_{\ell-1}(t) = p$, each predecessor is complete with
probability $p$, independently; hence the activation probability is $p^{K_\ell}$.
Once activated, the task completes at rate $\mu$, giving~\eqref{eq:dag_ode}.
Convergence $f_\ell^{(N)} \to f_\ell$ in probability at rate $O(1/\sqrt{N})$
follows from the McKean--Vlasov central limit theorem for exchangeable particle
systems~\cite{sznitman1991}.
\end{proof}

\subsubsection{Phase-transition character of the activation function}

The activation factor $\varphi_{K}(p) = p^K$ in~\eqref{eq:dag_ode}
exhibits qualitatively different behaviour depending on $K$:

\begin{itemize}
  \item \textbf{Finite $K$ (second-order crossover).}
        $\varphi_K$ is smooth on $[0,1]$; level $\ell$ tasks activate
        gradually as $f_{\ell-1}$ grows, producing a
        \emph{continuous crossover} between parallel and sequential regimes.
        This is a second-order (continuous) phase transition in the
        thermodynamic limit $M_\ell \to \infty$.
  \item \textbf{$K \to \infty$ (first-order transition).}
        $\varphi_K(p) \to \mathbf{1}[p = 1]$, a step function.
        Level-$\ell$ tasks are strictly inactive until $f_{\ell-1} = 1$,
        then activate discontinuously---a first-order (discontinuous)
        phase transition characterised by metastability and a latent
        synchronisation cost.
\end{itemize}

The \emph{critical fan-in} separating these regimes is
$K^* \approx 1/\varepsilon$ where $\varepsilon$ is the desired activation
precision: for $K \gg K^*$ the crossover is effectively step-like
(fork-join behaviour), whereas for $K = O(1)$ the crossover is smooth.
Figure~\ref{fig:dag_ode} illustrates the ODE dynamics for $K=1, 5, 50$.

\begin{figure}[h]
  \centering
  \includegraphics[width=\linewidth]{figure/dag_ode_dynamics.png}
  \caption{Mean-field DAG kinetics $f_\ell(t)$ for three levels
    ($\ell=0,1,2$) and three fan-in values.
    Left ($K=1$): smooth sequential cascade (2nd-order crossover).
    Centre ($K=5$): moderate sharpening.
    Right ($K=50$): near-step-function activation (1st-order regime),
    approximating the fork-join limit.
    Produced by \texttt{simulate\_dag\_scheduling.py}.}
  \label{fig:dag_ode}
\end{figure}

\begin{figure}[h]
  \centering
  \includegraphics[width=0.62\linewidth]{figure/dag_phase_transition.png}
  \caption{Activation function $\varphi_K(f_{\ell-1}) = f_{\ell-1}^K$ for
    $K = 1, 2, 5, 10, 50$.
    Increasing $K$ drives the smooth crossover toward a step function,
    marking the transition from second-order (finite $K$) to
    first-order ($K \to \infty$) phase-transition character.}
  \label{fig:dag_phase}
\end{figure}

% ------------------------------------------------------------
\subsection{Fork-Join DAG: First-Order Bounds}
% ------------------------------------------------------------

\begin{theorem}[Fork-join makespan, first order]\label{thm:fj_first}
Consider a fork-join DAG with $d$ levels, $M_\ell$ tasks at level $\ell$,
total work $T_1 = \sum_\ell M_\ell/\mu$, and $K$-way join
(all $K$ predecessors required, $K \geq 1$).
Under the two-phase protocol (Definition~\ref{def:two_phase})
with DTA scheduling at each level and offered load $\varrho_0 < 1$:
\begin{equation}\label{eq:fj_first}
  \mathbb{E}[C_{\max}^{\mathrm{FJ}}]
  \;\leq\;
  \frac{T_1}{N}
  \;+\;
  \frac{d\,H_K}{\mu}.
\end{equation}
\end{theorem}

\begin{proof}
Decompose the makespan into within-level processing time and inter-level
synchronisation time.

\smallskip\noindent
\emph{(i) Processing contribution.}
By Theorem~\ref{thm:cr_bot_worst} (worst-case BoT competitive ratio),
DTA is work-conserving under $\varrho_0 < 1$: every task is eventually
scheduled and completed.
The total processing time across all $d$ levels is bounded by the work
bound $T_1/N$ (since $N$ workers each receive at most $T_1/N$ of work in
expectation).

\smallskip\noindent
\emph{(ii) Synchronisation contribution.}
Between levels $\ell{-}1$ and $\ell$, the protocol waits for all $K$
predecessors of the next level's representative task to complete.
These $K$ predecessors have i.i.d.\ $\mathrm{Exp}(\mu)$ service times
(by exchangeability of DTA workers).
Model the join wait as a birth--death chain on states
$\{K, K-1, \ldots, 1, 0\}$, where state $m$ means $m$ predecessors
remain incomplete and transitions occur at rate $m\mu$ (any one of the $m$
remaining tasks may complete next).
By the optional stopping theorem~\cite{norris1997} applied to this
chain, the expected first-passage time from state $K$ to state $0$ is
\[
  \mathbb{E}[\tau_{\mathrm{join}}^{(K)}]
  \;=\; \sum_{j=1}^K \frac{1}{j\mu}
  \;=\; \frac{H_K}{\mu}.
\]
The two-phase protocol enforces that the $d$ sync events occur
sequentially and are conditionally independent given the level-completion
times; summing gives total sync cost $d H_K/\mu$.

\smallskip\noindent
Adding (i) and (ii) establishes~\eqref{eq:fj_first}.
\end{proof}

% ------------------------------------------------------------
\subsection{Second-Order Corrections}
% ------------------------------------------------------------

\begin{proposition}[Second-order correction $\delta^{(A)}$: Gumbel fluctuation]
\label{prop:dag_2A}
Under the conditions of Theorem~\ref{thm:fj_first}, the variance of
the total synchronisation cost is
\begin{equation}\label{eq:dag_varA}
  \mathrm{Var}\!\Bigl[\sum_{\ell=1}^d \tau_{\mathrm{join},\ell}^{(K)}\Bigr]
  \;=\;
  \frac{d}{\mu^2} \sum_{j=1}^K \frac{1}{j^2}
  \;\leq\;
  \frac{d\,\pi^2}{6\mu^2}.
\end{equation}
Consequently, the $(1-\alpha)$-confidence upper bound on the makespan is
\[
  C_{\max}^{(1-\alpha)}
  \;\leq\;
  \frac{T_1}{N} + \frac{d H_K}{\mu}
  + \frac{z_{1-\alpha}\,\pi\sqrt{d}}{\sqrt{6}\,\mu},
\]
where $z_{1-\alpha}$ is the standard normal quantile.
\end{proposition}

\begin{proof}
The join time $\tau_{\mathrm{join}}^{(K)} = \max(T_1,\ldots,T_K)$ for
i.i.d.\ $T_i \sim \mathrm{Exp}(\mu)$.
Using the same birth--death representation as in Theorem~\ref{thm:fj_first},
write $\tau_{\mathrm{join}}^{(K)} = X_K + X_{K-1} + \cdots + X_1$ where
$X_j \sim \mathrm{Exp}(j\mu)$ are independent holding times.
Then
\[
  \mathrm{Var}[\tau_{\mathrm{join}}^{(K)}]
  \;=\; \sum_{j=1}^K \mathrm{Var}[X_j]
  \;=\; \sum_{j=1}^K \frac{1}{j^2\mu^2}.
\]
The upper bound $\sum_{j=1}^\infty j^{-2} = \pi^2/6$ (Basel problem) gives
the right-hand side of~\eqref{eq:dag_varA}.
The $d$ sync events are independent by the two-phase protocol;
variances add.
The confidence bound follows from the central limit theorem
(applicable for $d \geq 2$) with the Gumbel variance $\pi^2/(6\mu^2)$.
\end{proof}

\begin{figure}[h]
  \centering
  \begin{minipage}{0.48\linewidth}
    \centering
    \includegraphics[width=\linewidth]{figure/dag_sync_cost_vs_K.png}
    \caption{Expected join time $\mathbb{E}[\tau_{\mathrm{join}}^{(K)}]$
      vs fan-in $K$.  The harmonic bound $H_K/\mu$ (solid) matches
      simulation (dots) to within statistical error.}
    \label{fig:dag_sync_cost}
  \end{minipage}
  \hfill
  \begin{minipage}{0.48\linewidth}
    \centering
    \includegraphics[width=\linewidth]{figure/dag_variance_vs_K.png}
    \caption{Standard deviation of join time vs $K$.
      The Gumbel asymptote $\pi/(\sqrt{6}\,\mu)$
      (dashed red) is reached already by $K \approx 20$.
      This is the second-order correction $\delta^{(A)}$.}
    \label{fig:dag_var}
  \end{minipage}
\end{figure}

\begin{proposition}[Second-order correction $\delta^{(B)}$: incomplete relaxation]
\label{prop:dag_2B}
Under the two-phase protocol, each DAG level starts with a batch-arrival
initial condition (all $M_\ell$ tasks submitted simultaneously at time
$\tau_\ell$) rather than the NESS initial condition assumed in the
BoT makespan bound.
The deviation of the per-level expected makespan from the NESS value is
bounded by
\begin{equation}\label{eq:dag_relax}
  \bigl|\mathbb{E}[C_\ell^{\mathrm{batch}}]
  - \mathbb{E}[C_\ell^{\mathrm{NESS}}]\bigr|
  \;\leq\;
  \frac{2}{\gamma},
\end{equation}
and summing over $d$ levels gives total correction
$\delta C_{\max}^{(B)} \leq 2d/\gamma$.
\end{proposition}

\begin{proof}
Let $\mu^{\mathrm{batch}}$ and $\mu^{\mathrm{NESS}}$ denote the
distribution of the system state at the start of level $\ell$ under
batch arrival and NESS respectively.
By Fill's spectral bound~\cite{fill1991}, for any initial distribution
$\nu$ and any $t \geq 0$:
\[
  \|\nu P^t - \pi\|_{\mathrm{TV}} \;\leq\; e^{-\gamma t},
\]
where $\pi$ is the stationary distribution and $P^t$ is the semigroup
of the per-worker birth--death chain.
Therefore, for the two initial distributions $\mu^{\mathrm{batch}}$
and $\mu^{\mathrm{NESS}}$:
\[
  \|\mu^{\mathrm{batch}} P^t - \mu^{\mathrm{NESS}} P^t\|_{\mathrm{TV}}
  \;\leq\;
  \|\mu^{\mathrm{batch}} P^t - \pi\|_{\mathrm{TV}}
  + \|\pi - \mu^{\mathrm{NESS}} P^t\|_{\mathrm{TV}}
  \;\leq\; 2e^{-\gamma t}.
\]
For any measurable event $A$ (in particular $A = \{C_\ell > t\}$):
\[
  \bigl|
    \mathbb{P}^{\mathrm{batch}}(C_\ell > t)
    - \mathbb{P}^{\mathrm{NESS}}(C_\ell > t)
  \bigr|
  \;\leq\;
  \|\mu^{\mathrm{batch}} P^t - \mu^{\mathrm{NESS}} P^t\|_{\mathrm{TV}}
  \;\leq\; 2e^{-\gamma t}.
\]
Integrating over $t \geq 0$ and using
$\mathbb{E}[C] = \int_0^\infty \mathbb{P}(C > t)\,dt$:
\[
  \bigl|\mathbb{E}[C_\ell^{\mathrm{batch}}]
  - \mathbb{E}[C_\ell^{\mathrm{NESS}}]\bigr|
  \;\leq\;
  \int_0^\infty 2e^{-\gamma t}\,dt
  \;=\; \frac{2}{\gamma}.
\]
Summing the bound independently over $d$ levels (each level is a
fresh application of the above argument) gives $\delta C_{\max}^{(B)} \leq 2d/\gamma$.
\end{proof}

\begin{remark}[Combined bound and high-load asymptotics]
\label{rem:dag_combined}
Combining Theorem~\ref{thm:fj_first} and
Propositions~\ref{prop:dag_2A}--\ref{prop:dag_2B},
the full makespan bound is
\begin{equation}\label{eq:fj_full}
  \mathbb{E}[C_{\max}^{\mathrm{FJ}}]
  \;\leq\;
  \underbrace{\frac{T_1}{N} + \frac{d\,H_K}{\mu}}_{\text{first order}}
  \;+\;
  \underbrace{\frac{\pi\sqrt{d}}{\sqrt{6}\,\mu}}_{\delta^{(A)}}
  \;+\;
  \underbrace{\frac{2d}{\gamma}}_{\delta^{(B)}}.
\end{equation}
As $\varrho_0 \to 1^-$, the spectral gap $\gamma \to 0$
(\hyperref[chap:stability]{Performance~IV}, Theorem~1.2),
so $\delta^{(B)} \to \infty$: the first-order approximation
becomes unreliable near the stability boundary.
In the light-to-moderate load regime ($\varrho_0 \leq 0.8$),
numerical evaluation of the DTA benchmark gives
$\gamma \approx 0.04\,\mu$ (\hyperref[chap:stability]{Performance~IV},
Figure~\ref{fig:stab_L2}), so $\delta^{(B)} \approx 50d/\mu$, which is
small relative to $T_1/N$ for sufficiently large batches.
\end{remark}

\begin{figure}[h]
  \centering
  \includegraphics[width=0.65\linewidth]{figure/dag_makespan_vs_depth.png}
  \caption{Fork-join makespan $\mathbb{E}[C_{\max}]$ vs DAG depth $d$
    ($K=4$, $N=4$, $M_\ell=16$ tasks per level).
    First-order bound $T_1/N + dH_K/\mu$ (solid), combined bound
    including $\delta^{(A)}$ (dashed), and Monte Carlo simulation
    with $95\%$ confidence intervals.
    Produced by \texttt{simulate\_dag\_scheduling.py}.}
  \label{fig:dag_makespan}
\end{figure}

% ------------------------------------------------------------
\subsection{Special Cases and Finite Fan-In}
% ------------------------------------------------------------

\begin{theorem}[Finite fan-in, mean-field correction]\label{thm:finite_fanin}
For the ODE system~\eqref{eq:dag_ode} with uniform fan-in $K_\ell = K$,
the completion time of level $\ell$ satisfies
$\tau_\ell = \inf\{t : f_\ell(t) = 1\}$.
The Level-1 BBGKY pair-correlation correction shifts $f_\ell(t)$
by at most $K^2/M_\ell$ uniformly in $t$, yielding a finite-$N$ correction
to $\tau_\ell$ of order $O(K^2/(M_\ell\,\mu))$ beyond the mean-field value.
\end{theorem}

\begin{proof}
The Level-1 BBGKY equation introduces the pair correlator
$g_2^{(\ell)}(v,v')$ for two tasks $v, v'$ at level $\ell$
(see \hyperref[chap:kinetic_theory]{Performance~II}, Section~3).
In the random-predecessor model, two level-$\ell$ tasks share a predecessor
with probability $K^2/M_{\ell-1}^2 \cdot M_{\ell-1} = K^2/M_{\ell-1}$.
The pair correlator is thus $O(K^2/M_{\ell-1})$, and its contribution
to $df_\ell/dt$ is bounded by $K^2/M_{\ell-1}$ times $O(\mu)$,
giving a correction to $\tau_\ell$ of $O(K^2/(M_{\ell-1}\mu))$.
\end{proof}

The following table summarises the first-order makespan bounds for
the DAG types studied.

\begin{center}
\small
\begin{tabular}{@{}llll@{}}
\toprule
DAG type & Fan-in $K$ & Depth $d$ & $\mathbb{E}[C_{\max}]$ first-order bound \\
\midrule
Fork-join & $K$ (fixed) & $d$ & $T_1/N + d H_K/\mu$ \\
Binary tree & $2$ & $\lceil\log_2 M\rceil$ & $T_1/N + \tfrac{3}{2\mu}\lceil\log_2 M\rceil$ \\
Series-parallel & $K_{\max}$ & $d_{\max}$ & $T_1/N + d_{\max} H_{K_{\max}}/\mu$ \\
Chain (serial) & $1$ & $M$ & $M/\mu$ \quad (no parallelism) \\
BoT ($d=1$) & $0$ & $1$ & $T_1/N + O(\ln N/\bar{q})$ \\
\bottomrule
\end{tabular}
\end{center}

\noindent
Second-order corrections $\delta^{(A)} = \pi\sqrt{d}/(\sqrt{6}\mu)$ and
$\delta^{(B)} = 2d/\gamma$ apply uniformly to all cases with $d \geq 2$.
For the binary tree specialisation ($K=2$, $d = \lceil\log_2 M\rceil$):
$H_2 = 3/2$, the sync cost per level is $3/(2\mu)$, and the total
synchronisation overhead grows only as $O(\log M)$---logarithmic in the
number of tasks.

% ------------------------------------------------------------
\subsection{Connection to Directed Percolation: Open Problems}
% ------------------------------------------------------------

The mean-field kinetic equation~\eqref{eq:dag_ode} describes scheduling
on the \emph{random DAG ensemble}; for a general (non-random) DAG,
the appropriate framework is the theory of
\emph{directed percolation} (DP)~\cite{hinrichsen2000}.

Consider a random DAG where each directed edge from level $\ell{-}1$
to level $\ell$ exists independently with probability $p$.
As $p$ decreases from $1$ (fully connected) to $0$ (empty), the
typical critical-path length $T_\infty$ undergoes a phase transition
at a threshold $p_c$:
\begin{itemize}
  \item For $p > p_c$ (dense): $T_\infty = O(d)$; scheduling is
        critical-path-bound.
  \item For $p < p_c$ (sparse): $T_\infty = O(\log d)$; scheduling is
        work-bound ($\approx T_1/N$).
  \item At $p = p_c$: universal scaling with DP critical exponents.
\end{itemize}
The scheduling order parameter
$\Phi = (C_{\max} - T_1/N)/(T_\infty - T_1/N)$
(fraction of makespan attributable to the critical path)
scales near $p_c$ as~\cite{hinrichsen2000}
\[
  \Phi \;\sim\; |p - p_c|^{\beta},
  \qquad
  \beta \approx 0.2765 \quad (d=1\text{ space} + 1\text{ time}).
\]
The temporal and spatial correlation lengths diverge as
$\xi_\parallel \sim |p-p_c|^{-\nu_\parallel}$ with $\nu_\parallel \approx 1.734$
and $\xi_\perp \sim |p-p_c|^{-\nu_\perp}$ with $\nu_\perp \approx 1.097$.
The dynamic exponent is $z = \nu_\parallel/\nu_\perp \approx 1.581$.

The Level-0 BBGKY closure~\eqref{eq:dag_ode} corresponds to
\emph{mean-field DP} (valid for $d \geq d_c = 4$ spatial dimensions).
In the physically relevant case $d = 1+1$ (one spatial dimension = the
worker ring, one temporal dimension = DAG depth), the mean-field exponents
are corrected by fluctuations; their values are known numerically to high
precision~\cite{hinrichsen2000} but a rigorous analytical derivation for
the scheduling model remains an open problem.

\begin{openq}[DP universality for DAG scheduling]
Establish rigorously that, in the random DAG ensemble with edge probability
$p$ near $p_c$, the scheduling phase transition is in the DP universality
class and the makespan scaling exponents equal the DP critical exponents
$\beta, \nu_\parallel, \nu_\perp$.
\end{openq}

\begin{openq}[BBGKY Level-2 for DAG kinetics]
Derive the Level-2 BBGKY correction to the ODE system~\eqref{eq:dag_ode},
analogous to the Level-1 correction in
\hyperref[chap:kinetic_theory]{Performance~II} (Section~3.2).
The Level-2 equation introduces triplet correlators
$g_3^{(\ell)}(v,v',v'')$ and is expected to yield the first
fluctuation correction to the DP mean-field exponents.
\end{openq}

\clearpage

% ========================================================================
% Adversarial Stability Comparison
% ========================================================================

\section*{Adversarial Stability Comparison}
\addcontentsline{toc}{section}{Adversarial Stability Comparison}
\label{chap:stability_compare}
\setcounter{section}{0}

\textbf{Notation.}
$N$ workers, load $\varrho_0 = \lambda_0/\mu$, self-consistent load $\varrho^*$,
per-worker capacity $H$, mean queue $\bar q$, warehouse capacity $C_W$
(DTA only).
Four strategies compared: \textbf{DTA} (hop-bounded deflection, this work),
\textbf{CQ} (Central Queue, M/M/N model), \textbf{Rand} (random assignment,
no rebalancing), \textbf{WS} (Work-Stealing with Chase--Lev deque).
spectral gap, burst capacity, first-passage crash time).

% ===========================================================================
\section{Strategies Under Comparison}
% ===========================================================================

Each strategy is a complete scheduling system, not just an assignment rule.
We characterise each by its steady-state queue model.

\begin{definition}[Strategy queue models]~
\begin{enumerate}
  \item \textbf{DTA:}
        $N$ workers, each an M/M/1/$H$ queue under mean-field deflection.
        Effective per-worker load $\varrho^* = \varrho_0(1 + p_d)$
        (self-consistent equation, \hyperref[chap:formal_model]{the Formal Model}).
        Warehouse buffer $C_W$ for overflow.  Hop limit $\lfloor N/2 \rfloor$.

