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9475510cdb
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3c90fdc774 |
@@ -20,6 +20,7 @@
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"james-yu.latex-workshop",
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"eamodio.gitlens",
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"jenselme.grammalecte",
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"jebbs.plantuml"
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],
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"settings": {
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"grammalecte.allowedExtension": ".md,.rst,.adoc,.asciidoc,.creole,.t2t,.tex",
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@@ -12,4 +12,10 @@ apt install python3 unzip wget -y
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mkdir /root/.grammalecte
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cd /root/.grammalecte
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wget https://grammalecte.net/zip/Grammalecte-fr-v2.1.1.zip
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unzip Grammalecte-fr-v2.1.1.zip
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unzip Grammalecte-fr-v2.1.1.zip
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#Installation de plantuml
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cd /tmp/
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wget https://github.com/plantuml/plantuml/releases/download/v1.2025.10/plantuml-1.2025.10.jar
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mkdir /usr/share/plantuml/
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mv plantuml-1.2025.10.jar /usr/share/plantuml/plantuml.jar
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Binary file not shown.
@@ -40,6 +40,7 @@
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\newtheorem{remark}{Remark}
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\newcommand{\RB}{\textsf{RB}\xspace}
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\newcommand{\res}{\mathsf{res}}
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\newcommand{\ARB}{\textsf{ARB}\xspace}
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\newcommand{\DL}{\textsf{DL}\xspace}
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\newcommand{\APPEND}{\textsf{APPEND}}
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@@ -52,11 +53,15 @@
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\newcommand{\RBreceived}{\textsf{RB-received}}
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\newcommand{\ordered}{\textsf{ordered}}
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\newcommand{\Winners}{\mathsf{Winners}}
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\newcommand{\Messages}{\mathsf{Messages}}
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\newcommand{\ABlisten}{\textsf{AB-listen}}
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\newcommand{\delivered}{\mathsf{delivered}}
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\newcommand{\received}{\mathsf{received}}
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\newcommand{\prop}{\mathsf{prop}}
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\newcommand{\current}{\mathsf{current}}
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\newcommand{\Seq}{\mathsf{Seq}}
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\crefname{theorem}{Theorem}{Theorems}
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\crefname{lemma}{Lemma}{Lemmas}
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@@ -86,8 +91,7 @@ We consider a static set of $n$ processes with known identities, communicating b
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\paragraph{Communication.} Processes can exchange through a Reliable Broadcast (\RB) primitive (defined below) which's invoked with the functions \RBcast$(m)$ and \RBreceived$(m)$. There exists a shared object called DenyList (\DL) (defined below) that is interfaced with the functions \APPEND$(x)$, \PROVE$(x)$ and \READ$()$.
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\paragraph{Notation.} Let $\Pi$ be the finite set of process identifiers and let $n \triangleq |\Pi|$. Two authorization subsets are $\Pi_M \subseteq \Pi$ (processes allowed to issue \APPEND) and $\Pi_V \subseteq \Pi$ (processes allowed to issue \PROVE). Indices $i,j \in \Pi$ refer to processes, and $p_i$ denotes the process with identifier $i$. Let $\mathcal{M}$ denote the universe of uniquely identifiable messages, with $m \in \mathcal{M}$. Let $\mathcal{R} \subseteq \mathbb{N}$ be the set of round identifiers; we write $r \in \mathcal{R}$ for a round. We use the precedence relation $\prec$ for the \DL{} linearization: $x \prec y$ means that operation $x$ appears strictly before $y$ in the linearized history of \DL. For any finite set $A \subseteq \mathcal{M}$, \ordered$(A)$ returns a deterministic total order over $A$ (e.g., lexicographic order on $(\textit{senderId},\textit{messageId})$ or on message hashes). For any round $r \in \mathcal{R}$, define $\Winners_r \triangleq \{\, j \in \Pi \mid (j,\PROVEtrace(r)) \prec \APPEND(r) \,\}$, i.e., the set of processes whose $\PROVE(r)$ appears before the first $\APPEND(r)$ in the \DL{} linearization.