  \item \textbf{Central Queue (CQ):}
        Single global FIFO, $N$ servers.  The classical M/M/$N$ model with
        Erlang-C steady state.  Tasks see mean wait
        $W_{\mathrm{CQ}} = C(N,\varrho_0) / (N\mu(1-\varrho_0/N))$,
        where $C(N,\varrho_0)$ is the Erlang-C formula.
        No per-worker queue capacity; no overflow buffer needed.
        \emph{Requires a single point of enqueue coordination.}

  \item \textbf{Random (Rand):}
        Each arriving task is assigned uniformly at random to one of $N$ workers,
        independently.  Per-worker load $\varrho_0$ (no rebalancing).
        Each worker is an independent M/M/1 queue, stable iff $\varrho_0 < 1$.

  \item \textbf{Work-Stealing (WS):}
        Tasks pushed to local deque; idle workers steal.  Equivalent to
        M/M/$N$ in steady-state throughput for BoT, but with different
        transient dynamics and per-operation overhead.  No finite queue capacity
        ($H = \infty$ effectively); no warehouse buffer.
        \emph{Requires atomic CAS on shared deque pointers.}
\end{enumerate}
\end{definition}

% ===========================================================================
\section{Steady-State Queue Comparison}
% ===========================================================================

\begin{proposition}[Mean queue comparison]\label{prop:mean_queue}
At offered load $\varrho_0 < 1$ with $N$ workers:
\begin{align}
  \bar q_{\mathrm{DTA}} &= \frac{\varrho^*}{1-\varrho^*}
     \cdot \frac{1 - (\varrho^*)^H (H(1-\varrho^*)+1)}{1 - (\varrho^*)^{H+1}}
     \;\le\; \frac{\varrho^*}{1-\varrho^*}, \label{eq:q_dta}\\
  \bar q_{\mathrm{CQ}} &= \frac{\varrho_0/N}{1 - \varrho_0/N}
     \;+\; \frac{C(N,\varrho_0)\,\varrho_0/N}{N(1-\varrho_0/N)^2 \cdot N!}
     \;\approx\; \frac{\varrho_0}{N - \varrho_0}\quad(N \text{ large}), \label{eq:q_cq}\\
  \bar q_{\mathrm{Rand}} &= \frac{\varrho_0}{1 - \varrho_0}, \label{eq:q_rand}\\
  \bar q_{\mathrm{WS}} &= \frac{\varrho_0/N}{1-\varrho_0/N}
     \;\approx\; \frac{\varrho_0}{N - \varrho_0}
     \quad(\text{BoT, steady state}). \label{eq:q_ws}
\end{align}
Hence $\bar q_{\mathrm{CQ}} \approx \bar q_{\mathrm{WS}} \le \bar q_{\mathrm{DTA}} \le \bar q_{\mathrm{Rand}}$,
with DTA converging to CQ/WS as $\varrho_0 \to 1$.
\end{proposition}

\begin{proof}
CQ and WS achieve global load balancing, giving effective per-worker load
$\varrho_0/N \ll \varrho^*$.  DTA's mean-field self-consistent load
$\varrho^* > \varrho_0/N$ but $< \varrho_0$ (deflection partially balances).
Rand has no rebalancing; per-worker load = $\varrho_0$.  The ordering follows
from the M/M/1 formula $\bar q = \varrho/(1-\varrho)$ applied to each regime.
\end{proof}

\begin{remark}[Physical reading]
The mean queue ordering reflects
\physics{equilibration efficiency}{how well each strategy spreads load}.
CQ and WS achieve \physics{global thermodynamic equilibrium}
{perfect load balance across all workers}.
DTA achieves \physics{local equilibrium}{mean-field: each worker
statistically equivalent} with an $O(1/N)$ correction from pair correlations
(see \hyperref[chap:kinetic_theory]{Performance~II}, Level~1).
Rand is a \physics{non-interacting system}{no load balancing at all};
individual queues are i.i.d.\ and the system cannot exploit diversity.
\end{remark}

% ===========================================================================

\begin{figure}[htbp]
\centering
\includegraphics[width=0.85\linewidth]{figure/burst_capacity_compare.png}
\caption{Burst capacity $B^*$ of DTA (base and with warehouse), CQ, Random, and Work-Stealing strategies versus offered load $\varrho_0$. DTA with warehouse uniquely exceeds CQ.}
\label{fig:burst_compare}
\end{figure}
\section{Burst Capacity $B^*$}
% ===========================================================================

\begin{definition}[Burst capacity]\label{def:burst}
For a strategy $\mathcal{S}$, the \emph{burst capacity} $B^*(\mathcal{S})$
is the maximum number of additional tasks that can be injected above the
steady-state arrival rate before the system saturates (any worker's queue
hits capacity, or the system becomes unstable).
\end{definition}

\begin{proposition}[Burst capacities]\label{prop:burst_compare}
At steady-state load $\varrho_0$ with per-worker capacity $L$ (where applicable):
\begin{align}
  B^*_{\mathrm{DTA}} &= N(L - \bar q_{\mathrm{DTA}}) + C_W \cdot K
     \;=\; N(L - \varrho^*/(1-\varrho^*)) + C_W \cdot K, \label{eq:B_dta}\\
  B^*_{\mathrm{CQ}} &= N \cdot L - \varrho_0/\mu \quad (\text{uniform buffer of } L \text{ per server}),
     \label{eq:B_cq}\\
  B^*_{\mathrm{Rand}} &= N(L - \varrho_0/(1-\varrho_0)), \label{eq:B_rand}\\
  B^*_{\mathrm{WS}} &\approx B^*_{\mathrm{CQ}} \quad (\text{no per-worker cap;  steals
     prevent local saturation}). \label{eq:B_ws}
\end{align}
\end{proposition}

\begin{proof}
DTA: Each worker has $L - \bar q$ tasks of remaining capacity in steady state;
warehouse adds $C_W \cdot K$ overflow slots.  Sum over $N$ workers (Theorem~3.1
of \hyperref[chap:stability]{Performance~IV}).  CQ: global buffer of $NL$ minus expected occupancy
$\varrho_0/\mu$ (Little's law).  Rand: per-worker M/M/1 queue with mean $\varrho_0/(1-\varrho_0)$;
same structure as DTA without deflection (no warehouse).  WS: effectively
infinite deque; burst capacity = total RAM capacity minus current load.
\end{proof}

\begin{corollary}[Comparison at fixed $\varrho_0$]\label{cor:burst_compare}
Define $B^*_{\mathrm{DTA,base}} = N(L - \bar q_{\mathrm{DTA}})$ (DTA without
the warehouse contribution).  Since $\varrho^*/(1-\varrho^*) \le \varrho_0/(1-\varrho_0)$,
we have $\bar q_{\mathrm{DTA}} \approx \bar q_{\mathrm{Rand}}$ (deflection has
exponentially small effect for large $H$), giving the \emph{base ordering}:
\begin{equation}\label{eq:burst_order_base}
  B^*_{\mathrm{Rand}} \;\lesssim\; B^*_{\mathrm{DTA,base}} \;\le\; B^*_{\mathrm{CQ}} \approx B^*_{\mathrm{WS}}.
\end{equation}
However, DTA's \emph{total} burst capacity includes the warehouse bonus:
\[
  B^*_{\mathrm{DTA,total}} = B^*_{\mathrm{DTA,base}} + C_W \cdot K
  \;\ge\; B^*_{\mathrm{CQ}} \quad \text{whenever } C_W K \ge B^*_{\mathrm{CQ}} - B^*_{\mathrm{DTA,base}}.
\]
The warehouse is a unique architectural advantage of DTA unavailable in CQ or WS.
Numerically (Table~1 of \texttt{simulate\_competitive.py}): at $\varrho_0 = 0.5$, $N=16$,
$B^*_{\mathrm{DTA,total}} = 2244 > B^*_{\mathrm{CQ}} = 1760$.
\end{corollary}

\begin{remark}[Physical analogy: network conductance]
The ordering in Equation~\eqref{eq:burst_order_base} mirrors the
\physics{network conductance hierarchy}{ability to redistribute load under stress}.
CQ/WS = fully connected network (any task can reach any server instantly);
DTA = ring network with bounded hop radius (short-range conductance);
Rand = disconnected islands (no conductance between workers).

The hop bound $\lfloor N/2 \rfloor$ in DTA is analogous to the
\physics{Cheeger constant of a ring graph}{connectivity bottleneck}~\cite{cheeger1970}.
For a ring of $N$ nodes with diameter $\lfloor N/2 \rfloor$, the Cheeger
constant scales as $\Phi \sim 2/N$, which by the Cheeger inequality
$\gamma \ge \Phi^2/2$ gives spectral gap $\gamma \ge 2/N^2$.
CQ has $\Phi = 1$ (complete graph); Rand has $\Phi = 0$ (no edges).
\end{remark}

% ===========================================================================

\begin{figure}[htbp]
\centering
\includegraphics[width=0.85\linewidth]{figure/spectral_gap_compare.png}
\caption{Spectral gap $\gamma$ comparison across strategies. DTA exhibits slower mixing ($\gamma_{\mathrm{DTA}} < \gamma_{\mathrm{CQ}}$) but compensates via the warehouse burst buffer.}
\label{fig:spec_compare}
\end{figure}
\section{Spectral Gap and Mixing Time}
% ===========================================================================

\begin{definition}[Per-strategy spectral gap]
Let $\gamma(\mathcal{S})$ denote the spectral gap of the Markov chain
describing the queue-length process under strategy $\mathcal{S}$,
and $t_{\mathrm{mix}}(\mathcal{S}) = \log(1/\varepsilon)/\gamma(\mathcal{S})$
the $\varepsilon$-mixing time.
\end{definition}

\begin{theorem}[Mixing time comparison]\label{thm:mixing_compare}
At load $\varrho_0 < 1$ with $\mu = 1$:
\begin{align}
  \gamma_{\mathrm{DTA}} &\approx \mu(1 - \sqrt{\varrho^*})^2
     \quad(H \gg 1), \label{eq:gap_dta}\\
  \gamma_{\mathrm{CQ}} &= \mu(1 - \varrho_0/N)
     \quad(\text{M/M/}N \text{ birth-death, leading eigenvalue}), \label{eq:gap_cq}\\
  \gamma_{\mathrm{Rand}} &= \mu(1 - \varrho_0)
     \quad(\text{M/M/1, standard result}), \label{eq:gap_rand}\\
  \gamma_{\mathrm{WS}} &\approx \gamma_{\mathrm{CQ}}
     \quad(\text{BoT: equivalent to M/M/}N \text{ in steady state}). \label{eq:gap_ws}
\end{align}
Consequently, the mixing-time ordering is:
\[
  t_{\mathrm{mix}}^{\mathrm{CQ}} \;\approx\; t_{\mathrm{mix}}^{\mathrm{WS}}
  \;\le\; t_{\mathrm{mix}}^{\mathrm{DTA}}
  \;\le\; t_{\mathrm{mix}}^{\mathrm{Rand}},
\]
with equality approached as $N \to \infty$ (DTA) and $\varrho_0 \to 1$ (all strategies diverge).
\end{theorem}

\begin{proof}
Equation~\eqref{eq:gap_dta}: from \hyperref[chap:stability]{Performance~IV}, Theorem~2.1 (spectral
gap of M/M/1/$H$ mean-field chain, large-$H$ approximation).
Equation~\eqref{eq:gap_cq}: the M/M/$N$ generator on total occupancy $Q$
has birth rate $\lambda_0$ and death rate $\min(Q,N)\mu$; the spectral gap
of this birth-death chain is $\mu - \lambda_0/N = \mu(1-\varrho_0/N)$
(see \cite{fill1991}).
Equation~\eqref{eq:gap_rand}: each M/M/1 chain has gap $\mu(1-\varrho_0)$.
The mixing time for the $N$-dimensional product chain is the max of individual gaps.
The ordering follows from $\varrho^* > \varrho_0/N$ (DTA per-worker load higher
than CQ), so $\sqrt{\varrho^*} > \sqrt{\varrho_0/N}$ and $\gamma_{\mathrm{DTA}} < \gamma_{\mathrm{CQ}}$.
\end{proof}

\begin{corollary}[Critical slowing-down near saturation]
All strategies exhibit critical slowing-down as $\varrho_0 \to 1$:
\[
  t_{\mathrm{mix}}^{\mathcal{S}} \;\sim\; \begin{cases}
    (1 - \varrho_0/N)^{-1} & \mathcal{S} \in \{\mathrm{CQ, WS}\} \\
    (1 - \sqrt{\varrho^*})^{-2} & \mathcal{S} = \mathrm{DTA} \\
    (1 - \varrho_0)^{-1} & \mathcal{S} = \mathrm{Rand}
  \end{cases}
  \quad \text{as } \varrho_0 \to 1.
\]
The exponent $-2$ for DTA (vs. $-1$ for CQ) is the
\physics{critical slowing-down exponent}{mixing time power law near saturation}.
This is the signature of the NESS (non-equilibrium steady state) nature
of DTA: deflection violates detailed balance, introducing an extra
factor of $(1-\sqrt{\varrho^*})^{-1}$ relative to CQ.
\end{corollary}

% ===========================================================================

\begin{figure}[htbp]
\centering
\includegraphics[width=0.85\linewidth]{figure/crash_time_compare.png}
\caption{First-passage crash time $T_{\mathrm{crash}}$ versus overload factor $\varrho_0$ for DTA and CQ. DTA maintains a finite crash time only when the warehouse is sufficiently large.}
\label{fig:crash_compare}
\end{figure}
\section{Adversarial Stability: First-Passage Crash Time}
% ===========================================================================

\begin{definition}[Crash time under adversarial injection]
An adversary injects tasks at rate $\lambda_0 + \epsilon$ for $\epsilon > 0$
starting from steady state.  The \emph{crash time} $T_{\mathrm{crash}}$ is
the first time any worker's queue (or warehouse) saturates.
\end{definition}

\begin{theorem}[Crash time comparison]\label{thm:crash_compare}
Under adversarial injection at rate $\lambda_0 + \epsilon$:
\begin{align}
  T_{\mathrm{crash}}^{\mathrm{DTA}} &= \frac{C_W}{\Lambda_W(\varrho_0+\epsilon/\mu) - \Lambda_D(\varrho_0)}
     \quad (\Lambda_W > \Lambda_D, \text{ Theorem~4.3 of \hyperref[chap:stability]{Performance~IV}}),
     \label{eq:T_dta}\\
  T_{\mathrm{crash}}^{\mathrm{CQ}} &= \frac{N \cdot L - \varrho_0/\mu}{\epsilon/\mu}
     \quad (\text{linear drain of global buffer}), \label{eq:T_cq}\\
  T_{\mathrm{crash}}^{\mathrm{Rand}} &= \frac{L - \varrho_0/(1-\varrho_0)}{\epsilon/\mu}
     \cdot \frac{1}{N} \quad (\text{each worker absorbs } 1/N \text{ of excess}),
     \label{eq:T_rand}\\
  T_{\mathrm{crash}}^{\mathrm{WS}} &\approx T_{\mathrm{crash}}^{\mathrm{CQ}}
     \quad (\text{deque unbounded; crash occurs at memory limit}).
     \label{eq:T_ws}
\end{align}
Hence the ordering $T_{\mathrm{crash}}^{\mathrm{Rand}} \le T_{\mathrm{crash}}^{\mathrm{DTA}}
\le T_{\mathrm{crash}}^{\mathrm{CQ}} \approx T_{\mathrm{crash}}^{\mathrm{WS}}$.
\end{theorem}

\begin{proof}
DTA: from \hyperref[chap:stability]{Performance~IV}, Theorem~4.3.  The warehouse fills at rate
$\Lambda_W - \Lambda_D > 0$ when $\varrho_0 + \epsilon/\mu > \varrho_0^\dagger$
(L4 threshold); crash time $= C_W / (\Lambda_W - \Lambda_D)$.
CQ: the global buffer of $NL$ tasks absorbs the excess $\epsilon$ rate;
time to fill = remaining capacity / excess rate.
Rand: each worker independently absorbs $\epsilon/N$ of excess;
crash time is determined by the \emph{first} worker to saturate, which is
$N$ times slower than CQ (no cross-worker load transfer).
WS: equivalent to CQ for BoT since steals prevent local saturation.
\end{proof}

\begin{remark}[Architectural trade-off]
CQ achieves longer crash time $T_{\mathrm{crash}}^{\mathrm{CQ}} = N \cdot T_{\mathrm{crash}}^{\mathrm{Rand}}$
but at the cost of a \textbf{single point of failure}: the central queue.
DTA sacrifices a factor of $\approx \Lambda_W/\epsilon$ in crash time
but gains \textbf{fault tolerance} (no single queue, warehouse as safety valve)
and \textbf{lower coordination overhead} (SPSC vs.\ global lock).
WS achieves CQ-level crash time but requires \textbf{shared memory}
(CAS operations), making it unsuitable for distributed-memory environments.
\end{remark}

% ===========================================================================
\section{Cheeger Conductance: Physical Synthesis}
\label{sec:cheeger}
% ===========================================================================

The spectral gap $\gamma$ of a reversible Markov chain is bounded below by
the \emph{Cheeger constant} $\Phi$ via the Cheeger inequality:
\begin{equation}\label{eq:cheeger}
  \frac{\Phi^2}{2} \;\le\; \gamma \;\le\; 2\Phi.
\end{equation}

In the context of load-balancing strategies, $\Phi$ measures the
\physics{conductance of the load-redistribution graph}{ability to move tasks between workers}:

\begin{proposition}[Conductance ordering]\label{prop:conductance}
Let $G_{\mathcal{S}}$ be the effective task-routing graph for strategy $\mathcal{S}$.
The Cheeger constants satisfy:
\[
  \Phi(\mathrm{Rand}) = 0 \;\le\; \Phi(\mathrm{DTA}) \;\le\; \Phi(\mathrm{CQ}) = \Phi(\mathrm{WS}) = 1.
\]
For DTA on a ring of $N$ workers with hop bound $h = \lfloor N/2 \rfloor$:
\begin{equation}\label{eq:cheeger_dta}
  \Phi(\mathrm{DTA}) \;=\; \frac{\min\text{-cut across ring}}{\text{volume of smaller partition}}
  \;\approx\; \frac{2}{N} \cdot p_d(1-p_d),
\end{equation}
where the factor $p_d(1-p_d)$ reflects that only deflected tasks contribute
to cross-worker flow.
\end{proposition}

\begin{proof}[Proof sketch]
Rand has no edges ($G_{\mathrm{Rand}}$ = isolated nodes); $\Phi = 0$.
CQ and WS have the complete bipartite structure (any task reachable from any worker);
$\Phi = 1$ (connected, min-cut = 1 for any balanced partition).
DTA: the routing graph is a ring; balanced cut at the ring's diameter removes
2 edges.  The volume of each half is $N/2$ workers.  Flow through the cut
= rate of deflected tasks crossing the cut = $\lambda_0 p_d \cdot$ (fraction of ring
covered by hop bound).  For $h = N/2$, each task can reach the diametrically
opposite worker, giving $\Phi \approx 2p_d/N$ up to constants.
\end{proof}

\begin{remark}
Equation~\eqref{eq:cheeger_dta} explains why DTA's spectral gap
$\gamma_{\mathrm{DTA}} \approx \mu(1-\sqrt{\varrho^*})^2$ is smaller than
$\gamma_{\mathrm{CQ}} = \mu(1-\varrho_0/N)$: the ring topology's bounded conductance
limits the speed of global load equilibration.  However, this \emph{locality}
also prevents the \physics{Braess paradox}{pathological degradation when adding
capacity in interconnected networks}: in fully-connected WS/CQ systems,
adding workers can paradoxically increase contention for the shared queue.
DTA's hop-bounded topology is immune to this failure mode.
\end{remark}

% ===========================================================================
\section{Summary Table}
% ===========================================================================

\begin{center}
\renewcommand{\arraystretch}{1.4}
\begin{tabularx}{\textwidth}{@{}lXXXX@{}}
\toprule
Metric & DTA & Central Queue & Random & Work-Stealing \\
\midrule
Mean queue $\bar q$
  & $\varrho^*/(1-\varrho^*)$
  & $\varrho_0/(N-\varrho_0)$
  & $\varrho_0/(1-\varrho_0)$
  & $\approx \bar q_{\mathrm{CQ}}$ \\
Burst capacity $B^*$
  & $N(L-\bar q)+C_W K$
  & $NL - \varrho_0/\mu$
  & $N(L-\bar q_{\mathrm{Rand}})$
  & $\approx B^*_{\mathrm{CQ}}$ \\
Spectral gap $\gamma$
  & $\mu(1-\sqrt{\varrho^*})^2$
  & $\mu(1-\varrho_0/N)$
  & $\mu(1-\varrho_0)$
  & $\approx \gamma_{\mathrm{CQ}}$ \\
Mixing time $t_{\mathrm{mix}}$
  & $\sim(1-\varrho^*)^{-2}$
  & $\sim(1-\varrho_0/N)^{-1}$
  & $\sim(1-\varrho_0)^{-1}$
  & $\approx t_{\mathrm{mix}}^{\mathrm{CQ}}$ \\
Crash time $T_{\mathrm{crash}}$
  & $C_W/(\Lambda_W-\Lambda_D)$
  & $N(L-\bar q_{\mathrm{CQ}})/\epsilon$
  & $(L-\bar q_{\mathrm{Rand}})/(N\epsilon)$
  & $\approx T_{\mathrm{crash}}^{\mathrm{CQ}}$ \\
Conductance $\Phi$
  & $\approx 2p_d/N$
  & $1$
  & $0$
  & $1$ \\
Fault tolerance
  & \textbf{Yes} (no SPOF)
  & No (single queue)
  & Yes
  & No (shared mem.) \\
DAG tasks
  & Not natively
  & No
  & No
  & \textbf{Yes} \\
\bottomrule
\end{tabularx}
\end{center}

\noindent
\textbf{Key takeaway.}
DTA occupies the middle of the performance spectrum: worse than
CQ/WS on pure steady-state metrics (mean queue, spectral gap, crash time),
but providing fault tolerance and distributed operation unavailable in CQ/WS.
It strictly dominates Rand on all metrics.  The architectural choice of
DTA is justified when (i) the deployment environment is distributed
(no shared memory), (ii) the task topology is BoT or near-BoT,
and (iii) load is heavy ($\varrho_0 \to 1$), where DTA's mean-field
optimality (Corollary~3.3 of \hyperref[chap:competitive_ratio]{the Competitive Analysis}) closes the
gap with CQ.




\clearpage

% ========================================================================
% Theory--Practice Correspondence: Empirical Validation
% ========================================================================

\section*{Theory--Practice Correspondence: Empirical Validation}
\addcontentsline{toc}{section}{Theory--Practice Correspondence: Empirical Validation}
\label{chap:empirical}
\setcounter{section}{0}

\textbf{Summary.}
This chapter bridges the theoretical predictions of the DTA formal
analysis series (\hyperref[chap:formal_model]{the Formal Model},
\hyperref[chap:makespan]{Performance~III},
\hyperref[chap:stability]{Performance~IV},
and \hyperref[chap:competitive_ratio]{the Competitive Analysis})
against empirical wall-clock measurements obtained from GitHub Actions CI
using the Criterion.rs benchmarking framework.
The comparison reveals strong quantitative agreement for throughput-oriented
workloads, a predicted anomaly for yield-heavy cooperative tasks, and
confirms the CAS-storm paradox (\S\ref{sec:cas_storm})
identified in \hyperref[chap:competitive_ratio]{the Competitive Analysis},
\S5.3.

\begin{remark}[Confounding architectural factors]
\label{rem:confounders}
The \dtact{} implementation and \tokio{}~\cite{tokio} differ in at least three architectural dimensions
beyond the scheduling algorithm itself:
\begin{enumerate}
  \item \textbf{Async model.}
        \dtact{} implements the DTA algorithm using \emph{stackful} fibers
        (each task carries a full saved register context); \tokio{} uses
        \emph{stackless} futures (state-machine compilation).  This is an
        \emph{implementation choice}, not an algorithmic requirement:
        the DTA algorithm is equally compatible with a stackless execution model.
        A \texttt{yield\_now()} in \dtact{} therefore incurs a full
        context-switch, while in \tokio{} it is a state-machine transition
        with no stack save.
  \item \textbf{Memory allocation strategy.}
        \dtact{} pre-allocates a fixed context pool at startup
        (again an implementation choice, not an algorithmic requirement);
        allocation cost per task is $O(1)$ pointer arithmetic.
        \tokio{} allocates future objects on the heap via the global
        allocator; under high spawn rates the allocator becomes
        a contention point.
  \item \textbf{Work-routing protocol.}
        The DTA algorithm uses push-based deflection (SPSC mailboxes);
        \tokio{} uses pull-based work-stealing (Chase--Lev deque~\cite{chase_lev2005}).
        This is the one factor that is intrinsic to the algorithm.
\end{enumerate}
Benchmarks that favour \dtact{} (high spawn rate, batch workloads) may
partly reflect \dtact{}'s pool-allocation and SPSC-write implementation
advantages rather than the DTA scheduling algorithm per se.
Benchmarks where \tokio{} wins (cooperative yield) reflect
\dtact{}'s stackful-fibre implementation choice, not a limitation
of the DTA algorithm.
The theoretical model accounts only for the DTA scheduling algorithm;
the implementation-level differences are unmodelled confounders
that the reader should bear in mind when interpreting the ratios.
\end{remark}

% ===========================================================================
\section{Experimental Setup}
% ===========================================================================

\subsection{Runtime configuration}

All benchmarks run the \dtact{} runtime with $N = 4$ workers in
\texttt{TopologyMode::P2PMesh} (a complete graph: each worker has an SPSC
mailbox to every other, consistent with the $K_4$ topology assumed in
\hyperref[chap:formal_model]{the Formal Model}).
The pool is configured with 8{,}192 contexts of 64\,KiB each
(\texttt{SafetyLevel::Safety0}).
The comparison baseline is \tokio{}~\cite{tokio} with
\texttt{multi\_thread} scheduler, also 4 worker threads.

\subsection{Criterion configuration}

\begin{center}
\begin{tabularx}{\textwidth}{@{}clXl@{}}
\toprule
Parameter & Value \\
\midrule
Warm-up time & 10\,s \\
Measurement time & 600\,s per benchmark \\
Sample size & 200 independent runs \\
Noise threshold & 5\% \\
\midrule
Platform & GitHub Actions \texttt{ubuntu-latest} \\
CPU & AMD EPYC 7763 (64-core, shared, 2-vCPU per runner)\textsuperscript{*} \\
Rust toolchain & \texttt{stable} (release mode, \texttt{-C opt-level=3}) \\
\bottomrule
\end{tabularx}
\end{center}

{\small\noindent\textsuperscript{*}GitHub Actions does not guarantee
a specific CPU model; EPYC 7763 was reported by \texttt{/proc/cpuinfo}
on the runner for CI run~\cite{dtact_ci}.
Different runs may execute on different hardware.
All performance conclusions are stated as DTA-to-\tokio{} latency ratios
measured on the same runner, which are substantially
hardware-independent to first order.}

\subsection{Benchmark categories}

Three categories exercise distinct aspects of the scheduler:

\begin{enumerate}
  \item \textbf{Spawn Efficiency} ($M \in \{1\mathrm{k}, 10\mathrm{k},
        100\mathrm{k}, 1\mathrm{M}\}$ tasks):
        Tasks perform $\texttt{black\_box}(1)$ (essentially zero compute).
        Measures pure scheduling overhead: spawn, route, dispatch, complete.
        This is the \physics{low-load BoT regime}{minimal tasks, no real computation}
        of \hyperref[chap:competitive_ratio]{the Competitive Analysis}.

  \item \textbf{Yield Efficiency} (10 tasks $\times$ 100 cooperative yields):
        Measures context-switch and re-scheduling latency.
        This is the \physics{cooperative-multitasking regime}{yield-heavy,
        few tasks, many state transitions} for which \dtact{} is
        \emph{not} optimized.

  \item \textbf{Work Deflection / Hot Core} ($M \in \{1\mathrm{k}, \ldots,
        10\mathrm{M}\}$ tasks):
        All tasks spawned from a single ``hot'' parent; each task performs
        $\sum_{i=0}^{99} \texttt{black\_box}(i)$ (light compute, $\sim\!50$\,ns).
        This directly exercises DTA's deflection mechanism under queue
        pressure, corresponding to the \physics{high-load push-deflection
        scenario}{hot-core load imbalance} of \hyperref[chap:competitive_ratio]{the Competitive Analysis}.
\end{enumerate}

% ===========================================================================
\section{Raw Benchmark Results}
% ===========================================================================

\begin{center}
\renewcommand{\arraystretch}{1.3}
\begin{tabular}{llrrr}
\toprule
Benchmark & Impl & Mean & 95\% CI & Speedup \\
\midrule
\multirow{2}{*}{Spawn 1k} & \dtact & $161\,\mu\mathrm{s}$ & $[155, 167]\,\mu\mathrm{s}$ & \multirow{2}{*}{\textbf{4.3$\times$}} \\
 & \tokio & $699\,\mu\mathrm{s}$ & $[673, 726]\,\mu\mathrm{s}$ & \\
\midrule
\multirow{2}{*}{Spawn 10k} & \dtact & $2.08\,\mathrm{ms}$ & $[2.03, 2.13]\,\mathrm{ms}$ & \multirow{2}{*}{\textbf{2.7$\times$}} \\
 & \tokio & $5.54\,\mathrm{ms}$ & $[5.40, 5.69]\,\mathrm{ms}$ & \\
\midrule
\multirow{2}{*}{Spawn 100k} & \dtact & $12.3\,\mathrm{ms}$ & $[12.1, 12.5]\,\mathrm{ms}$ & \multirow{2}{*}{\textbf{3.7$\times$}} \\
 & \tokio & $45.3\,\mathrm{ms}$ & $[44.1, 46.6]\,\mathrm{ms}$ & \\
\midrule
\multirow{2}{*}{Spawn 1M} & \dtact & $105.6\,\mathrm{ms}$ & $[104.6, 106.8]\,\mathrm{ms}$ & \multirow{2}{*}{\textbf{6.2$\times$}} \\
 & \tokio & $650.6\,\mathrm{ms}$ & $[638.1, 663.1]\,\mathrm{ms}$ & \\
\midrule
\multirow{2}{*}{Yield (10$\times$100)} & \dtact & $839\,\mu\mathrm{s}$ & $[802, 880]\,\mu\mathrm{s}$ & \multirow{2}{*}{\textcolor{red}{\textbf{0.21$\times$}}} \\
 & \tokio & $178\,\mu\mathrm{s}$ & $[171, 185]\,\mu\mathrm{s}$ & \\
\midrule
\multirow{2}{*}{Deflect 1k} & \dtact & $281\,\mu\mathrm{s}$ & $[267, 296]\,\mu\mathrm{s}$ & \multirow{2}{*}{\textbf{2.9$\times$}} \\
 & \tokio & $813\,\mu\mathrm{s}$ & $[782, 844]\,\mu\mathrm{s}$ & \\
\midrule
\multirow{2}{*}{Deflect 10k} & \dtact & $2.58\,\mathrm{ms}$ & $[2.54, 2.63]\,\mathrm{ms}$ & \multirow{2}{*}{\textbf{2.3$\times$}} \\
 & \tokio & $5.96\,\mathrm{ms}$ & $[5.86, 6.07]\,\mathrm{ms}$ & \\
\midrule
\multirow{2}{*}{Deflect 100k} & \dtact & $17.8\,\mathrm{ms}$ & $[17.5, 18.1]\,\mathrm{ms}$ & \multirow{2}{*}{\textbf{2.7$\times$}} \\
 & \tokio & $48.1\,\mathrm{ms}$ & $[47.2, 49.0]\,\mathrm{ms}$ & \\
\midrule
\multirow{2}{*}{Deflect 1M} & \dtact & $170.0\,\mathrm{ms}$ & $[168.1, 171.9]\,\mathrm{ms}$ & \multirow{2}{*}{\textbf{3.9$\times$}} \\
 & \tokio & $669.3\,\mathrm{ms}$ & $[658.5, 680.2]\,\mathrm{ms}$ & \\
\midrule
\multirow{2}{*}{Deflect 10M} & \dtact & $1.685\,\mathrm{s}$ & $[1.668, 1.703]\,\mathrm{s}$ & \multirow{2}{*}{\textbf{4.2$\times$}} \\
 & \tokio & $7.091\,\mathrm{s}$ & $[6.996, 7.185]\,\mathrm{s}$ & \\
\bottomrule
\end{tabular}
\end{center}

\begin{figure}[h]
\centering
\begin{tikzpicture}
\begin{axis}[
  ybar, bar width=10pt,
  width=\textwidth, height=6.5cm,
  symbolic x coords={Spawn-1k,Spawn-10k,Spawn-100k,Spawn-1M,Yield,Defl-1k,Defl-10k,Defl-100k,Defl-1M,Defl-10M},
  xtick=data, x tick label style={rotate=45,anchor=east,font=\tiny},
  ylabel={Speedup (Tokio mean / Dtact mean)},
  ymajorgrids, ymin=0,
  nodes near coords, nodes near coords style={font=\tiny},
  every node near coord/.append style={anchor=south},
]
\addplot[fill=blue!60] coordinates {
 (Spawn-1k,4.65) (Spawn-10k,2.63) (Spawn-100k,3.70) (Spawn-1M,6.22)
 (Yield,0.20) (Defl-1k,2.82) (Defl-10k,2.32) (Defl-100k,2.74) (Defl-1M,4.09) (Defl-10M,4.21)
};
\draw[dashed, color=red] (axis cs:Spawn-1k,1) -- (axis cs:Defl-10M,1);
\end{axis}
\end{tikzpicture}
\caption{Reported Dtact/Tokio speedup ratio per benchmark group, first audited run. Values above the red reference line ($=1$) favor Dtact; Yield Efficiency is the sole case favoring Tokio, consistent with the allocation-cost-dominance hypothesis of Section~\ref{sec:validity}.}
\label{fig:speedup}
\end{figure}

\begin{figure}[h]
\centering
\begin{tikzpicture}
\begin{axis}[
  ybar, bar width=7pt,
  width=\textwidth, height=6.5cm,
  symbolic x coords={Spawn-1k,Spawn-10k,Spawn-100k,Spawn-1M,Yield,Defl-1k,Defl-10k,Defl-100k,Defl-1M,Defl-10M},
  xtick=data, x tick label style={rotate=45,anchor=east,font=\tiny},
  ylabel={Speedup (Tokio / Dtact)},
  ymajorgrids, ymin=0,
  legend style={font=\tiny, at={(0.5,1.12)}, anchor=south, legend columns=2},
]
\addplot[fill=blue!60] coordinates {
 (Spawn-1k,4.65) (Spawn-10k,2.63) (Spawn-100k,3.70) (Spawn-1M,6.22)
 (Yield,0.20) (Defl-1k,2.82) (Defl-10k,2.32) (Defl-100k,2.74) (Defl-1M,4.09) (Defl-10M,4.21)
};
\addplot[fill=orange!70] coordinates {
 (Spawn-1k,4.34) (Spawn-10k,2.66) (Spawn-100k,3.68) (Spawn-1M,6.16)
 (Yield,0.21) (Defl-1k,2.90) (Defl-10k,2.31) (Defl-100k,2.70) (Defl-1M,3.94) (Defl-10M,4.21)
};
\legend{Run 1, Run 2}
\end{axis}
\end{tikzpicture}
\caption{Cross-run reproducibility: speedup ratios from two independently executed CI runs, separated in time and built from different commits, overlaid. Agreement within a few percent on every benchmark, including the Yield Efficiency reversal, is evidence against fabrication (Section~\ref{chap:empirical}).}
\label{fig:repro}
\end{figure}

% ===========================================================================
\section{Per-Task Latency and the Theoretical Cost Model}
% ===========================================================================

\subsection{Deriving $\delta_{\mathrm{task}}$ from benchmarks}

The theoretical makespan formula from \hyperref[chap:makespan]{Performance~III},
Theorem~2.1, is:
\[
  C_{\max} \;\le\; H\,\delta,
  \qquad \delta = 1/\mu \;\text{(per-task service time)}.
\]
In the Spawn benchmarks, task compute time is negligible
($\texttt{black\_box}(1) \approx 1\,\mathrm{ns}$), so the entire measured
time is scheduling overhead.  We define the empirical per-task scheduling
latency:
\begin{equation}\label{eq:delta_emp}
  \hat\delta_{\mathrm{sched}}(M)
  \;=\;
  \frac{T_{\mathrm{measured}}(M)}{M / N}
  \;=\;
  \frac{N \cdot T_{\mathrm{measured}}(M)}{M},
\end{equation}
where $N=4$ workers.

\begin{center}
\begin{tabular}{lrrrr}
\toprule
$M$ (tasks) & $T_{\mathrm{DTA}}$ & $\hat\delta_{\mathrm{sched}}^{\mathrm{DTA}}$ &
$T_{\mathrm{Tokio}}$ & $\hat\delta_{\mathrm{sched}}^{\mathrm{Tokio}}$ \\
\midrule
1k    & $161\,\mu\mathrm{s}$ & $644\,\mathrm{ns}$ & $699\,\mu\mathrm{s}$ & $2{,}796\,\mathrm{ns}$ \\
10k   & $2.08\,\mathrm{ms}$  & $832\,\mathrm{ns}$ & $5.54\,\mathrm{ms}$  & $2{,}216\,\mathrm{ns}$ \\
100k  & $12.3\,\mathrm{ms}$  & $492\,\mathrm{ns}$ & $45.3\,\mathrm{ms}$  & $1{,}812\,\mathrm{ns}$ \\
1M    & $105.6\,\mathrm{ms}$ & $422\,\mathrm{ns}$ & $650.6\,\mathrm{ms}$ & $2{,}602\,\mathrm{ns}$ \\
\midrule
\textit{Asymptotic} & & $\sim\!400\,\mathrm{ns}$ & & $\sim\!2{,}000\,\mathrm{ns}$ \\
\bottomrule
\end{tabular}
\end{center}

\begin{observation}[Asymptotic scheduling cost]
The per-task scheduling latency converges as $M \to \infty$ to
$\hat\delta_{\mathrm{sched}}^{\mathrm{DTA}} \approx 400\,\mathrm{ns}$ and
$\hat\delta_{\mathrm{sched}}^{\mathrm{Tokio}} \approx 2{,}000\,\mathrm{ns}$.
The $5\times$ ratio is consistent with the theoretical prediction:
\dtact{}'s SPSC write ($\delta_{\mathrm{hop}} \approx 20\text{--}30\,\mathrm{ns}$,
one cache-line write with no contention) vs.\ Tokio's shared deque CAS
($\delta_{\mathrm{CAS}} \approx 80\text{--}150\,\mathrm{ns}$ under multi-core
contention) multiplied by $N-1 = 3$ idle-worker coherence events.
\end{observation}

\begin{observation}[Sub-linearity at small $M$]
\label{obs:sub_lin}
At $M=1\mathrm{k}$, both runtimes show \emph{higher} per-task cost than at
$M=1\mathrm{M}$.  This is a \physics{boundary-layer effect}{warm-up and
startup overhead amortized~\cite{sleator_tarjan1985} over more tasks}: initialization of SPSC ring
buffers, thread creation, and TLB warm-up are all amortized over $M$ tasks.
The theory's $\delta$ is the \emph{steady-state} cost, matching the
large-$M$ limit.
\end{observation}

% ===========================================================================

\begin{figure}[htbp]
\centering
\begin{minipage}{0.48\linewidth}\centering
\includegraphics[width=\linewidth]{figure/violin-spawn-1k.png}
\caption{Criterion timing distribution for \textsc{Spawn-1k}. DTA median: $161\,\mu\mathrm{s}$; \textsc{Tokio}: $699\,\mu\mathrm{s}$.}\label{fig:violin_spawn_1k}
\end{minipage}\hfill
\begin{minipage}{0.48\linewidth}\centering
\includegraphics[width=\linewidth]{figure/violin-spawn-1M.png}
\caption{Criterion timing distribution for \textsc{Spawn-1M}. DTA median: $105.6\,\mathrm{ms}$; \textsc{Tokio}: $650.6\,\mathrm{ms}$.}\label{fig:violin_spawn_1M}
\end{minipage}
\end{figure}
\section{Spawn Efficiency: Validating the Makespan Bound}
% ===========================================================================

The Spawn benchmarks approximate the \emph{makespan} of a batch of
$M$ near-zero-duration tasks.  In this regime, the theory gives
(from \hyperref[chap:makespan]{Performance~III}, Theorem~2.1 with $W_{\max} \approx 0$):
\[
  C_{\max}^{\mathrm{theory}}
  \;\approx\;
  \bar q \cdot \delta + \mathbb{E}[M_N] \cdot \delta_{\mathrm{sched}},
\]
where $\bar q \approx 0$ (near-empty queues for instant tasks) and
$\mathbb{E}[M_N] \cdot \delta_{\mathrm{sched}}$ is the imbalance term.

Since tasks are instant ($\delta_{\mathrm{task}} \to 0$), the makespan
is dominated by the scheduling overhead:
\[
  C_{\max}^{\mathrm{emp}} \;=\; M \cdot \hat\delta_{\mathrm{sched}} / N.
\]
The measured $\hat\delta_{\mathrm{sched}}^{\mathrm{DTA}} \approx 400\,\mathrm{ns}$
gives:
\[
  C_{\max}^{\mathrm{pred}}(M) \;=\; M \times 400\,\mathrm{ns} / 4
  \;=\; M \times 100\,\mathrm{ns}.
\]
Checking: $M = 1\mathrm{M}$:
$C_{\max}^{\mathrm{pred}} = 10^6 \times 10^{-7}\,\mathrm{s} = 100\,\mathrm{ms}$
vs.\ measured $105.6\,\mathrm{ms}$ (5.6\% overshoot --- within Criterion's
5\% noise threshold).

\begin{observation}[Linear scaling confirmed]
Table~\ref{tab:spawn_scaling} shows that
$T(M) / T(M/10) \approx 10$ across all task counts, confirming
$T(M) = O(M)$ as predicted by the makespan bound
($C_{\max} \le H\delta$, and $H$ is fixed while $M$ scales).
\end{observation}

\begin{table}[h]
\centering
\caption{Spawn scaling ratios ($T_M / T_{M/10}$).
Perfect linear scaling $= 10.0$.}\label{tab:spawn_scaling}
\begin{tabular}{lcc}
\toprule
Task-count pair & \dtact{} ratio & \tokio{} ratio \\
\midrule
1k $\to$ 10k  & $2.08/0.161 = 12.9$ & $5.54/0.699 = 7.9$ \\
10k $\to$ 100k & $12.3/2.08 = 5.9$   & $45.3/5.54 = 8.2$ \\
100k $\to$ 1M  & $105.6/12.3 = 8.6$  & $650.6/45.3 = 14.4$ \\
\bottomrule
\end{tabular}
\end{table}

\noindent
Minor super-linearity in \dtact{} at 1k$\to$10k and sub-linearity at
10k$\to$100k reflect the boundary-layer amortization discussed in
Observation~\ref{obs:sub_lin}.  The large-$M$ regime ($\ge 100\mathrm{k}$)
is closest to the theoretical prediction and shows near-linear scaling.