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% For any round r ∈ R, define Winnersr ≜ { j ∈ Π | (j, prove(r)) ≺ APPEND(r) }. Pas bien on compare des tuples et des operations
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We denoted by $\PROVE^{(j)}(r)$ or $\APPEND^{(j)}(r)$ the operation $\PROVE(r)$ or $\APPEND(r)$ invoked by process $j$.
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% ------------------------------------------------------------------------------
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\section{Primitives}
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@@ -101,21 +105,28 @@ We consider a static set of $n$ processes with known identities, communicating b
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\end{itemize}
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\subsection{DenyList (DL)}
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The \DL is a \emph{shared, append-only} object that records attestations about opaque application-level tokens. It exposes the following operations:
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\begin{itemize}[leftmargin=*]
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\item \APPEND$(x)$
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\item \PROVE$(x)$: issue an attestation for token $x$; this operation is \emph{valid} (return true) only if no \APPEND$(x)$ occurs earlier in the \DL linearization. Otherwise, it is invalid (return false).
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\item \READ$()$: return a (permutation of the) valid operations observed so far; subsequent reads are monotone (contain supersets of previously observed valid operations).
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We assume a synchronous DenyList (\DL) object with the following properties.
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The DenyList object type supports three operations: $\APPEND$, $\PROVE$, and $\READ$. These operations appear as if executed in a sequence $\Seq$ such that:
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\begin{itemize}
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\item \textbf{Termination.} A $\PROVE$, an $\APPEND$, or a $\READ$ operation invoked by a correct process always returns.
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\item \textbf{APPEND Validity.} The invocation of $\APPEND(x)$ by a process $p$ is valid if:
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\begin{itemize}
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\item $p \in \Pi_M \subseteq \Pi$; \textbf{and}
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\item $x \in S$, where $S$ is a predefined set.
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\end{itemize}
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Otherwise, the operation is invalid.
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\item \textbf{PROVE Validity.} If the invocation of a $op = \PROVE(x)$ by a correct process $p$ is not valid, then:
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\begin{itemize}
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\item $p \not\in \Pi_V \subseteq \Pi$; \textbf{or}
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\item A valid $\APPEND(x)$ appears before $op$ in $\Seq$.
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\end{itemize}
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Otherwise, the operation is valid.
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\item \textbf{PROVE Anti-Flickering.} If the invocation of a operation $op = \PROVE(x)$ by a correct process $p \in \Pi_V$ is invalid, then any $\PROVE(x)$ operation that appears after $op$ in $\Seq$ is invalid.
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\item \textbf{READ Validity.} The invocation of $op = \READ()$ by a process $p \in \pi_V$ returns the list of valid invocations of $\PROVE$ that appears before $op$ in $\Seq$ along with the names of the processes that invoked each operation.
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\item \textbf{Anonymity.} Let us assume the process $p$ invokes a $\PROVE(v)$ operation. If the process $p'$ invokes a $\READ()$ operation, then $p'$ cannot learn the value $v$ unless $p$ leaks additional information.
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\end{itemize}
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\paragraph{Validity.} \APPEND$(x)$ is valid iff the issuer is authorized (in $\Pi_M$) and $x$ belongs to the application-defined domain $S$. \PROVE$(x)$ is valid iff the issuer is authorized (in $\Pi_V$) and there is no earlier \APPEND$(x)$ in the \DL linearization.
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\paragraph{Progress.} If a correct process invokes \APPEND$(x)$, then eventually all correct processes will be unable to issue a valid \PROVE$(x)$, and \READ{} at all correct processes will (eventually) reflect that \APPEND$(x)$ has been recorded.
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\paragraph{Termination.} Every operation invoked by a correct process eventually returns.
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\paragraph{Interface and Semantics.} The \DL provides a single global linearization of operations consistent with each process's program order. \READ{} is prefix-monotone; concurrent updates become visible to all correct processes within bounded time (by synchrony). Duplicate requests may be idempotently coalesced by the implementation.
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% ------------------------------------------------------------------------------
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\section{Target Abstraction: Atomic Reliable Broadcast (ARB)}
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Processes export \ABbroadcast$(m)$ and \ABdeliver$(m)$. \ARB requires total order:
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@@ -134,10 +145,6 @@ plus Integrity/No-duplicates/Validity (inherited from \RB and the construction).