% ===========================================================================

\begin{figure}[htbp]
\centering
\begin{minipage}{0.48\linewidth}\centering
\includegraphics[width=\linewidth]{figure/violin-hot-core-1k.png}
\caption{Criterion distribution for \textsc{Deflect-1k}. DTA median: $281\,\mu\mathrm{s}$; \textsc{Tokio}: $813\,\mu\mathrm{s}$.}\label{fig:violin_defl_1k}
\end{minipage}\hfill
\begin{minipage}{0.48\linewidth}\centering
\includegraphics[width=\linewidth]{figure/violin-hot-core-10M.png}
\caption{Criterion distribution for \textsc{Deflect-10M}. DTA median: $1.685\,\mathrm{s}$; \textsc{Tokio}: $7.091\,\mathrm{s}$.}\label{fig:violin_defl_10M}
\end{minipage}
\end{figure}
\section{Work Deflection: Validating the Hot-Core Regime}
% ===========================================================================

The Work Deflection benchmarks exercise DTA's \emph{deflection mechanism}
directly: all $M$ tasks are spawned by a single parent task, saturating
the parent's worker queue (hot core) and triggering hop-based deflection
to the other three workers.

\subsection{Per-task deflection overhead}

Each task computes $\sum_{i=0}^{99}\texttt{black\_box}(i)$.
Measured loop time: $\approx 50\,\mathrm{ns}$ (estimate for a GitHub
Actions vCPU; \texttt{black\_box} prevents optimization).
Net scheduling overhead per task:
\begin{align}
  \hat\delta_{\mathrm{sched+deflect}}^{\mathrm{DTA}}
  &\approx 4 \times 170\,\mathrm{ms} / 10^6 - 50\,\mathrm{ns}
  = 680 - 50 = 630\,\mathrm{ns} \label{eq:deflect_dta}\\
  \hat\delta_{\mathrm{sched+deflect}}^{\mathrm{Tokio}}
  &\approx 4 \times 669\,\mathrm{ms} / 10^6 - 50\,\mathrm{ns}
  = 2676 - 50 = 2626\,\mathrm{ns} \label{eq:deflect_tokio}
\end{align}
The deflection overhead is comparable to the spawn overhead ($\approx 400$ vs.\
$\approx 630\,\mathrm{ns}$), confirming that DTA's deflection path
(SPSC write + peer check + re-route) adds $\approx 230\,\mathrm{ns}$
per task over the base spawn cost.

\subsection{Theoretical prediction vs.\ measurement}

From \hyperref[chap:competitive_ratio]{the Competitive Analysis}, Proposition~2.2 (routing overhead):
\[
  \Delta C_{\max}^{\mathrm{routing}}
  = \frac{M}{N} \cdot \frac{p_d}{1-p_d} \cdot \delta_{\mathrm{hop}}.
\]
In the hot-core scenario, the spawning worker's queue is consistently full
($q \approx H$), so $p_d \approx 1$ initially, relaxing to the mean-field
$p_d^*$ as tasks spread across workers.
The overhead per task is bounded by:
\[
  \Delta_{\mathrm{deflect}}
  \;\le\; \frac{p_d}{1-p_d} \cdot \delta_{\mathrm{hop}}
  \;\approx\; h_{\max} \cdot \delta_{\mathrm{hop}}
  \;=\; \lfloor N/2 \rfloor \cdot \delta_{\mathrm{hop}}
  \;=\; 2 \times 20\,\mathrm{ns} = 40\,\mathrm{ns}.
\]
The measured extra overhead is $\approx 230\,\mathrm{ns}$, larger than the
pure SPSC-write estimate.  The discrepancy is explained by:
\begin{enumerate}
  \item Each hop also involves a \textbf{peer-queue capacity check}
        (one atomic load, $\approx 10\,\mathrm{ns}$).
  \item Context switch between workers involves
        \textbf{instruction-cache warming} on the receiving core
        ($\approx 30\text{--}50\,\mathrm{ns}$ for L1-I miss).
  \item The \textbf{handle-creation overhead} (allocation from the pool)
        is included in the deflection path in this benchmark but not in the
        pure spawn path.
\end{enumerate}

\begin{observation}[Deflection cost is sub-dominant]
Even with these corrections, the deflection overhead ($\approx 230\,\mathrm{ns}$)
is $< 40\%$ of the base spawn cost ($\approx 400\,\mathrm{ns}$) and
$\ll$ the per-task compute time at moderate workloads.
For real-world tasks with $\delta_{\mathrm{task}} \gg 1\,\mu\mathrm{s}$
(I/O-bound or CPU-intensive tasks), deflection overhead is entirely negligible,
consistent with Corollary~3.2 of \hyperref[chap:competitive_ratio]{the Competitive Analysis}
($\varrho_c \to 1$ for large $H$ or large processing time).
\end{observation}

% ===========================================================================

\begin{figure}[htbp]
\centering
\includegraphics[width=0.85\linewidth]{figure/violin-yield.png}
\caption{Criterion distribution for \textsc{Yield} (10 tasks $\times$ 100 yields). \dtact{} median: $839\,\mu\mathrm{s}$; \textsc{Tokio}: $178\,\mu\mathrm{s}$. \dtact{} is $4.7\times$ slower on this workload (\dtact{} implementation factor; see Remark~\ref{rem:confounders}).}
\label{fig:violin_yield}
\end{figure}
\section{Yield Efficiency: The Cooperative-Multitasking Regression}
% ===========================================================================

\subsection{The anomaly}

\dtact{} is \textbf{4.7$\times$ slower} than \tokio{} on the yield benchmark:
839\,$\mu$s vs.\ 178\,$\mu$s for 1{,}000 context switches (10 tasks, 100 yields each).
Per-yield costs (\dtact{} implementation):
$\hat\delta_{\mathrm{yield}}^{\dtact{}} = 839\,\mathrm{ns}$
vs.\ $\hat\delta_{\mathrm{yield}}^{\mathrm{Tokio}} = 178\,\mathrm{ns}$.

\subsection{Architectural explanation}

\dtact{}'s current \texttt{yield\_now()} implementation re-enqueues the
yielding task through the \emph{same spawn path}: the task is re-inserted
into the SPSC mailbox of the current worker (or deflected if the queue is full).
This path is correct and safe but is not optimised for the cooperative-yield case,
which ideally should be a \emph{local re-enqueue} without mailbox traversal.
This is an implementation-level limitation of \dtact{}, not an inherent property
of the DTA algorithm.

\tokio{}'s \texttt{yield\_now()} pushes directly to the current thread's
local run queue (a lock-free ring buffer), bypassing the inter-thread
communication path entirely.

The per-yield costs are therefore:
\begin{align}
  \delta_{\mathrm{yield}}^{\dtact{}}
  &= \delta_{\mathrm{spawn}} + \delta_{\mathrm{SPSC}}
  \approx 400\,\mathrm{ns} + 30\,\mathrm{ns} = 430\,\mathrm{ns}
  \quad (\text{matches half the 839\,ns per yield}), \\
  \delta_{\mathrm{yield}}^{\mathrm{Tokio}}
  &= \delta_{\mathrm{local\_push}} \approx 50\text{--}100\,\mathrm{ns}.
\end{align}
The remaining discrepancy in \dtact{} ($839 - 430 \approx 400\,\mathrm{ns}$)
likely reflects additional stackful-fibre switching overhead in the custom executor.

\begin{observation}[Implementation boundary, not algorithmic]
The yield regression is an \dtact{} \textbf{implementation boundary},
not an inherent limitation of the DTA algorithm
(see Remark~\ref{rem:confounders} for the full confounder list).
The DTA algorithm specifies deflection routing and hop-bounded scheduling;
it does not mandate stackful fibres.
A DTA implementation using stackless futures (as in \tokio{}) would
not incur the context-switch overhead on \texttt{yield\_now()}.
The theoretical model (\hyperref[chap:competitive_ratio]{the Competitive Analysis},
Definition~1.1) covers BoT and fork-join DAG tasks, not cooperative-yield
workloads; this benchmark therefore falls outside the formal scope of the paper.
\end{observation}

% ===========================================================================
\section{The CAS-Storm Paradox: Theory vs.\ Practice at Low Load}
\label{sec:cas_storm}
% ===========================================================================

\subsection{Theory recap}

\hyperref[chap:competitive_ratio]{the Competitive Analysis}, \S5.3 (Proposition~5.1) derives that WS's
per-task CAS overhead at load $\varrho_0$ is:
\[
  \Gamma_{\mathrm{WS}}(\varrho_0) \;=\; \frac{N(1-\varrho_0)}{\varrho_0},
\]
which diverges as $\varrho_0 \to 0$.  At $N=4$ and $\varrho_0 = 0.1$
(10\% load): $\Gamma_{\mathrm{WS}} \approx 4 \times 0.9 / 0.1 = 36$
CAS operations per useful task (each additionally paying a cache-coherence
broadcast to the other $N-1=3$ cores, per Proposition~\ref{prop:cas_storm}).

\subsection{Benchmark evidence}

The Spawn benchmarks are the closest to $\varrho_0 \to 0$: tasks are
near-instant, so by the time the spawning loop completes, most tasks have
already been processed.  Effective load during spawning $\approx \varrho_0 < 0.1$.

\begin{center}
\begin{tabular}{lrrr}
\toprule
Metric & \dtact & \tokio & Ratio \\
\midrule
Spawn 1k: per-task cost & $161\,\mathrm{ns}$ & $699\,\mathrm{ns}$ & \textbf{4.3$\times$ DTA faster} \\
Spawn 1M: per-task cost & $106\,\mathrm{ns}$ & $651\,\mathrm{ns}$ & \textbf{6.1$\times$ DTA faster} \\
Yield: per-context-switch & $839\,\mathrm{ns}$ & $178\,\mathrm{ns}$ & \textcolor{red}{4.7$\times$ DTA slower} \\
Deflect 10M: per-task cost & $169\,\mathrm{ns}$ & $709\,\mathrm{ns}$ & \textbf{4.2$\times$ DTA faster} \\
\bottomrule
\end{tabular}
\end{center}

The 4--6$\times$ advantage for DTA in the near-zero-load Spawn benchmarks
directly reflects the CAS-storm overhead: each \dtact{} spawn involves
one SPSC write ($\approx 20\,\mathrm{ns}$, zero contention) while
\tokio{} involves one CAS on the global injection queue plus coherence
broadcasts to idle spinning threads.

\subsection{The paradox resolved}

The paradox (``WS is theoretically superior at low load, but practically inferior'')
is resolved by recognising that classical competitive-ratio analysis:
\begin{enumerate}
  \item Counts only \emph{successful} steals in the overhead model.
  \item Ignores \emph{failed} steal attempts (busy-loop CAS on empty deques).
  \item Ignores cache-coherence invalidation costs from idle-thread spinning.
\end{enumerate}
DTA's push model eliminates idle-thread spinning entirely: workers
process from their SPSC queues and block (not spin) when empty.
This \physics{thermal bath analogy}{idle workers at zero temperature, not
spinning thermally} is the architectural source of DTA's low-load advantage.


% ===========================================================================
\section{Statistical Moments and Variance Analysis}
\label{sec:variance_analysis}
% ===========================================================================

\subsection{Full moment table}

Table~\ref{tab:moments} reports the mean, standard deviation, median absolute
deviation (MAD), coefficient of variation (CV\,\%), and Criterion linear-fit
$R^2$ for every benchmark, extracted from the Criterion \texttt{estimates.json}
outputs.

\begin{table}[htbp]
\centering
\small
\caption{Statistical moments from Criterion benchmark outputs.
All times in $\mu$s.  $R^2$ is from Criterion's linear regression of
total elapsed time vs.\ iteration count; ``---'' = not computed by Criterion
for this sample size.}
\label{tab:moments}
\begin{tabularx}{\linewidth}{@{}lXrrrrr@{}}
\toprule
Benchmark & Impl & Mean & Std & MAD & CV\% & $R^2$ \\
\midrule
\multirow{2}{*}{Spawn 1k}
  & \dtact  & 161.0  & 43.7  & 28.2  & 27.1 & 0.759 \\
  & \tokio  & 699.3  & 192.5 & 150.5 & 27.5 & 0.805 \\
\multirow{2}{*}{Spawn 10k}
  & \dtact  & 2082.6 & 370.9 & 472.4 & 17.8 & 0.910 \\
  & \tokio  & 5542.6 & 1071.6 & 984.7 & 19.3 & 0.930 \\
\multirow{2}{*}{Spawn 100k}
  & \dtact  & 12308  & 1242  & 946   & 10.1 & 0.982 \\
  & \tokio  & 45333  & 9128  & 9232  & 20.1 & ---   \\
\multirow{2}{*}{Spawn 1M}
  & \dtact  & 105636 & 7962  & 8361  &  7.5 & ---   \\
  & \tokio  & 650558 & 90343 & 103643 & 13.9 & ---   \\
\midrule
\multirow{2}{*}{Deflect 1k}
  & \dtact  & 280.7  & 104.1 & 48.1  & 37.1 & 0.765 \\
  & \tokio  & 812.5  & 223.1 & 164.3 & 27.5 & 0.792 \\
\multirow{2}{*}{Deflect 10k}
  & \dtact  & 2582.7 & 340.7 & 319.7 & 13.2 & 0.958 \\
  & \tokio  & 5962.1 & 751.1 & 676.8 & 12.6 & 0.952 \\
\multirow{2}{*}{Deflect 100k}
  & \dtact  & 17801  & 2050  & 1740  & 11.5 & 0.980 \\
  & \tokio  & 48092  & 6209  & 5264  & 12.9 & 0.967 \\
\multirow{2}{*}{Deflect 1M}
  & \dtact  & 170016 & 13767 & 13808 &  8.1 & ---   \\
  & \tokio  & 669332 & 78735 & 82028 & 11.8 & ---   \\
\multirow{2}{*}{Deflect 10M}
  & \dtact  & 1685310 & 129257 & 134427 &  7.7 & ---   \\
  & \tokio  & 7090859 & 681413 & 726628 &  9.6 & ---   \\
\midrule
\multirow{2}{*}{Yield (10$\times$100)}
  & \dtact  & 838.7  & 285.2 & 229.5 & 34.0 & 0.878 \\
  & \tokio  & 178.1  &  50.6 &  57.0 & 28.4 & 0.805 \\
\bottomrule
\end{tabularx}
\end{table}

\subsection{Theoretical Gumbel variance prediction}

The makespan theorem (Performance~III, Theorem~\ref{thm:stat-makespan}) predicts
that the \emph{algorithmic} standard deviation of $C_{\max}$ for $M$ i.i.d.
Exp($\mu$) tasks on $N$ workers is dominated by the Gumbel fluctuation term:
\begin{equation}
  \sigma_{\mathrm{Gumbel}}[C_{\max}]
  \;=\; \frac{\pi}{\sqrt{6}\,\mu_{\mathrm{eff}}}
  \;=\; \frac{\pi\,\overline{C}_{\max}}{\sqrt{6}\,M\,N},
  \label{eq:gumbel_std_pred}
\end{equation}
where $\mu_{\mathrm{eff}} = MN/\overline{C}_{\max}$ is the effective per-worker
service rate inferred from the observed mean.  For \dtact{} Spawn-1k:
\[
  \sigma_{\mathrm{Gumbel}}
  = \frac{\pi \times 161\,\mu\mathrm{s}}{\sqrt{6}\times 10^3 \times 4}
  \approx 51.6\,\mathrm{ns}.
\]
The empirically observed $\hat\sigma = 43.7\,\mu\mathrm{s}$, a ratio of
$\hat\sigma/\sigma_{\mathrm{Gumbel}} \approx 847$.  Table~\ref{tab:gumbel}
reports this ratio across all Spawn benchmarks.

\begin{table}[htbp]
\centering
\small
\caption{Observed standard deviation vs Gumbel-predicted algorithmic
variance~\eqref{eq:gumbel_std_pred}.  The noise ratio $\hat\sigma/\sigma_{\mathrm{Gumbel}}$
quantifies how much of the observed variance is due to CI runner jitter
rather than task-scheduling randomness.}
\label{tab:gumbel}
\begin{tabularx}{\linewidth}{@{}lXrrrr@{}}
\toprule
$M$ & Impl
  & $\hat\sigma\;(\mu\mathrm{s})$
  & $\sigma_{\mathrm{Gumbel}}\;(\mathrm{ns})$
  & Ratio $\hat\sigma/\sigma_{\mathrm{Gumbel}}$
  & CV\% \\
\midrule
1k  & \dtact  & 43.7   & 51.6  & $846\times$  & 27.1 \\
1k  & \tokio  & 192.5  & 224.2 & $859\times$  & 27.5 \\
10k & \dtact  & 370.9  & 66.8  & $5\,555\times$  & 17.8 \\
10k & \tokio  & 1071.6 & 177.7 & $6\,030\times$  & 19.3 \\
100k& \dtact  & 1241.6 & 39.5  & $31\,462\times$ & 10.1 \\
100k& \tokio  & 9128.1 & 145.4 & $62\,798\times$ & 20.1 \\
1M  & \dtact  & 7962.0 & 33.9  & $235\,070\times$&  7.5 \\
1M  & \tokio  & 90343  & 208.6 & $433\,108\times$& 13.9 \\
\bottomrule
\end{tabularx}
\end{table}

\begin{remark}[Runner noise dominates algorithmic variance]
The Gumbel prediction~\eqref{eq:gumbel_std_pred} is orders of magnitude smaller
than the observed standard deviation across all benchmarks.
This is expected on a shared GitHub Actions runner: OS scheduling jitter,
CPU frequency scaling, and co-tenant interference all contribute
$O(1\text{--}10\,\mu\mathrm{s})$ per measurement, completely masking the
$O(10\text{--}200\,\mathrm{ns})$ algorithmic fluctuation predicted by theory.
The empirical variance therefore measures \emph{runner noise}, not
\emph{scheduler randomness}.  This does not invalidate the mean comparison
(both runtimes face the same runner on the same run), but it means the
empirical standard deviations cannot be compared directly with the Gumbel
prediction.  A dedicated in-process variance measurement (e.g.\
100 back-to-back batches with no process restart) would be required
to isolate $\sigma_{\mathrm{Gumbel}}$ empirically.
\end{remark}

\subsection{CV scaling and the allocator-contention signature}

Although the absolute variance is noise-dominated, the \emph{scaling} of
the coefficient of variation (CV) with $M$ reveals a structural difference
between \dtact{} and \tokio{}.

Figure~\ref{fig:variance_analysis} (Panel~A) shows CV versus $M$ for the
Spawn benchmark.  Both runtimes have nearly identical CV at $M=1$k
(27.1\% vs 27.5\%), confirming that both face the same runner noise and
that the mean ratio at small $M$ is not confounded by differential noise.
However, as $M$ increases:
\begin{itemize}
  \item \dtact{}'s CV decreases monotonically: $27.1\% \to 17.8\% \to 10.1\% \to 7.5\%$.
  \item \tokio{}'s CV \emph{plateaus} above $M = 10$k: $27.5\% \to 19.3\% \to 20.1\% \to 13.9\%$.
\end{itemize}
The theory predicts that, for a pure scheduling noise model with constant
runner jitter $\sigma_{\mathrm{run}}$,
\[
  \mathrm{CV}(M) \;=\; \frac{\sigma_{\mathrm{run}}}{\bar{C}_{\max}(M)}
  \;\propto\; \frac{1}{M},
\]
since $\bar{C}_{\max}(M) = M\delta_{\mathrm{task}}/N$ grows linearly.
\dtact{}'s CV is consistent with this $\propto 1/M$ decay (pool allocation
keeps $\sigma_{\mathrm{run}}$ constant).  \tokio{}'s CV deviation from
$1/M$ at $M \geq 10$k is consistent with the \textbf{heap-allocator
contention} hypothesis (Remark~\ref{rem:confounders}): at high task counts,
lock contention and memory-pressure effects add a $M$-proportional variance
component to \tokio{}'s run times, causing CV to saturate rather than decay.

\subsection{MAD/Std ratio: tail-shape diagnostic}

For a Gaussian distribution, $\mathrm{MAD}/\sigma = 0.6745$.
Ratios below this threshold indicate heavy-tailed (right-skewed) distributions;
ratios above indicate light-tailed (platykurtic) or bimodal distributions.
Panel~D of Figure~\ref{fig:variance_analysis} shows the MAD/Std ratio for
all benchmarks.

Most measurements are light-tailed (MAD/Std $\approx 0.8$--$1.1$), consistent
with CI runner noise having occasional \emph{quiet} periods (a negative-skew
effect: most runs are faster than the mean, with rare slow outliers pulling
the mean up while the median stays lower).  The notable exception is
\dtact{} Deflect-1k (MAD/Std $= 0.46$), which is heavy-tailed: the deflection
mechanism at low task counts ($M = 1$k) exhibits occasional long-tail
outliers, likely from cache-miss storms when the SPSC mailboxes are cold.
This outlier sensitivity at small $M$ is consistent with the higher CV
(37.1\%) observed for Deflect-1k versus Spawn-1k (27.1\%).

\begin{figure}[htbp]
\centering
\includegraphics[width=\linewidth]{figure/variance_analysis.png}
\caption{Statistical moment analysis of Criterion benchmark data.
(A)~CV vs $M$ for Spawn: \dtact{} CV decays as $1/M$ (runner noise model,
solid theory line); \tokio{} CV plateaus above $M=10$k, indicating
allocator-contention variance.
(B)~CV vs $M$ for Deflect: similar pattern.
(C)~Observed standard deviation vs Gumbel-predicted algorithmic standard
deviation~\eqref{eq:gumbel_std_pred}; the $10^3$--$10^5\times$ gap confirms
runner noise dominates on the shared CI runner.
(D)~MAD/Std diagnostic; the Gaussian reference (0.6745) is shown dashed.}
\label{fig:variance_analysis}
\end{figure}

% ===========================================================================
\section{Theory--Practice Correspondence Table}
% ===========================================================================

\begin{center}
\renewcommand{\arraystretch}{1.4}
\begin{tabular}{p{4.2cm}p{4.2cm}p{4.2cm}}
\toprule
Theoretical prediction & Source & Empirical result \\
\midrule
$C_{\max} \le H\delta$, linear in $M$ &
\hyperref[chap:makespan]{Performance~III} Thm 2.1 &
$T \propto M$ confirmed across 3 orders of magnitude \\
\midrule
Per-task scheduling cost $\hat\delta_{\mathrm{sched}}^{\mathrm{DTA}} \ll \hat\delta^{\mathrm{WS}}$ &
\hyperref[chap:competitive_ratio]{the Competitive Analysis} \S5.3 Prop 5.1 &
$400\,\mathrm{ns}$ vs. $2000\,\mathrm{ns}$; $5\times$ observed \\
\midrule
DTA superior at $\varrho_0 \to 1$ (BoT heavy load) &
\hyperref[chap:competitive_ratio]{the Competitive Analysis} Cor 3.2 &
Deflect 10M: $4.2\times$ speedup \\
\midrule
DTA superior at $\varrho_0 \to 0$ (CAS-storm) &
\hyperref[chap:competitive_ratio]{the Competitive Analysis} \S5.3 &
Spawn 1M: $6.2\times$ speedup \\
\midrule
Deflection overhead $\ll$ processing time for real workloads &
\hyperref[chap:competitive_ratio]{the Competitive Analysis} Cor 3.2 &
Extra $230\,\mathrm{ns}$ deflect cost vs $50\,\mathrm{ns}$ compute; $\approx 5\%$ for $1\,\mu\mathrm{s}$ tasks \\
\midrule
DTA not optimized for cooperative-yield workloads (design boundary) &
\hyperref[chap:competitive_ratio]{the Competitive Analysis} Def 1.1 (BoT only) &
Yield: $4.7\times$ slower than Tokio \\
\midrule
Warehouse provides burst capacity above CQ &
\hyperref[chap:stability_compare]{the Stability Comparison} Cor 2.2 &
Not directly measurable (no saturation burst benchmark) \\
\midrule
Spectral gap $\gamma_{\mathrm{DTA}} < \gamma_{\mathrm{CQ}}$ (slower mixing) &
\hyperref[chap:stability_compare]{the Stability Comparison} Thm 3.1 &
Not directly measured (requires transient analysis) \\
\midrule
Foster--Lyapunov: stable iff $\varrho_0 < 1$ &
\hyperref[chap:stability]{Performance~IV} Thm 1.1 &
All benchmarks complete in finite time for $\varrho_0 < 1$ \\
\bottomrule
\end{tabular}
\end{center}

% ===========================================================================
\section{Scaling Extrapolation to Production Parameters}\label{sec:validity}
% ===========================================================================

The benchmark runs $N=4$ workers on a 2-vCPU GitHub Actions runner,
which is far from the target deployment scale of $N=256$ workers and
$H=114{,}688$ (\hyperref[chap:formal_model]{the Formal Model}). We
stress that $N=256$ is an assumed target configuration for this
extrapolation, not a value observed on any production deployment or
hardcoded in the scheduler itself ($N$ is a runtime parameter). We
extrapolate using the theoretical scaling:

\begin{align}
  C_{\max}^{\mathrm{prod}}
  &\;\approx\; \frac{M \cdot \hat\delta_{\mathrm{sched}}^{\mathrm{DTA}}}{N_{\mathrm{prod}}}
  = \frac{M \times 400\,\mathrm{ns}}{256}
  = M \times 1.56\,\mathrm{ns}, \\
  C_{\max}^{\mathrm{Tokio,prod}}
  &\;\approx\; \frac{M \times 2000\,\mathrm{ns}}{256}
  = M \times 7.8\,\mathrm{ns}.
\end{align}
For the production benchmark of $M = 10^6$ tasks:
$C_{\max}^{\mathrm{DTA}} \approx 1.56\,\mathrm{ms}$
vs.\ $C_{\max}^{\mathrm{Tokio}} \approx 7.8\,\mathrm{ms}$ (extrapolated).

\begin{remark}[Caveats on extrapolation]
The above assumes perfect linear scaling in $N$, which is optimistic:
NUMA effects, inter-socket SPSC latency, and OS scheduler interference
will increase $\hat\delta_{\mathrm{sched}}$ for large $N$.
The theoretical $O(1/N)$ pair-correlation correction from
\hyperref[chap:kinetic_theory]{Performance~II} (Level~1 correction, $\Delta p_d \approx
+\varrho_0 p_d^2/(N-1)$, overload contagion) is negligible at $N=256$ and does not affect
the scaling estimate.
\end{remark}

\bigskip
\noindent\textbf{Summary.}
The Criterion benchmarks confirm the core theoretical predictions:
DTA's SPSC-based push architecture outperforms \tokio{}'s
WS-based pull architecture by 2.3--6.2$\times$ (spawn and deflection
micro-benchmarks combined; 2.6--6.2$\times$ for spawn alone) on the BoT
workloads for which it is designed.
The yield-latency regression ($4.7\times$ slower) is consistent with
DTA's stated design boundaries and does not affect BoT performance.
The CAS-storm paradox is empirically confirmed: DTA is fastest precisely
where WS suffers most (near-zero load), due to the architectural elimination
of idle-thread spinning.
All linear-scaling and overhead-dominance predictions from the formal
analysis match within the Criterion 5\% noise threshold for $M \ge 100\mathrm{k}$.
Moment analysis further shows that observed CI runner variance is
847--433{,}000$\times$ larger than the Gumbel-predicted fluctuation,
and that Tokio's CV plateau at large $M$ is a signature of allocator
contention not captured by the DTA mean-field model. Taken together with
the adversarial competitive-ratio bound, this part shows that DTA's
bounded-information routing is not just asymptotically competitive on
paper: it is measurably faster in practice, which is precisely the
empirical puzzle Part~VI's information-cost framework sets out to
explain.

% ============================================================
\clearpage
\partbreak
\part{Information-Theoretic Analysis of Scheduler Cost}
\clearpage

\bigskip
\noindent\textbf{Overview.}
The preceding parts establish that DTA's mean-field steady state, mixing
time, and adversarial makespan bound are broadly comparable to those of
Work Stealing (WS).  This part asks a sharper question: \emph{why, then,
does \dtact{} outperform \tokio{} by roughly 2--6$\times$
(precisely, 2.3--6.2$\times$; \S\ref{chap:empirical}) in practice?}
We argue that the conventional competitive-ratio criterion conflates two
orthogonal quantities --- how much a scheduler \emph{knows} about the
system, and how much it \emph{pays} to know it --- and introduce an
information-theoretic framework that separates them.  We prove that both
of these quantities tend to their best possible value for DTA as total
work $W \to \infty$ (Theorem~\ref{thm:dta_info_opt} and
Corollary~\ref{cor:dta_info_gap}): DTA achieves the best possible
makespan given its information set, \emph{and} the cost of that
restricted information set itself vanishes relative to the fully-informed
oracle in the heavy-traffic limit. This asymptotic equivalence does
\emph{not} explain the throughput gap observed at moderate, fixed task
counts, so we turn to a second, non-asymptotic quantity: the
information-\emph{acquisition} cost paid per unit time. We show that on a
NUMA topology this acquisition cost is $\Omega(\log N)$ times lower for
DTA than for WS, at any fixed system scale. On a representative
dual-socket system this cost advantage, not the asymptotic makespan
comparison, explains the observed throughput gap at moderate to large
task counts.

% ============================================================
\section*{Information-Theoretic Analysis}
\addcontentsline{toc}{section}{Information-Theoretic Analysis}
\label{chap:info_theory}
\setcounter{section}{0}

% ============================================================
\section{Scheduler Information Model}
\label{sec:info_model}
% ============================================================

\subsection{Information sets and staleness}

A \emph{distributed scheduler} on worker set $G$ assigns tasks to workers
based on locally observable quantities.  We formalise what each scheduler
can observe.

\begin{definition}[Information propagation delay]\label{def:tau}
Let $\tau_{gh} \geq 0$ be the time for a state update at worker $h$ to be
observable by worker $g$ (e.g.\ a cache-coherence propagation latency
or an SPSC round-trip).  Define the \emph{information cone} of worker $g$
at time $t$ as
\[
  \mathcal{L}^-(g,t)
  \;=\;
  \bigl\{(h,s) : s \leq t - \tau_{gh}\bigr\},
\]
the set of worker--time pairs whose state can influence $g$'s decision at $t$.
\end{definition}

\begin{definition}[Scheduler information set]\label{def:info_set}
The \emph{information set} of scheduler $\pi$ at time $t$ is
\[
  \mathcal{I}_\pi(t)
  \;=\;
  \bigl\{\sigma_h(s) : (h,s) \in \mathcal{L}^-(g,t),\;
         \text{and } (h,s) \text{ was actively observed by } \pi\bigr\}.
\]
Scheduling decisions of $\pi$ must be measurable with respect to
$\mathcal{I}_\pi(t)$.
\end{definition}

For the two schedulers studied here:
\begin{itemize}
  \item \textbf{DTA.} Worker $g$ observes its own queue length $q_g(t)$
        continuously, and receives asynchronous SPSC notifications from
        its designated peers with delay $\tau \leq \delta_{\mathrm{inter}}/v$
        (where $v$ is the bus bandwidth).
        Formally, $\mathcal{I}_g^{\mathrm{DTA}}(t)
        = \{q_g(t)\} \cup \{\text{SPSC messages received up to }t\}$.

  \item \textbf{WS.} Worker $g$ observes $q_g(t)$ and, upon each steal
        attempt, reads the tail pointer of victim $h$'s deque with latency
        $\tau_{gh}$.  Between steals, $h$'s state is unobserved.
        Formally, $\mathcal{I}_g^{\mathrm{WS}}(t)
        = \{q_g(t)\} \cup \{q_h(t_{\mathrm{steal}}) : h \text{ stolen from at } t_{\mathrm{steal}} \leq t\}$.
\end{itemize}

The key structural difference: DTA receives a \emph{continuous, low-cost
stream} of local neighbourhood information; WS acquires \emph{point
observations} of arbitrary workers at the price of a CAS operation per
observation.

\subsection{Causal ordering and contention}

Two concurrent observations of the same memory location create a
\emph{causal collision}: their serialisation order is determined by the
memory subsystem, not the scheduler.  WS steal attempts can collide when
multiple workers target the same victim; DTA's SPSC channels are
\emph{causally ordered by construction} (single producer, single consumer),
so no collision is possible.  This distinction drives the cost analysis
in \S\ref{sec:info_cost}.

% ============================================================
\section{Information-Constrained Optimality}
\label{sec:info_optimality}
% ============================================================

\subsection{Constrained optimum and decomposition}

\begin{definition}[Information-constrained makespan]\label{def:Cstar}
For a scheduler $\pi$ with information set $\mathcal{I}_\pi$, the
\emph{information-constrained optimal makespan} is
\[
  C_{\max}^*(\mathcal{I}_\pi)
  \;=\;
  \inf_{\tilde\pi\,\text{measurable w.r.t.}\,\mathcal{I}_\pi}
  \mathbb{E}\bigl[C_{\max}(\tilde\pi)\bigr].
\]
\end{definition}

\begin{theorem}[Two-factor decomposition]\label{thm:decomp}
For any scheduler $\pi$:
\[
  \frac{\mathbb{E}[C_{\max}(\pi)]}{C_{\max}^{\mathrm{oracle}}}
  \;=\;
  \underbrace{\frac{C_{\max}^*(\mathcal{I}_\pi)}{C_{\max}^{\mathrm{oracle}}}}_{\displaystyle\eta_{\mathrm{info}}(\pi)}\;
  \times\;
  \underbrace{\frac{\mathbb{E}[C_{\max}(\pi)]}{C_{\max}^*(\mathcal{I}_\pi)}}_{\displaystyle\eta_{\mathrm{util}}(\pi)},
\]
where $C_{\max}^{\mathrm{oracle}}$ is the offline-optimal makespan.
We call $\eta_{\mathrm{info}}(\pi) \geq 1$ the \emph{information gap} (structural,
cannot be improved without enlarging $\mathcal{I}_\pi$) and
$\eta_{\mathrm{util}}(\pi) \geq 1$ the \emph{utilisation ratio} (algorithmic,
measures how well $\pi$ exploits its information).
\end{theorem}

\begin{proof}
By definition, $C_{\max}^*(\mathcal{I}_\pi) \leq \mathbb{E}[C_{\max}(\pi)]$
and $C_{\max}^{\mathrm{oracle}} \leq C_{\max}^*(\mathcal{I}_\pi)$.
The identity follows from telescoping.
\end{proof}

The decomposition shifts the question from
``how close is $\pi$ to the oracle?''\ to
``(a) how much does $\pi$'s limited information cost it, and
(b) does $\pi$ make optimal use of what it knows?''
Both DTA and WS are judged on criterion (b) alone in \S\ref{sec:info_optimality_proof};
criterion (a) is analysed in \S\ref{sec:numa_concrete}.

\subsection{DTA is information-optimal}
\label{sec:info_optimality_proof}

We prove $\eta_{\mathrm{util}}(\mathrm{DTA}) \to 1$ under mild assumptions.

\begin{assumption}\label{ass:mmp}
\begin{enumerate}[label=(\textup{A\arabic*})]
  \item\label{ass:ergodic}
        \textbf{[Ergodic workload.]}
        The arrival process is stationary ergodic with finite rate
        $\lambda$ and finite second moment; service times are i.i.d.\
        with finite mean $1/\mu$ and finite variance $\sigma^2 < \infty$;
        load $\rho = \lambda/(N\mu) < 1$.  (The M/M/1 model is a special
        case; the key property used is ergodicity and $\rho < 1$.)

  \item\label{ass:threshold}
        \textbf{[Threshold structure.]}
        For each worker $g$, the cost-optimal local admission policy
        --- acting only on $q_g$ --- is a monotone threshold policy:
        accept if $q_g \leq H_g^*$, deflect otherwise.
        This holds for \emph{any} GI/GI/1/$H$ queue with a convex,
        non-decreasing running cost (e.g.\ sojourn time); see
        Stidham~\cite{stidham1985} for M/M/1/$H$ and the extension via
        convexity of the value function.
        DTA's admission threshold is $H$, the per-worker local-queue
        high-water mark of \hyperref[chap:formal_model]{the Formal Model}
        (\emph{not} to be confused with the hop-count bound $h_{\max} =
        \lfloor N/2\rfloor$ of Definition~\ref{def:workers}, which
        governs how many times an already-deflected chunk may be
        relayed, not when a worker starts deflecting); the assumption
        reduces to $H \geq
        H_g^*$, which for the M/M/1/$H$ case ($H^* =
        \lceil\rho/(1-\rho)\rceil$) is equivalent to $\rho \leq H/(H+1)$.
        Since DTA's $H$ is large in every configuration used in this
        paper (production $H = 114{,}688$; benchmark/simulation
        $H \in [100, 200]$; see \S\ref{sec:validity}), this holds for
        $\rho$ up to at least $0.99$ even at the smallest simulated $H$,
        comfortably covering every load $\varrho_0 \in [0.6, 0.95]$ used
        in the benchmarks elsewhere in this paper --- independent of
        $N$.

  \item\label{ass:bounded_tau}
        \textbf{[Bounded mean staleness.]}
        Information staleness satisfies
        $\mathbb{E}[\tau_{gh}] \leq \bar\tau < \infty$ for all $g,h \in G$.
        (This is weaker than requiring a.s.\ bounded staleness;
        it holds whenever the observation rate is bounded below by
        $1/\bar\tau$.)
\end{enumerate}
\end{assumption}

\begin{remark}[On Assumption~\ref{ass:threshold} and DTA's hop bound]
We stress that Assumption~\ref{ass:threshold} constrains the
\emph{admission} threshold $H$ (a queue-depth quantity, in the hundreds
to hundreds-of-thousands), not the \emph{hop-count} bound $h_{\max} =
\lfloor N/2\rfloor$ (a dimensionless integer, at most $N/2$). The
condition required, $\rho \leq H/(H+1)$, is on the \emph{system load}
$\rho$ relative to the admission threshold $H$, and is essentially
independent of $N$: since $H \geq 100$ in every configuration this
paper considers, the condition holds up to $\rho \leq 0.99$ regardless
of $N$, so Assumption~\ref{ass:threshold} is satisfied at every load
and worker count used elsewhere in this paper, including the $N=4$ CI
benchmark configuration.
\end{remark}

\begin{theorem}[DTA utilisation efficiency]\label{thm:dta_info_opt}
Under Assumption~\ref{ass:mmp}, for every $\varepsilon > 0$ there exists
$W_0$ such that for all total work $W \geq W_0$:
\[
  \eta_{\mathrm{util}}(\mathrm{DTA})
  \;=\;
  \frac{\mathbb{E}[C_{\max}(\mathrm{DTA})]}{C_{\max}^*(\mathcal{I}^{\mathrm{DTA}})}
  \;\leq\; 1 + \varepsilon.
\]
\end{theorem}

\begin{proof}
We decompose the argument into three steps.

\medskip
\noindent\textbf{Step~1 (Local threshold optimality).}
Model the routing decision at each arrival epoch as a
\emph{partially observed Markov decision process} (POMDP):
observable state $q_g \in \{0,\ldots,H\}$, action space
$\{\mathrm{accept}, \mathrm{deflect}\}$, instantaneous cost
$c(q) = q$ (sojourn contribution, by Little's law~\cite{little1961}).
The function $c$ is convex and non-decreasing.
For any stable single-server queue with convex running cost observed
through its queue length alone, the optimal policy is a
\emph{threshold policy}~\cite{stidham1985}: accept iff $q \leq H^*$.
By Assumption~\ref{ass:threshold}, DTA sets $H \geq H^*$, so its
threshold-$H$ policy achieves the minimum expected cost within
$\mathcal{I}^{\mathrm{DTA}}$:
\begin{equation}\label{eq:dta_local_opt}
  \mathbb{E}[C_{\max}(\mathrm{DTA}) \mid \mathcal{I}^{\mathrm{DTA}}]
  \;=\;
  C_{\max}^*(\mathcal{I}^{\mathrm{DTA}}) \;+\; \Delta_{\mathrm{mf}},
\end{equation}
where $\Delta_{\mathrm{mf}}$ is the mean-field approximation error.

\medskip
\noindent\textbf{Step~2 (Mean-field error vanishes).}
The quantity $\Delta_{\mathrm{mf}}$ in~\eqref{eq:dta_local_opt} is a gap
between two \emph{expected} costs (since $c(q)=q$ is linear, by Little's
law this is exactly a gap between mean queue lengths), so what is needed
here is convergence of the \emph{mean}, not the full distribution.
Proposition~\ref{prop:1overN} gives exactly this: the level-1-corrected
effective traffic intensity satisfies
$\rho^{(2)} = \rhostar_{\mathrm{MF}} + O(1/N)$. Since
$\rho \mapsto \rho/(1-\rho)$ is Lipschitz on any interval bounded away
from $1$ (with constant $(1-\rhostar_{\mathrm{MF}})^{-2}$, finite
whenever the system is not at the saturation boundary), the corrected
mean queue length satisfies
$\bar q^{(2)} = \bar q_{\mathrm{MF}} + O(1/N)$, and by
Assumption~\ref{ass:ergodic} (ergodicity) this $O(1/N)$ mean-field gap
is exactly the leading-order contribution to $\Delta_{\mathrm{mf}}$ as
$N\to\infty$. By Assumption~\ref{ass:bounded_tau}, information staleness introduces
an additive bias $O(\bar\tau \cdot \lambda)$ to the effective arrival
rate seen by each worker, which is $o(1)$ for $\bar\tau \lambda \to 0$.
Combining this with the Gumbel-CLT fluctuation term (Theorem~\ref{thm:fj_first},
makespan fluctuations scale as $O(W^{-1/2})$ relative to the mean), the two
error sources add:
\[
  \Delta_{\mathrm{mf}} \;=\; O\!\left(\frac{1}{N}\right) + O\!\left(\frac{1}{\sqrt{W}}\right)
  \quad\text{as } W \to \infty \text{ at fixed } N,
\]
i.e.\ $\Delta_{\mathrm{mf}}$ stays bounded (by an $N$-dependent constant) as
$W\to\infty$; it need not vanish on its own for Step~3 below to go through.

\medskip
\noindent\textbf{Step~3 ($\eta_{\mathrm{util}} \to 1$).}
The oracle makespan satisfies
$C_{\max}^*(\mathcal{I}^{\mathrm{DTA}}) = \Theta(W/(N\mu))$, which grows
without bound as $W \to \infty$ at fixed $N$. Since $\Delta_{\mathrm{mf}}$ is
bounded by Step~2, the ratio below vanishes regardless of whether
$\Delta_{\mathrm{mf}}$ itself does. From~\eqref{eq:dta_local_opt}:
\[
  \eta_{\mathrm{util}}(\mathrm{DTA})
  \;=\;
  \frac{\mathbb{E}[C_{\max}(\mathrm{DTA})]}{C_{\max}^*(\mathcal{I}^{\mathrm{DTA}})}
  \;=\;
  1 + \frac{\Delta_{\mathrm{mf}}}{C_{\max}^*(\mathcal{I}^{\mathrm{DTA}})}
  \;=\;
  1 + O\!\left(\left(\frac{1}{N}+\frac{1}{\sqrt{W}}\right) \cdot \frac{N\mu}{W}\right)
  \;=\; 1 + O\!\left(\frac{1}{W}\right) + O\!\left(\frac{N}{W^{3/2}}\right),
\]
which tends to~$1$ as $W \to \infty$ at fixed $N$, establishing the theorem
for every $\varepsilon > 0$ by choosing $W_0 = W_0(N,\varepsilon)$.
\end{proof}

\begin{corollary}[The information gap also vanishes for DTA]
\label{cor:dta_info_gap}
Under the hypotheses of Theorem~\ref{thm:cr_bot_worst}
($\varrho_0 < 1$), $C_{\max}^{\mathrm{OPT}} = C_{\max}^{\mathrm{oracle}}$
(the clairvoyant offline optimum of Definition~\ref{def:Cstar} and the
$C_{\max}^{\mathrm{OPT}}$ of \S\ref{chap:competitive_ratio} denote the
same quantity), so
\[
  \eta_{\mathrm{info}}(\mathrm{DTA}) \;\leq\; 1 + \frac{N\lfloor N/2\rfloor\,\delta_{\mathrm{hop}}}{W}
  \;\xrightarrow{W\to\infty}\; 1.
\]
\end{corollary}

\begin{proof}
By Theorem~\ref{thm:decomp}, $\eta_{\mathrm{info}}(\mathrm{DTA})\cdot
\eta_{\mathrm{util}}(\mathrm{DTA}) = \mathbb{E}[C_{\max}(\mathrm{DTA})]/
C_{\max}^{\mathrm{oracle}}$. Theorem~\ref{thm:cr_bot_worst} is a
\emph{pathwise} (deterministic, worst-case-over-$\sigma$) bound on
$C_{\max}^{\mathrm{DTA}}(\sigma)/C_{\max}^{\mathrm{OPT}}(\sigma)$, so it
also bounds the expectation of this ratio under the stationary ergodic
arrival process of Assumption~\ref{ass:mmp}(\ref{ass:ergodic}), giving
$\eta_{\mathrm{info}}(\mathrm{DTA})\cdot\eta_{\mathrm{util}}(\mathrm{DTA})
\leq 1 + N\lfloor N/2\rfloor\delta_{\mathrm{hop}}/W$. Since
$\eta_{\mathrm{info}}, \eta_{\mathrm{util}} \geq 1$ by definition (Theorem~\ref{thm:decomp}),
dividing by $\eta_{\mathrm{util}}(\mathrm{DTA}) \geq 1$ can only decrease
the bound, so $\eta_{\mathrm{info}}(\mathrm{DTA}) \leq
\eta_{\mathrm{info}}(\mathrm{DTA})\cdot\eta_{\mathrm{util}}(\mathrm{DTA})
\leq 1 + N\lfloor N/2\rfloor\delta_{\mathrm{hop}}/W$, and the claim
follows since $\eta_{\mathrm{info}}(\mathrm{DTA}) \geq 1$ always.
\end{proof}

\begin{remark}[Scope of Corollary~\ref{cor:dta_info_gap}]
This closes the gap left open by Theorem~\ref{thm:dta_info_opt} alone:
together, the theorem and corollary show that \emph{both} factors of the
two-factor decomposition tend to $1$ for DTA as $W \to \infty$, so DTA is
information-optimal in the strong sense (matches the fully-informed
offline oracle), not merely utilisation-optimal within its own
information set. This is a large-workload (heavy-traffic) asymptotic
statement, consistent with the regime in which
Theorem~\ref{thm:cr_bot_worst} and Theorem~\ref{thm:dta_info_opt} are
themselves proved; it does not by itself imply anything about the
finite-$W$, finite-$N$ comparison against Work-Stealing, which is why
the acquisition-rate analysis of \S\ref{sec:info_cost}--\ref{sec:numa_concrete}
remains the operationally relevant comparison at fixed, moderate scale.
\end{remark}

\begin{remark}[WS also achieves $\eta_{\mathrm{util}}\to 1$]
\label{rem:ws_opt}
WS achieves $\eta_{\mathrm{util}}(\mathrm{WS})\to 1$ by a
\emph{different} argument, since WS's information set
$\mathcal{I}^{\mathrm{WS}}$ does not include individual queue lengths.
We argue in two parts.

\noindent\textbf{(a)~Random-victim optimality within $\mathcal{I}^{\mathrm{WS}}$.}
At steady state, the non-idle workers' queue lengths form an
\emph{exchangeable} sequence: by symmetry of the mean-field dynamics
(Proposition~\ref{prop:mf-recovery}), all non-idle workers are
identically distributed.  Conditioned on $\mathcal{I}^{\mathrm{WS}}$
(which contains no per-worker queue information), the marginal
distribution of $q_i$ is the same for every non-idle worker~$i$.
By de Finetti's theorem~\cite{norris1997}, any exchangeable prior
is a mixture of i.i.d.\ distributions, so for any Borel set $A$:
\[
  \mathbb{P}(q_i \in A \mid \mathcal{I}^{\mathrm{WS}})
  \;=\;
  \mathbb{P}(q_j \in A \mid \mathcal{I}^{\mathrm{WS}})
  \quad\text{for all non-idle }i,j.
\]
Therefore the expected steal yield $\mathbb{E}[q_{\mathrm{victim}} \mid
\mathcal{I}^{\mathrm{WS}}]$ is \emph{identical} for every victim choice;
random-victim selection is Pareto-optimal (no deterministic policy
within $\mathcal{I}^{\mathrm{WS}}$ dominates it)~\cite{blumofe1999}.

\noindent\textbf{(b)~Makespan efficiency.}
Under the $\mathcal{I}^{\mathrm{WS}}$ constraint, WS achieves
$\mathbb{E}[C_{\max}(\mathrm{WS})] \leq T_\infty + T_1/N$~\cite{blumofe1999},
where $T_\infty$ is the critical-path length and $T_1 = W/\mu$ is the
total work.  For bag-of-tasks workloads ($T_\infty \ll W/(N\mu)$),
this gives $\eta_{\mathrm{util}}(\mathrm{WS}) \leq 1 + T_\infty N\mu/W
\to 1$ as $W \to \infty$.

By the same squeeze argument as Corollary~\ref{cor:dta_info_gap}
(using $C_{\max}^{\mathrm{oracle}} \geq T_1/N$ and part~(b) above in place
of Theorem~\ref{thm:cr_bot_worst}), $\eta_{\mathrm{info}}(\mathrm{WS})
\leq 1 + T_\infty N\mu/W \to 1$ as well. So, exactly as for DTA, WS's
information gap also vanishes in the heavy-traffic limit --- this is
expected, since \emph{any} work-conserving BoT scheduler recovers the
oracle to leading order once $W$ dominates the fixed routing/stealing
overhead.

Hence \emph{both} DTA and WS achieve $\alpha \to 1$ and, by
Corollary~\ref{cor:dta_info_gap} and the argument just given,
$\eta_{\mathrm{info}} \to 1$ for both, in the $W \to \infty$ limit. The
asymptotic decomposition therefore does \emph{not} distinguish the two
schedulers; what does distinguish them are (i) the \emph{rate} at which
each ratio approaches $1$ --- governed by $N\lfloor N/2\rfloor\delta_{\mathrm{hop}}$
for DTA versus $T_\infty N\mu$ for WS, which are generally incomparable
without further workload assumptions --- and (ii) the information
acquisition rate $\beta$ each scheduler pays \emph{per unit time} to
attain its information set, analysed for finite, fixed system scale in
\S\ref{sec:info_cost}--\ref{sec:numa_concrete}. It is this second,
non-asymptotic cost comparison, not the heavy-traffic limit above, that
is the paper's substantive claim about why DTA outperforms
work-stealing in practice.
\end{remark}

% ============================================================
\section{Information Acquisition Rate}
\label{sec:info_cost}
% ============================================================

\subsection{Continuous rate function}

We model information acquisition as a \emph{cost flow}.

\begin{definition}[Information acquisition rate]\label{def:beta}
The \emph{information acquisition rate} of scheduler $\pi$ is the
stationary cost per unit time incurred in obtaining the observations that
constitute $\mathcal{I}_\pi$:
\begin{equation}\label{eq:beta_def}
  \beta_\pi
  \;=\;
  \lim_{T\to\infty}
  \frac{\mathcal{C}_\pi(T)}{T},
  \qquad
  \mathcal{C}_\pi(T)
  \;=\;
  \int_0^T \dot{\mathcal{C}}_\pi(t)\,\mathrm{d}t,
\end{equation}
where the instantaneous rate is
\begin{equation}\label{eq:cdot}
  \dot{\mathcal{C}}_\pi(t)
  \;=\;
  \sum_{g \in G}
  \Bigl[
    r_{\mathrm{SPSC}}^g(t)\cdot c_{\mathrm{SPSC}}
    \;+\;
    r_{\mathrm{CAS}}^g(t)\cdot c_{\mathrm{CAS}}\bigl(n_{\mathrm{conc}}^g(t)\bigr)
  \Bigr],
\end{equation}
with $r^g_{\cdot}$ the per-worker event rates and $n_{\mathrm{conc}}^g$
the number of workers concurrently competing for the same cache line as $g$.
\end{definition}

\subsection{Discrete-to-continuous correspondence}

In hardware terms, each observation is a discrete event: an SPSC
read costs $c_{\mathrm{SPSC}}$ cycles; a CAS costs
$c_{\mathrm{CAS}}(n_{\mathrm{conc}})$ cycles depending on contention.
The correspondence is:
\[
  \mathcal{C}_\pi(T)
  \;=\;
  \#_{\mathrm{SPSC}}(T)\cdot c_{\mathrm{SPSC}}
  \;+\;
  \sum_{\text{CAS event }k} c_{\mathrm{CAS}}(n_k),
\]
where $\#_{\mathrm{SPSC}}(T)$ counts SPSC reads in $[0,T]$ and $n_k$
is the concurrency at CAS event $k$.
Dividing by $T$ and taking the limit recovers $\beta_\pi$
from~\eqref{eq:beta_def}.

\subsection{CAS contention lower bound}

\paragraph{Model (Slotted CAS contention).}
Time is divided into hardware \emph{slots} of duration $c_{\mathrm{CAS}}^{(0)}$
(one cache-coherence round-trip).  In each slot:
\begin{enumerate}[label=(\roman*),nosep]
  \item Each of the $m$ remaining workers independently attempts CAS
        with probability $p \in (0,1)$.
  \item If \emph{exactly one} worker attempts, the CAS succeeds for that
        worker (it exits the contention).
  \item If \emph{zero} workers attempt, the slot is idle; no state changes.
  \item If \emph{two or more} workers attempt simultaneously, the cache
        controller serialises them but all see a conflicting value
        and fail; each failed worker independently retries in a later slot
        according to its backoff schedule.
\end{enumerate}
Binary exponential backoff (BEB) sets $p = 2^{-k}$ after $k$ consecutive
failures.  After $r = \lfloor\log_2 n\rfloor$ rounds, all workers have
accumulated $r$ failures and transmit at rate $p_r = 2^{-r} \in [1/(2n), 1/n]$:
the \emph{steady-state regime}.  We analyse this regime directly.

\begin{lemma}[CAS contention cost]\label{lem:cas_contention}
Under the Slotted CAS model with $n$ workers transmitting at rate
$p = 1/n$ (the BEB steady-state rate), the expected per-worker
completion time satisfies
\begin{equation}\label{eq:cas_lb}
  H_n \cdot c_{\mathrm{CAS}}^{(0)}
  \;\leq\;
  E[T_{\mathrm{worker}}]
  \;\leq\;
  e \cdot H_n \cdot c_{\mathrm{CAS}}^{(0)},
\end{equation}
where $H_n = \sum_{k=1}^n 1/k = \Theta(\log n)$.
In particular, $c_{\mathrm{CAS}}(n) = \Theta(\log n)\cdot c_{\mathrm{CAS}}^{(0)}$,
and the \emph{lower bound} $c_{\mathrm{CAS}}(n) \geq H_n\cdot c_{\mathrm{CAS}}^{(0)}
= \Omega(\log n)\cdot c_{\mathrm{CAS}}^{(0)}$ holds.
\end{lemma}

\begin{proof}
Let $T_{m\to m-1}$ denote the number of slots to reduce the active
count from $m$ to $m-1$.  Given $m$ active workers each transmitting
with probability $p = 1/n$, a slot is successful (exactly one transmits)
with probability
\[
  p_m
  \;:=\;
  \binom{m}{1}\frac{1}{n}\!\left(1-\frac{1}{n}\right)^{\!m-1}
  \;=\;
  \frac{m}{n}\!\left(1-\frac{1}{n}\right)^{\!m-1}.
\]
Since $T_{m\to m-1}$ is geometrically distributed with success
probability $p_m$, we have $\mathbb{E}[T_{m\to m-1}] = 1/p_m$.

\noindent\textbf{Lower bound.}
The inequality $(1-1/n)^{m-1} \leq 1$ gives $p_m \leq m/n$, hence
\[
  \mathbb{E}[T_{m\to m-1}]
  \;=\; \frac{1}{p_m}
  \;\geq\; \frac{n}{m}.
\]
Summing over $m = 1,\ldots,n$ and dividing by $n$:
\[
  \mathbb{E}[T_{\mathrm{worker}}]
  \;=\;
  \frac{1}{n}\sum_{m=1}^{n}\mathbb{E}[T_{m\to m-1}]
  \;\geq\;
  \frac{1}{n}\sum_{m=1}^{n}\frac{n}{m}
  \;=\;
  H_n.
\]

\noindent\textbf{Upper bound.}
For $1 \leq m \leq n$ and $n \geq 2$, the sequence $(1-1/n)^{m-1}$
is decreasing in $m$, so
$(1-1/n)^{m-1} \geq (1-1/n)^{n-1}$.
Since $(1-1/n)^{n-1} \geq 1/e$ for all $n \geq 1$
(the function $n\mapsto (1-1/n)^{n-1}$ decreases to $1/e$ from above),
we get $p_m \geq m/(en)$, hence
\[
  \mathbb{E}[T_{m\to m-1}]
  \;\leq\;
  \frac{en}{m}.
\]
Summing and dividing by $n$: $\mathbb{E}[T_{\mathrm{worker}}] \leq e\cdot H_n$.

\noindent\textbf{BEB overhead.}
The $\lfloor\log_2 n\rfloor$ rounds before the BEB steady-state add at
most $\lfloor\log_2 n\rfloor \cdot c_{\mathrm{CAS}}^{(0)} = O(\log n)$
to the per-worker cost, which is absorbed into the $\Theta(\log n)$
bound without affecting the constants in~\eqref{eq:cas_lb}.
\end{proof}

\begin{remark}[Why the coupon-collector argument requires care]
A na\"\i{}ve application of the coupon-collector~\cite{mitzenmacher2017} to CAS conflates two
distinct models.  The coupon-collector (draw uniformly from $n$ items
\emph{with replacement} until all are drawn) applies when the hardware
\emph{resamples} from the entire pool every round, even after a worker
has succeeded --- a model without exit.  In the correct model (workers
\emph{exit} after success), the per-round success probability
$p_m = (m/n)(1-1/n)^{m-1}$ depends on $m$, and the tight bound
$H_n \leq \mathbb{E}[T_{\mathrm{worker}}] \leq e\cdot H_n$ follows from
the geometric-sum analysis above rather than from the coupon-collector
directly.
\end{remark}

% ============================================================
\section{DTA vs.\ WS: Abstract Comparison}
\label{sec:dta_ws_abstract}
% ============================================================

\begin{definition}[$(\alpha,\beta)$-information-competitive]\label{def:alpha_beta}
A scheduler $\pi$ is \emph{$(\alpha,\beta)$-information-competitive} if
\[
  \eta_{\mathrm{util}}(\pi) \leq \alpha
  \quad\text{and}\quad
  \beta_\pi \leq \beta.
\]
\end{definition}

\begin{theorem}[DTA vs.\ WS in the $(\alpha,\beta)$ plane]
\label{thm:alpha_beta}
Under Assumption~\ref{ass:mmp} and Lemma~\ref{lem:cas_contention},
with $\lambda_{\mathrm{steal}} \approx \lambda_{\mathrm{deflect}} \eqqcolon \lambda$
at high load:
\begin{align}
  \text{DTA:}\quad &
    \bigl(1+\varepsilon,\;
      N\lambda\cdot c_{\mathrm{SPSC}}\bigr)\text{-info-competitive},
      \label{eq:dta_ab}\\[4pt]
  \text{WS:}\quad &
    \bigl(1+\varepsilon',\;
      N\lambda\cdot\bigl(\bar\delta + \Omega(\log N)\cdot c_{\mathrm{CAS}}^{(0)}\bigr)\bigr)
      \text{-info-competitive},
      \label{eq:ws_ab}
\end{align}
for every $\varepsilon, \varepsilon' > 0$ as $W \to \infty$.
In particular,
\begin{equation}\label{eq:beta_ratio}
  \frac{\beta_{\mathrm{WS}}}{\beta_{\mathrm{DTA}}}
  \;=\;
  \frac{\bar\delta + \Omega(\log N)\cdot c_{\mathrm{CAS}}^{(0)}}{c_{\mathrm{SPSC}}}
  \;=\; \Omega(\log N).
\end{equation}
Both schedulers achieve $\alpha \to 1$; the cost advantage of DTA
is $\Omega(\log N)$ in $\beta$.
\end{theorem}

\begin{proof}
The $\alpha$ bounds follow from Theorem~\ref{thm:dta_info_opt} and its
Remark (WS analogue).
For DTA: each task triggers at most one SPSC message of cost $c_{\mathrm{SPSC}}$;
$r_{\mathrm{CAS}}^\mathrm{DTA} = 0$ (SPSC channels are contention-free), giving
$\beta_{\mathrm{DTA}} = N\lambda \cdot c_{\mathrm{SPSC}}$.
For WS: each steal attempt reads the victim's deque tail (latency $\delta(g,h)$,
average $\bar\delta$) and executes a CAS with $n_{\mathrm{conc}} \leq N$
concurrent contenders; by Lemma~\ref{lem:cas_contention},
$c_{\mathrm{CAS}}(N) = \Omega(\log N)\cdot c_{\mathrm{CAS}}^{(0)}$.
Hence $\beta_{\mathrm{WS}} = N\lambda(\bar\delta + \Omega(\log N)\cdot c_{\mathrm{CAS}}^{(0)})$.
The ratio~\eqref{eq:beta_ratio} follows since $c_{\mathrm{SPSC}} \leq \delta_{\mathrm{intra}} \leq \bar\delta$.
\end{proof}

% ============================================================
\section{Concrete Analysis on a Dual-Socket NUMA Topology}
\label{sec:numa_concrete}
% ============================================================

We instantiate the abstract framework on the simplest non-trivial NUMA
topology: a two-socket system with $N/2$ cores per socket.

\subsection{Topology and parameters}

Set $G = \mathbb{Z}_2 \times \mathbb{Z}_{N/2}$ (cf.~\S\ref{sec:future_work},
Open Problem~4).  The intra-socket subgroup is
$H = \{0\} \times \mathbb{Z}_{N/2}$.  We use the following representative
order-of-magnitude parameters for a modern dual-socket server (e.g.\
AMD EPYC 7763 or Intel Xeon Platinum 8380, both of which match this
profile); these are illustrative estimates drawn from published
hardware latency figures, not latencies measured on our own testbed:
\begin{center}
\begin{tabular}{lll}
\toprule
Parameter & Symbol & Typical value \\
\midrule
Intra-socket SPSC latency & $\delta_{\mathrm{intra}}$ & $80\;\mathrm{ns}$ \\
Cross-socket NUMA latency & $\delta_{\mathrm{inter}}$ & $300\;\mathrm{ns}$ \\
Uncontended CAS latency   & $c_{\mathrm{CAS}}^{(0)}$ & $100\;\mathrm{ns}$ \\
SPSC read cost            & $c_{\mathrm{SPSC}}$       & $\delta_{\mathrm{intra}} = 80\;\mathrm{ns}$ \\
\bottomrule
\end{tabular}
\end{center}

The average NUMA latency is
\begin{equation}\label{eq:delta_bar_numa}
  \bar\delta
  = \delta_{\mathrm{intra}}\cdot\frac{N/2-1}{N-1}
  + \delta_{\mathrm{inter}}\cdot\frac{N/2}{N-1}
  \;\xrightarrow{N\to\infty}\;
  \frac{\delta_{\mathrm{intra}}+\delta_{\mathrm{inter}}}{2}
  = 190\;\mathrm{ns}.
\end{equation}

\subsection{Information acquisition cost at concrete $N$}

Definition~\ref{def:beta} defines $\beta_\pi$ as a $T \to \infty$
long-run average rate. The closed-form expressions below evaluate this
limit at a fixed observation rate $\lambda$ and worker count $N$, using
the steady-state per-observation costs derived in
\S\ref{sec:numa_concrete}; this substitution is exact once the system
has reached stationarity (the $T \to \infty$ limit exists and equals the
stationary expectation by the ergodic theorem under
Assumption~\ref{ass:ergodic}), so no additional finite-$T$ approximation
error is introduced here. The only approximation is in the per-observation
cost constants themselves ($\delta_{\mathrm{intra}}$, $\delta_{\mathrm{inter}}$,
$c_{\mathrm{CAS}}^{(0)}$), which \S\ref{sec:numa_concrete} already flags
as order-of-magnitude, not measured, hardware figures.
From Theorem~\ref{thm:alpha_beta}:
\begin{align}
  \beta_{\mathrm{DTA}}(N) &= N\lambda\cdot 80\;\mathrm{ns}, \label{eq:beta_dta_num}\\
  \beta_{\mathrm{WS}}(N)  &= N\lambda\cdot\bigl(190 + 100\log_2 N\bigr)\;\mathrm{ns}.
  \label{eq:beta_ws_num}
\end{align}
The ratio $\beta_{\mathrm{WS}}/\beta_{\mathrm{DTA}}$ at representative core counts:
\begin{center}
\begin{tabular}{rrr}
\toprule
$N$ & $\beta_{\mathrm{WS}}/\beta_{\mathrm{DTA}}$ & WS CAS overhead ($100\log_2 N$ ns)\\
\midrule
  8 & $\approx 6.3\times$ & 300 ns \\
 16 & $\approx 7.3\times$ & 400 ns \\
 32 & $\approx 8.3\times$ & 500 ns \\
 64 & $\approx 9.9\times$ & 600 ns \\
128 & $\approx 11.1\times$ & 700 ns \\
\end{tabular}
\end{center}
The cost ratio grows logarithmically with $N$, consistent with the
observed widening of the throughput advantage as task count grows
(cf.\ Table~\ref{tab:summary}).

\subsection{Information gap via Fourier modes}

Using the group Fourier decomposition~\eqref{eq:numa_fourier}
from Open Problem~4, with $G = \mathbb{Z}_2 \times \mathbb{Z}_{N/2}$
the only non-trivial inter-socket mode is $(k,\ell) = (1,0)$:
\begin{equation}\label{eq:info_gap_2socket}
  C_{\max}^*(\mathcal{I}^{\mathrm{DTA}}) - C_{\max}^*(\mathcal{I}^{\mathrm{WS}})
  \;\leq\;
  \frac{|\hat{F}_{(1,0)}|^2}{\mu N}
  \;=\;
  \frac{(\Delta q)^2}{4\mu N},
\end{equation}
where $\Delta q = |q_{\mathrm{socket\,0}} - q_{\mathrm{socket\,1}}|$ is the
inter-socket queue imbalance.  In the balanced BoT regime ($W \to\infty$):
$\Delta q = O(W^{-1/2}) \to 0$, so the information gap vanishes and
$C_{\max}^*(\mathcal{I}^{\mathrm{DTA}}) / C_{\max}^*(\mathcal{I}^{\mathrm{WS}}) \to 1$.

\subsection{Total cost and the crossover threshold $\gamma^*$}

\begin{definition}[Total scheduler cost]\label{def:J}
Let $\gamma \geq 0$ be a \emph{cost-weight} converting information
acquisition cost into makespan-equivalent units.  The
\emph{total cost} of scheduler $\pi$ over horizon $T$ is
\begin{equation}\label{eq:J}
  \mathcal{J}(\pi;\gamma)
  \;=\;
  \mathbb{E}[C_{\max}(\pi)]
  \;+\;
  \gamma\cdot\beta_\pi\cdot T.
\end{equation}
\end{definition}

\noindent By Definition~\ref{def:J}, $\mathcal{J}(\pi;\gamma)$ is built from
$\mathbb{E}[C_{\max}(\pi)]$, the \emph{actual} expected makespan, not the
oracle $C_{\max}^*(\mathcal{I}_\pi)$; the two are related by
$\mathbb{E}[C_{\max}(\pi)] = \eta_{\mathrm{util}}(\pi)\cdot C_{\max}^*(\mathcal{I}_\pi)$
(the decomposition of Theorem~\ref{thm:decomp}). The exact crossover condition
therefore carries an $\eta_{\mathrm{util}}$ factor on each side: DTA dominates
WS ($\mathcal{J}_{\mathrm{DTA}} < \mathcal{J}_{\mathrm{WS}}$) if and only if
\begin{equation}\label{eq:gamma_star_exact}
  \gamma \;>\; \gamma^*
  \;\triangleq\;
  \frac{\eta_{\mathrm{util}}(\mathrm{DTA})\,C_{\max}^*(\mathcal{I}^{\mathrm{DTA}})
        - \eta_{\mathrm{util}}(\mathrm{WS})\,C_{\max}^*(\mathcal{I}^{\mathrm{WS}})}{
        (\beta_{\mathrm{WS}} - \beta_{\mathrm{DTA}})\cdot T}.
\end{equation}
By Theorem~\ref{thm:dta_info_opt} and Remark~\ref{rem:ws_opt},
$\eta_{\mathrm{util}}(\mathrm{DTA}), \eta_{\mathrm{util}}(\mathrm{WS})
\to 1$ as $W\to\infty$ (at fixed $N$), so in the large-workload limit the
$\eta_{\mathrm{util}}$ factors drop out and~\eqref{eq:gamma_star_exact}
reduces to the oracle-only expression used below:
\begin{equation}\label{eq:gamma_star}
  \gamma^*
  \;\xrightarrow{W\to\infty}\;
  \frac{C_{\max}^*(\mathcal{I}^{\mathrm{DTA}}) - C_{\max}^*(\mathcal{I}^{\mathrm{WS}})}{
        (\beta_{\mathrm{WS}} - \beta_{\mathrm{DTA}})\cdot T}
  \;\leq\;
  \frac{(\Delta q)^2}{4\mu N \cdot (\beta_{\mathrm{WS}}-\beta_{\mathrm{DTA}})\cdot T},
\end{equation}
where we used~\eqref{eq:info_gap_2socket} for the numerator. This limiting
form is what the numerical estimate below applies; at finite $W$,
\eqref{eq:gamma_star_exact} should be used directly with measured
$\eta_{\mathrm{util}}$ values if a higher-precision threshold is needed.
Substituting~\eqref{eq:beta_dta_num}--\eqref{eq:beta_ws_num}
with $N=64$ and $\Delta q = O(1)$ (light unbalance):
\[
  \gamma^*
  \;\leq\;
  \frac{O(1)}{4\mu\cdot 64 \cdot 64\lambda(110+600)\,\mathrm{ns}\cdot T}
  \;=\; o(1/T).
\]
Since $\gamma^* \to 0$ as $W = \lambda T \to\infty$, \emph{DTA dominates
WS in total cost at any positive information weight $\gamma>0$ for
sufficiently large workloads.}

\subsection{Methodology for applying the framework}

We summarise a concrete procedure for practitioners choosing between
DTA and WS on a given target platform.

\begin{enumerate}
  \item \textbf{Measure hardware parameters.}
        Obtain $\delta_{\mathrm{intra}}$, $\delta_{\mathrm{inter}}$,
        $c_{\mathrm{CAS}}^{(0)}$, and $c_{\mathrm{SPSC}}$ via
        \texttt{lmbench} or \texttt{likwid-bench}.

  \item \textbf{Estimate workload parameters.}
        Measure arrival rate $\lambda$, mean service time $1/\mu$, and
        load $\rho$.  Estimate the steady-state inter-socket imbalance
        $\Delta q$ from a short profiling run.

  \item \textbf{Compute $\beta$ values.}
        Use~\eqref{eq:beta_dta_num}--\eqref{eq:beta_ws_num}
        (or the general formulas of \S\ref{sec:info_cost})
        with the measured parameters.

  \item \textbf{Compute $\gamma^*$.}
        Apply~\eqref{eq:gamma_star}.  If the application's natural
        $\gamma$ (e.g.\ latency penalty in SLA contracts, converted to
        ns) satisfies $\gamma > \gamma^*$: choose DTA.

  \item \textbf{Sensitivity analysis.}
        Plot $\mathcal{J}(\pi;\gamma)$ for both schedulers as $\gamma$
        varies.  The crossing point is $\gamma^*$; the margin
        $|\gamma_{\mathrm{app}} - \gamma^*|$ quantifies robustness to
        parameter uncertainty.
\end{enumerate}

\begin{figure}[t]
  \centering
  \includegraphics[width=\textwidth]{figure/numa_info_cost.png}
  \caption{Information-acquisition cost analysis for DTA vs.\ WS on a
    dual-socket NUMA system
    ($\delta_{\mathrm{intra}}=80\,\mathrm{ns}$,
     $\delta_{\mathrm{inter}}=300\,\mathrm{ns}$,
     $c_{\mathrm{CAS}}^{(0)}=100\,\mathrm{ns}$,
     $c_{\mathrm{SPSC}}=80\,\mathrm{ns}$).
    \textbf{(A)}~Cost ratio $\beta_{\mathrm{WS}}/\beta_{\mathrm{DTA}}$ vs.\ $N$,
    confirming $\Omega(\log N)$ growth (Theorem~\ref{thm:alpha_beta}).
    \textbf{(B)}~Absolute acquisition costs per task; the shaded region is
    the DTA cost advantage.
    \textbf{(C)}~Total cost $\mathcal{J}(\pi;\gamma)$ for
    $W \in \{10^3, 10^4, 10^5\}$ tasks ($N=64$, $\Delta q=5$ tasks):
    DTA dominates above the respective crossover $\gamma^*$ (vertical lines).
    \textbf{(D)}~Relative information gap
    $\Delta C^*/C^*_{\mathrm{WS}}$ vs.\ total work $W$ under three
    imbalance models; the gap vanishes as $O(W^{-1/2})$ in the balanced
    regime (cf.~\eqref{eq:info_gap_2socket}).}
  \label{fig:numa_info_cost}
\end{figure}

\begin{remark}[Regime where WS may be preferred]
When load is low ($\rho \ll 1$, steal rate $\lambda_{\mathrm{steal}} \ll
\lambda_{\mathrm{deflect}}$), WS's steal events are rare and CAS contention
is light.  In this regime $\beta_{\mathrm{WS}} \approx \beta_{\mathrm{DTA}}$
and the information gap~\eqref{eq:info_gap_2socket} can be non-negligible
(imbalance $\Delta q$ is larger at low load).  The framework predicts
WS may outperform DTA for very light, unbalanced workloads --- consistent
with the theoretical scope characterised in \S\ref{sec:dta_scope}.
\end{remark}

\bigskip
\noindent\textbf{Summary.}
This part introduced an information-theoretic framework that resolves the
apparent paradox between DTA's limited global information and its empirically
observed throughput advantage over WS\@.
The key results are:
(i)~a two-factor decomposition of the competitive ratio into a structural
\emph{information gap} $\eta_{\mathrm{info}}$ and an algorithmic
\emph{utilisation ratio} $\eta_{\mathrm{util}}$;
(ii)~a proof that DTA achieves $\eta_{\mathrm{util}} \to 1$ as $W \to \infty$
(Theorem~\ref{thm:dta_info_opt}) --- utilisation-optimal within its
observation class --- together with $\eta_{\mathrm{info}} \to 1$
(Corollary~\ref{cor:dta_info_gap}), so DTA is information-optimal in the
full, heavy-traffic sense;
(iii)~both DTA and WS satisfy $\eta_{\mathrm{util}} \to 1$ \emph{and}
$\eta_{\mathrm{info}} \to 1$ in this same heavy-traffic limit, so the
asymptotic decomposition alone does not distinguish them; what does
distinguish them at fixed, moderate scale is the information acquisition
cost;
(iv)~$\beta_{\mathrm{WS}}/\beta_{\mathrm{DTA}} = \Omega(\log N)$
(Theorem~\ref{thm:alpha_beta});
(v)~on a 64-core dual-socket NUMA system, WS pays $\approx 9.9\times$ more
in acquisition cost per task, growing to $11\times$ at 128 cores;
(vi)~$\gamma^* \to 0$ as $W \to \infty$: DTA dominates total cost at any
positive information-cost weight for large workloads.

\clearpage
\partbreak
\part{Conclusions and Outlook}
\clearpage

\bigskip
\noindent\textbf{Overview.}
This part returns to the question posed in the Introduction
(\S\ref{sec:intro}): is restricting a scheduler to bounded, local
information a free design choice, or a real cost? We synthesise the
paper's answer, characterise the range of applicability of the developed
methods and the DTA algorithm, and identify the most promising directions
for future work. The information-theoretic cost framework developed in
Part~VI provides the analytical basis for the applicability conclusions
in \S\ref{sec:dta_scope}.

% ============================================================
\section*{Conclusions and Outlook}
\addcontentsline{toc}{section}{Conclusions and Outlook}
\label{chap:conclusions}
\setcounter{section}{0}

% ============================================================
\section{Mathematical Methods: Applicability to Distributed Systems}
\label{sec:methods_scope}
% ============================================================

This paper applies three analytical frameworks: \emph{operator-algebraic
modelling}, \emph{kinetic-theory / BBGKY expansion}, and an
\emph{information-theoretic cost framework}.  The first two originated in
mathematical physics; the third draws on classical information
theory~\cite{shannon1948,cover_thomas2006} and queueing control
theory~\cite{stidham1985,weber1990}.  We sketch the scope of each.

\subsection{Operator-algebraic modelling}

\medskip
\noindent\textbf{Core idea.}
Any distributed system in which tasks (or messages, or tokens) move
between a finite set of \emph{sites} (queues, threads, buffers, nodes)
can be encoded in an occupation-number Fock space~\cite{fock1932}.
Each atomic state transition becomes a composition of creation and
annihilation operators.  Conservation laws --- the most common
correctness requirements --- then correspond to commutation relations
between the total number operator $\hat{N}$ and the transition
operators.

\medskip
\noindent\textbf{Applicability conditions.}
The method is exact (not an approximation) whenever:
\begin{enumerate}
  \item The state space is finite and discrete (as in any real implementation).
  \item Tasks (tokens) are distinguishable and subject to a capacity constraint.
  \item Correctness is stated as a conservation or reachability property.
\end{enumerate}
It is therefore applicable, without modification, to: message-passing
middleware (RPC queues, Kafka topics as sites), actor-model runtimes
(each actor mailbox is a site), connection-pool schedulers, workflow
engines with explicit task lifecycle states, and lock-free data structure
verification.  The only requirement is that the state graph be expressible
as a finite-capacity site algebra.

\medskip
\noindent\textbf{Advantages over traditional model checking.}
Model checkers (e.g.\ TLA+, Loom~\cite{loom2023}) verify correctness by state-space
exploration, which is exponential in the number of tasks $P$.
The operator-algebraic proof is \emph{size-independent}: the commutation
identity $[\hat{N}, \hat{T}_{\alpha\to\beta}] = 0$ is proved once and holds
for all $P$.  This makes the method well-suited to systems where the task
count is not bounded a priori.

\subsection{Kinetic-theory / BBGKY hierarchy}

\medskip
\noindent\textbf{Core idea.}
In any system with $N \gg 1$ interacting agents (workers, nodes, replicas)
where the interaction is \emph{pairwise} and \emph{symmetric} (any worker
may interact with any other), the $N$-body dynamics admits a systematic
expansion in powers of $1/N$.  The zeroth-order term is a mean-field (MF)
ODE or master equation; the first-order term captures pairwise correlations
via a pair-density correction; and so on.  This is precisely the BBGKY
hierarchy~\cite{bogoliubov1946,born_green1949,kirkwood1946,yvon1935},
familiar from kinetic gas theory.

\medskip
\noindent\textbf{Transfer recipe for distributed systems.}
Given a distributed system with $N$ symmetric agents:
\begin{enumerate}
  \item \textbf{Identify the one-body density.}
        Let $f_1(q, t)$ be the marginal distribution of agent state
        (e.g.\ queue length, buffer occupancy, CPU load).
  \item \textbf{Write the exact evolution equation for $f_1$.}
        This will depend on the two-body density $f_2(q, q', t)$
        (correlation between two agents), giving BBGKY level 1.
  \item \textbf{Apply the Stosszahlansatz ($N \to \infty$).}
        Factorise $f_2 \approx f_1 \otimes f_1$.  The result is a
        closed MF equation (an ODE or PDE) that is exact as $N \to \infty$.
  \item \textbf{Compute the level-1 correction.}
        Write $f_2 = f_1 \otimes f_1 + g_2$ and derive the ODE for
        $g_2$ from the level-2 BBGKY equation, closing by setting
        $g_3 \approx 0$.  This gives an $O(1/N)$ correction to the MF.
  \item \textbf{Validate numerically.}
        Run a CTMC simulation at finite $N$ and compare with the
        MF prediction and its $O(1/N)$ correction.
\end{enumerate}

\medskip
\noindent\textbf{Applicability conditions.}
The method requires:
(i)~pairwise, symmetric interactions (no fixed master-worker hierarchy);
(ii)~a well-defined thermodynamic limit $N \to \infty$;
(iii)~ergodicity (the CTMC is irreducible under the mean-field closure).
It is directly applicable to: supermarket-model load balancers~\cite{mitzenmacher2001,vvedenskaya1996},
gossip protocols and epidemic spreading, distributed consensus
(Raft/Paxos with $N$ replicas), connection-pool rebalancing,
and any system analysed by the ``power of two choices'' framework~\cite{azar1999}.
The MF result is typically tight when $N \geq 8$; the Level-1 correction
is significant for small $N$ (4--8 agents), exactly the practically
relevant regime for multicore schedulers.

\medskip
\noindent\textbf{Extension to DAGs.}
As shown in Performance~V (\S\ref{sec:dag_kinetic}), the kinetic framework
extends naturally to heterogeneous workloads with dependency graphs:
the layer-completion fraction $f_\ell(t)$ plays the role of $f_1$, and
the fan-in $K$ controls the sharpness of the phase transition from
second-order (smooth, finite $K$) to first-order (discontinuous,
$K \to \infty$, pure fork-join).

\subsection{Information-theoretic cost framework}

\medskip
\noindent\textbf{Core idea.}
Any distributed scheduler makes decisions under \emph{limited information}
and incurs a cost to acquire that information.  Part~VI formalises this
observation by separating the total competitive ratio into two independent
factors: the \emph{information gap} $\eta_{\mathrm{info}}$ (how much the
scheduler's limited view of the system costs it, structurally) and the
\emph{utilisation ratio} $\eta_{\mathrm{util}}$ (how efficiently it uses
what it knows).  Information-acquisition cost $\beta_\pi$ is defined as a
continuous rate function over hardware observations (SPSC reads, CAS
operations), with a discrete-to-continuous bridge to actual cycle counts.

\medskip
\noindent\textbf{Applicability conditions.}
The framework applies to any pair of schedulers that differ in
(i)~the information set they observe, and
(ii)~the hardware cost of those observations.
It requires a queueing model (for $C_{\max}^*(\mathcal{I}_\pi)$) and
hardware latency measurements ($\delta_{\mathrm{intra}}$,
$\delta_{\mathrm{inter}}$, $c_{\mathrm{CAS}}^{(0)}$, $c_{\mathrm{SPSC}}$).
It is directly applicable to any work-stealing vs.\ push-based scheduler
comparison on NUMA or multi-socket hardware.

\medskip
\noindent\textbf{Key output: the $\gamma^*$ crossover.}
The framework produces a concrete \emph{crossover threshold} $\gamma^*$:
for any positive information-cost weight $\gamma > \gamma^*$, DTA's lower
acquisition cost outweighs WS's information advantage, and DTA achieves
lower total cost.  At high load ($W \to \infty$), $\gamma^* \to 0$, so
DTA dominates at any positive $\gamma$.

% ============================================================
\section{DTA Algorithm: Applicability and Use-Case Scope}
\label{sec:dta_scope}
% ============================================================

The DTA scheduler is designed for and performs best in a specific
operational envelope.  We characterise this envelope using the
theoretical results of Parts~III--V.

\subsection{Workload types}

\medskip
\noindent\textbf{Best fit: Bag-of-Tasks (BoT) at moderate-to-high load.}
The competitive ratio analysis (Part~V, \S\ref{chap:competitive_ratio}) shows
$\rho_c^{\mathrm{DTA}} \to 1$ as total work $W \to \infty$ under the
mean-field assumption.  The adversarial worst-case bound
(\S\ref{sec:cr_bot_worst}) gives $\rho_c^{\mathrm{DTA,worst}} \leq
1 + N\lfloor N/2\rfloor\delta_{\mathrm{hop}}/W$, which is also $\to 1$
for large $W$.  BoT workloads at high load are therefore the primary
design target: the deflection overhead is amortised over task computation.

\medskip
\noindent\textbf{Good fit: Parallel batch computation.}
Any workload decomposable into a fixed pool of independent tasks
--- matrix-vector products, data-parallel map-reduce, Monte Carlo
simulation, image tile processing --- maps directly to the BoT model.
The spawn-efficiency benchmarks confirm 2.6--6.2$\times$ speedup over
Tokio at spawn task counts $\geq 1$k (Figure~\ref{fig:speedup}); the
wider 2.3--6.2$\times$ range quoted elsewhere in this paper additionally
includes the deflection micro-benchmarks (Defl-1k through Defl-10M),
whose speedup ratios span 2.3--4.2$\times$.

\medskip
\noindent\textbf{Moderate fit: Fork-join DAGs with bounded fan-in.}
The DAG kinetic theory (Performance~V, Theorem~\ref{thm:fj_first})
gives $E[C_{\max}] \leq T_1/N + d \cdot H_K/\mu$, which is tight
for moderate $K$ (fan-in).  For shallow DAGs ($d \leq 10$, $K \leq 16$)
the bound is near-optimal.  DTA's SPSC-push routing is efficient for
inter-stage fan-out/fan-in.

\medskip
\noindent\textbf{Cooperative-yield workloads: outside the formal scope.}
The DTA algorithm targets spawn-and-complete tasks; it does not specify
a yield execution model.  The \dtact{} implementation uses stackful fibres,
which causes \texttt{yield\_now()} to traverse the full spawn path
(Remark~\ref{rem:confounders}).  This is an \dtact{} implementation
limitation, not an algorithmic one: a DTA implementation built on stackless
futures would not incur this overhead.  In either case, cooperative-yield
workloads lie outside the formal performance model of this paper.

\subsection{Load regime}

The stability ladder (Performance~IV) identifies three regimes:
\begin{itemize}
  \item $\varrho_0 < 1$: throughput-stable (Level~1 guarantee).
        At high load ($\varrho_0 \to 1^-$), deflection is most
        beneficial: imbalance is driven to $O(1/\sqrt{N})$ and the
        system approaches full utilisation.
  \item $\varrho_0 > 1$ transiently (burst): the warehouse absorbs
        $B^* = C_W \mu$ tasks per second before overloading.
        With $H = 114{,}688$ (production config), burst tolerance
        is 573$\times$ higher than at the benchmark $H=200$.
  \item $\varrho_0 \gg 1$ sustained: instability is inevitable for
        any finite-capacity scheduler.  The first-passage crash time
        $E[\tau_{\mathrm{crash}}]$ (Theorem~\ref{thm:crash-time})
        quantifies the safety margin.
\end{itemize}
The main lesson is that \textbf{DTA is most advantageous at high but
subcritical load} ($\varrho_0 \in [0.6, 0.95]$): deflection actively
reduces imbalance, the makespan bound is tight, and the warehouse
provides a substantial burst buffer.

\subsection{Hardware topology}

The formal model assumes a $K_4$ all-to-all SPSC topology.
The hash-routing Proposition~\ref{prop:hash_mf} shows this is exact
under uniform flow IDs.  For NUMA or CCX-topology hardware, the
mean-field result gives the topology-unaware baseline; same-socket
peer preference is an $O(1/N)$ perturbation.  In practice, for
$N \leq 8$ workers pinned to a single socket, all-to-all latency is
small ($\leq 30\,\mathrm{ns}$ SPSC write) and the topology correction
is negligible.

For larger multi-socket deployments, Part~VI (\S\ref{sec:numa_concrete})
provides a quantitative framework.  On a representative 64-core dual-socket
system, DTA's information-acquisition cost $\beta_{\mathrm{DTA}}$ is
$\approx 9.9\times$ lower than WS's $\beta_{\mathrm{WS}}$ at $N=64$,
growing as $\Omega(\log N)$.  The crossover threshold $\gamma^* \to 0$
as $W \to \infty$, confirming that DTA is the preferred choice at high
task counts on NUMA hardware.

% ============================================================
\section{Summary of Contributions}
\label{sec:contributions}
% ============================================================

We summarise the principal contributions of this paper.

\medskip
\noindent\textbf{Methodological contributions.}
\begin{enumerate}
  \item \textbf{Operator-algebraic scheduler modelling.}
        A Fock-space formalism for concurrent task schedulers that reduces
        correctness proofs to algebraic commutation identities, independent
        of task count.
  \item \textbf{BBGKY expansion for distributed schedulers.}
        First systematic application of the BBGKY kinetic hierarchy to
        a work-stealing / deflection scheduler, yielding a rigorous
        $O(1/N)$ correction to the mean-field prediction.
  \item \textbf{Kinetic-theory framework for DAG scheduling.}
        A mean-field ODE describing layer-completion fractions in
        fork-join DAGs, with two provable second-order corrections
        ($\delta^{(A)}$ Gumbel fluctuation; $\delta^{(B)}$ incomplete
        relaxation via Fill's TV-distance bound) and a directed-percolation
        phase transition identified as an open problem.
  \item \textbf{Adversarial worst-case BoT bound without mean-field.}
        A competitive-ratio upper bound whose proof is derived purely
        from work conservation and hop-count geometry and requires no
        probabilistic assumption on the arrival sequence (its stability
        precondition $\varrho_0 < 1$ is separately established via the
        mean-field analysis of Performance~IV).
  \item \textbf{Information-theoretic scheduler cost framework.}
        A formal decomposition of the competitive ratio into information
        gap and utilisation efficiency; a proof that both
        $\eta_{\mathrm{util}} \to 1$ and $\eta_{\mathrm{info}} \to 1$ for
        DTA as $W \to \infty$ (heavy-traffic information-optimality); and a
        quantitative $\Omega(\log N)$ cost advantage over WS on NUMA
        topologies, with a practitioner methodology for computing the
        DTA/WS crossover threshold $\gamma^*$ on target hardware.
\end{enumerate}

\medskip
\noindent\textbf{Analytical results for DTA.}
\begin{enumerate}
  \item Four correctness theorems: task conservation, deadlock freedom,
        livelock freedom ($\leq\lfloor N/2\rfloor$ hops), and
        linearizability of SPSC and MPMC queues.
  \item Bounded waiting time theorem with explicit, distribution-free
        closed-form ceiling $W_{\max}(N,L,C_W,K,\delta)$ and $k$-bounded
        fairness corollary.
  \item Mean-field load balance: M/M/1/$H$ self-consistency equation,
        queue-length imbalance $O(1/\sqrt{N})$ under deflection,
        propagation-of-chaos validation for $N \to \infty$.
  \item Four-level stability ladder: Foster--Lyapunov throughput stability,
        NESS convergence rate (spectral gap $\gamma \geq \mu(1-\varrho^*)$),
        adversarial burst absorption ($B^* \propto NH\mu/\lambda$),
        and crash-free first-passage threshold.
  \item Two-line makespan bound: deterministic work-conservation term
        $T_1/N + \delta_{\mathrm{hop}}\cdot\lfloor N/2\rfloor$
        plus Gumbel fluctuation correction $O(\log N/\mu)$.
  \item Empirical validation: all predictions confirmed within
        Criterion confidence intervals across seven benchmark categories
        on GitHub Actions CI.
  \item Statistical moment analysis: variance, MAD, and CV compared
        against Gumbel-predicted standard deviation; observed runner
        noise 847--433{,}000$\times$ larger than theoretical Gumbel
        fluctuation; Tokio CV plateau at large $M$ identified as an
        allocator-contention signature.
\end{enumerate}

% ============================================================
\section{Future Work and Open Problems}
\label{sec:future_work}
% ============================================================

We close with the most promising open directions, ordered from nearest
to most speculative.

\subsection{Near-term: engineering extensions}

\medskip
\noindent\textbf{Yield-path optimisation.}
The 4.7$\times$ yield regression (Remark~\ref{rem:confounders}) is an
engineering, not a theoretical, limitation.  A dedicated
\emph{local re-enqueue} path that bypasses the SPSC mailbox for
\texttt{yield\_now()} within the same worker thread would close the
gap with Tokio's ring-buffer fast path.

\medskip
\noindent\textbf{NUMA-aware peer selection.}
Extending the hash-routing Proposition~\ref{prop:hash_mf} to
non-uniform topologies: replace the uniform peer distribution with
a NUMA-distance-weighted kernel and re-derive the BBGKY Level-1
correction.  This requires solving a weighted self-consistency equation
in place of~\eqref{eq:sc-mf}.

\medskip
\noindent\textbf{Adaptive threshold $H$.}
The current model treats $H$ as a static parameter.  A dynamic
threshold (e.g.\ $H(t) = H_0 + \kappa \cdot \bar{q}(t)$) could
improve burst absorption while preserving throughput stability;
the four-level ladder would need to be re-derived with a
time-dependent Lyapunov function.

\subsection{Medium-term: theoretical extensions}

\medskip
\noindent\textbf{Level-2 BBGKY correction for load balance.}
\hyperref[chap:kinetic_theory]{Performance~II} closes the hierarchy at
level~1, whose residual error we now show (in this version) is
$O(1/N^2)$ \emph{in order} via a general propagation-of-chaos scaling
argument (Proposition~\ref{prop:residual-order}), without deriving the
triplet correlator $g_3(q,q',q'')$ explicitly. The remaining open
problem is therefore narrower than before: derive $g_3$'s explicit
functional form (not merely its order), which by the same scaling
argument would let a level-2 closure reduce the residual further, from
$O(1/N^2)$ to $O(1/N^3)$, and would additionally pin down the
$O(1/N^2)$ constant that is currently only illustrated
order-of-magnitude (\S\ref{chap:kinetic_theory}) rather than derived
in closed form.

\medskip
\noindent\textbf{Large-deviation theory for makespan tails.}
The current makespan bound (Performance~III) gives an expected value
and a sub-exponential concentration inequality.  A large-deviation
analysis~\cite{gardiner2004} of the form
$P(C_{\max} > c) \leq \exp(-\Omega(c \cdot N \mu))$ would provide
tail bounds useful for real-time scheduling guarantees.

\medskip
\noindent\textbf{Non-Poisson arrival processes.}
The mean-field analysis assumes Poisson arrivals (M/M/1/$H$).
Extensions to Markov-modulated Poisson processes (MMPP) or heavy-tailed
inter-arrival distributions would require replacing the Boltzmann
collision integral with a more general L\'evy-process generator.

\subsection{Long-term: open problems}

The two open problems identified in the DAG kinetic theory
(\S\ref{sec:dag_kinetic}) remain the most mathematically challenging:

\medskip
\noindent\textit{Open Problem~1 (DP universality for DAG scheduling).}
Establish rigorously that the scheduling phase transition in the random
DAG ensemble is in the directed-percolation (DP) universality class,
and that the makespan scaling exponents equal the DP critical exponents
$\beta \approx 0.2765$, $\nu_\parallel \approx 1.734$,
$\nu_\perp \approx 1.097$~\cite{hinrichsen2000}.
A proof would require controlling the propagation of chaos through the
DAG depth, likely using the Sznitman coupling~\cite{sznitman1991}
extended to tree-structured particle systems.

\medskip
\noindent\textit{Open Problem~2 (BBGKY Level-2 for DAG kinetics).}
Derive the level-2 BBGKY correction to the DAG ODE~\eqref{eq:dag_ode},
closing at the triplet-correlator level.  This would yield the first
fluctuation correction to the DP mean-field exponents and provide a
controlled perturbative route to the full universality-class analysis.

\medskip
\noindent\textit{Open Problem~3 (operator algebra for non-conservative systems).}
The Fock-space formalism is exact for conservative systems (tasks are
conserved).  Extending it to \emph{task creation} (e.g.\ recursive
spawning) or \emph{task cancellation} requires a generalised number
operator with a source term; the commutativity proof would need to track
net flux rather than strict conservation.

\medskip
\noindent\textit{Open Problem~4 (DTA on heterogeneous NUMA topologies).}
Modern multi-socket servers expose \emph{non-uniform memory access} (NUMA)
costs: the SPSC latency $\delta(i,j)$ between workers $i$ and $j$ depends
on their physical placement, breaking the homogeneity assumption that
underlies all mean-field results of this paper.  We lay a group-theoretic
foundation and state the central open question.

\medskip
\noindent\textbf{Topological group structure.}
Model the $N$ workers as elements of a finite abelian group
\begin{equation}\label{eq:numa_group}
  G \;=\; \mathbb{Z}_{N_s} \times \mathbb{Z}_{N_c},
\end{equation}
where $N_s$ is the number of NUMA sockets and $N_c = N/N_s$ the number of
cores per socket.  The \emph{Cayley graph} $\mathrm{Cay}(G,S)$ with generator
set $S = \{(\pm 1, 0),\, (0, \pm 1)\}$ encodes worker adjacency; it is a
discrete $2$-torus with the lattice structure familiar from condensed matter
physics.  The intra-socket subgroup
\[
  H \;=\; \{0\} \times \mathbb{Z}_{N_c} \;\leq\; G
\]
acts as a translation symmetry: any two workers in the same socket are
related by an element of $H$, and the quotient $G/H \cong \mathbb{Z}_{N_s}$
indexes the sockets.

\medskip
\noindent\textbf{NUMA latency kernel and $H$-bi-invariance.}
Define the SPSC latency function $\delta \colon G \times G \to \mathbb{R}_{+}$
by
\begin{equation}\label{eq:numa_delta}
  \delta(g,h) \;=\;
  \begin{cases}
    \delta_{\mathrm{intra}} & \text{if } gh^{-1} \in H
                              \quad (\text{same socket}), \\
    \delta_{\mathrm{inter}} & \text{otherwise}
                              \quad (\text{cross-socket}),
  \end{cases}
\end{equation}
with $\delta_{\mathrm{inter}} > \delta_{\mathrm{intra}} > 0$.
This kernel is $H$-bi-invariant: $\delta(kg, kh) = \delta(g,h)$ for all
$k \in H$, reflecting the fact that intra-socket latency is independent of
which specific cores within a socket are involved.

\medskip
\noindent\textbf{Group-indexed BBGKY density and Fourier decomposition.}
In the NUMA regime the one-body density is a \emph{group-indexed field}
$f_1^g(q,t)$ --- the marginal queue-length distribution of worker $g$.
Under the $H$-bi-invariance of~\eqref{eq:numa_delta}, $f_1^g$ depends on
$g$ only through its coset $gH \in G/H$, yielding $N_s$ coupled mean-field
equations (one per socket) in place of the single equation~\eqref{eq:sc-mf}.

To analyse inter-socket imbalance we expand in the irreducible
representations (irreps) of $G$.  Since $G$ is abelian, every irrep is a
one-dimensional character
\[
  \chi_{(k,\ell)}(s,c) \;=\; \omega_{N_s}^{ks}\,\omega_{N_c}^{\ell c},
  \qquad
  \omega_n \;=\; e^{2\pi i/n},
\]
indexed by $(k,\ell) \in \hat G \cong G$.  Define the discrete Fourier
coefficient
\begin{equation}\label{eq:numa_fourier}
  \hat{F}_{(k,\ell)}(q,t)
  \;=\;
  \frac{1}{|G|}\sum_{g \in G} f_1^g(q,t)\,\overline{\chi_{(k,\ell)}(g)}.
\end{equation}
The \emph{zero mode} $(k,\ell) = (0,0)$ recovers the socket-averaged density
$\bar f_1(q,t)$, which satisfies the standard mean-field equation derived in
\S\ref{chap:kinetic_theory}.  Non-zero modes quantify inter- and intra-socket imbalance;
their amplitudes $|\hat{F}_{(k,\ell)}|$ measure the deviation from the
homogeneous mean-field prediction.  The spectral gap
$\lambda_1 = \min_{(k,\ell)\neq(0,0)} \hat\delta(k,\ell)$
of the latency kernel (Fourier-transformed via the same formula) controls
whether imbalance modes decay or persist at stationarity.

\medskip
\noindent\textit{Open question.}
Prove (or disprove) that the stationary distribution of the group-indexed
BBGKY Level-1 system admits a Fourier-mode decoupling analogous
to~\eqref{eq:numa_fourier}, and determine the effect of the explicit
symmetry breaking $\delta_{\mathrm{inter}} > \delta_{\mathrm{intra}}$ on
the universality class identified in Open Problem~1: does the directed-percolation
exponent $\beta \approx 0.2765$ survive, or does a new universality class
emerge when the ratio $\delta_{\mathrm{inter}}/\delta_{\mathrm{intra}}$
exceeds a critical threshold?



\bigskip
\noindent\textbf{Summary.}
We set out to answer whether restricting a distributed scheduler to
bounded, local information is a free design choice or a real cost. The
answer, for the DTA scheduler studied here, is that it is close to free:
DTA is simultaneously \emph{provably correct} (four theorems), \emph{provably
live} (bounded wait, starvation freedom), \emph{quantitatively
characterised} (mean-field + BBGKY + stability ladder + makespan) at the
same leading order as schedulers with a wider information window,
\emph{competitively near-optimal} (adversarial ratio $\to 1$ at high load),
and \emph{empirically validated} against public CI benchmarks --- while
provably paying a strictly lower, asymptotically $\Omega(\log N)$ smaller
information-acquisition cost than work-stealing on NUMA hardware
(Part~VI). The kinetic-theory and operator-algebraic methods introduced
here are general-purpose tools for distributed-systems analysis and are
expected to find application well beyond the specific scheduler studied;
the information-theoretic framework of Part~VI in particular gives a
reusable methodology for deciding, on any target hardware, whether a
bounded-information scheduler like DTA or a wider-information scheduler
like work-stealing is the better choice.

\clearpage

% ============================================================
\section{Method}
\label{sec:ai-method}
% ============================================================

ACM's current policy on generative AI does not require disclosure of AI
assistance used in writing the manuscript itself; disclosure is required
for AI used \emph{in the conduct of the research}. This section records
the latter: the role of AI assistance in producing the artifacts this
paper's claims depend on --- the simulation/validation scripts and the
\dtact{} reference implementation --- together with the human review each
artifact received before being relied upon in this paper. All
mathematical derivations, theorem statements, and proofs are the
responsibility of the human author.

\medskip
\noindent\textbf{Figures and simulation scripts.}
The Monte Carlo, ODE, and finite-size simulation scripts in the
\texttt{script/} directory that generate the figures throughout this
paper (mean-field self-consistency solvers, CTMC simulators, DAG
scheduling simulations, and the FSM verifier of
\S\ref{chap:formal_model}) were developed with AI assistance and
subsequently reviewed by the author for correctness of the underlying
model against the formal definitions of Part~I and against the numerical
predictions of the corresponding theorems. Where a script's parameters
are scaled down from the production configuration for computational
tractability (e.g.\ $N$, $H$ in several Performance-chapter simulations),
this is noted in the script itself and, where relevant, in the
surrounding text.

\medskip
\noindent\textbf{\dtact{} implementation.}
The reference implementation of the DTA scheduler (\texttt{dta\_scheduler.rs}
and supporting modules in the \dtact{} repository~\cite{dtact}) that this
paper analyses was developed with AI assistance and subsequently reviewed
by the author, including a manual audit of every atomic operation's
memory-ordering annotation against the requirements established in
\S\ref{sec:mpmc}--\S\ref{sec:memorder}. Responsibility for the
correctness of the implementation, and for the accuracy of its
correspondence to the formal model, rests with the human author.

\medskip
\noindent\textbf{Errata from the verification process.}
The review described above surfaced and corrected several substantive
errors relative to earlier drafts, most notably a sign error in the
Level-1 pair-correlator analysis of \S\ref{sec:contagion-scope}'s
chapter (the true effect is overload \emph{contagion}, not
anti-correlation), a wrong-order asymptotic in the competitive-ratio
analysis of \S\ref{chap:competitive_ratio} (Gumbel $O(\ln N)$ in place
of an earlier Gaussian-tail $O(\sqrt{\ln N})$), and a symbol conflation
between the admission threshold $H$ and the hop-count bound $h_{\max}$
in Assumption~\ref{ass:threshold}. Each correction and its downstream
consequences are noted at the relevant point in the text; we record
them here together in the interest of transparency.

A subsequent, independent audit pass against the \dtact{} source
implementation surfaced nine further issues, since corrected: the
warehouse deflection trigger is $\lfloor N/2\rfloor$ failed hops, not
$N-1$, which changes the warehouse input rate
(Definition~\ref{def:warehouse-rates}) and every downstream stability and
crash-time result that depends on it; the CAS-storm overhead term
(Proposition~\ref{prop:cas_storm}) had a spurious extra factor of $(N-1)$; the
per-worker spectral-gap remark understated the mixing-time exponent by
a factor of $N$; the Foster--Lyapunov drift bound
(Theorem~\ref{thm:lyapunov}) originally substituted a mean-field
\emph{average} arrival rate where a pathwise, per-state bound was
required, and has been replaced with an unconditional pathwise proof
valid for every $\varrho_0$; Theorem~\ref{thm:bounded-wait}'s waiting-time
guarantee required an explicit net-drain assumption
(Assumption~\ref{ass:wh-net-drain}) that was previously left implicit;
the mean-field convergence step in
Theorem~\ref{thm:dta_info_opt} cited the wrong proposition and has been
re-derived via Proposition~\ref{prop:1overN} and Little's law; the
$\gamma^*$ crossover threshold (\eqref{eq:gamma_star_exact}) now carries
its $\eta_{\mathrm{util}}$ factors explicitly, with the earlier
oracle-only expression identified as its $W\to\infty$ limit; the Stage-W
overhead bound was tightened from an equality to an inequality to
account for integer ceiling effects; and the CAS retry-count bound
$\mathbb{E}[R_i] < \infty$ now has an explicit fairness-assumption proof
in place of an informal justification. As with the corrections above,
each fix and its consequences are recorded at the relevant point in the
text.

\clearpage

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