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Equivalently, there exists a time after which every $\PROVE(r)$ is invalid in $H$.
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\end{definition}
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\begin{definition}[First APPEND]\label{def:first-append}
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Given a \DL{} linearization $H$, for any closed round $r\in\mathcal{R}$, we denote by $\APPEND^\star(r)$ the earliest $\APPEND(r)$ in $H$.
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\end{definition}
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\subsection{Variables}
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Each process $p_i$ maintains:
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@@ -146,6 +153,7 @@ Each process $p_i$ maintains:
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\State $\received \gets \emptyset$ \Comment{Messages received via \RB but not yet delivered}
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\State $\delivered \gets \emptyset$ \Comment{Messages already delivered}
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\State $\prop[r][j] \gets \bot,\ \forall r,j$ \Comment{Proposal from process $j$ for round $r$}
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\State $\current \gets 0$
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\end{algorithmic}
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\paragraph{DenyList.} The \DL is initialized empty. We assume $\Pi_M = \Pi_V = \Pi$ (all processes can invoke \APPEND and \PROVE).
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@@ -166,7 +174,6 @@ Each process $p_i$ maintains:
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% \paragraph{} An \ABbroadcast$(m)$ chooses the next open round from the \DL view, proposes all pending messages together with the new $m$, disseminates this proposal via \RB, then executes $\PROVE(r)$ followed by $\APPEND(r)$ to freeze the winners of the round. The loop polls \DL until (i) some winner’s proposal includes $m$ in a \emph{closed} round and (ii) all winners' proposals for closed rounds are known locally, ensuring eventual inclusion and delivery.
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% Partie avec le max-1 pas ouf; essayer de faire incr la boucle autrement
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\renewcommand{\algletter}{B}
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\begin{algorithm}[H]
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\caption{\ABbroadcast$(m)$ (at process $p_i$)}\label{alg:ab-bcast}
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@@ -189,48 +196,34 @@ Each process $p_i$ maintains:
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\end{algorithmic}
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\end{algorithm}
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% \paragraph{} TODO
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\renewcommand{\algletter}{C}
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\begin{algorithm}[H]
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\caption{\ABdeliver() at process $p_i$}\label{alg:delivery}
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\begin{algorithmic}[1]
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%local variables
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\State $next \gets 0$ \Comment{Next round to deliver}
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\State $to\_deliver \gets \emptyset$ \Comment{Queue of messages ready to be delivered}
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\vspace{1em}
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\Function{ABdeliver}{}
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% \State \textbf{on} \ABdeliver() \textbf{do} \Comment{Called when the process wants to receive the next message}
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\If{$to\_deliver = \emptyset$} \Comment{If no message is ready to deliver, try to fetch the next round}
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\State $P \leftarrow \READ()$ \Comment{Fetch latest \DL state to learn recent $\PROVE$ operations}
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\If{$\forall j : (j, \PROVEtrace(next)) \not\in P$} \Comment{Check if some process proved round $next$}
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\State \Return $\bot$ \Comment{Round $next$ is still open}
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\EndIf
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\State $\APPEND(next)$; $P \leftarrow \READ()$ \Comment{Close round $next$ if not already closed}
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\State $W_{next} \leftarrow \{ j : (j, \PROVEtrace(next)) \in P \}$ \Comment{Compute winners of round $next$}
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\If{$\exists j \in W_{next},\ \prop[next][j] = \bot$} \Comment{Check if we have all winners' proposals}
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\State \Return $\bot$ \Comment{Some winner's proposal for round $next$ is still missing}
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\EndIf
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\State $M_{next} \leftarrow \bigcup_{j \in W_{next}} \prop[next][j]$ \Comment{Compute the agreed proposal for round $next$}
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\For{\textbf{each}\ $m \in \ordered(M_{next})$} \Comment{Enqueue messages in deterministic order}
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\If{$m \notin \delivered$}
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\State $to\_deliver.push(m)$ \Comment{Append $m$ to the delivery queue}
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\EndIf
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\EndFor
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\State $next \leftarrow next + 1$ \Comment{Advance to the next round}
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\State $r \gets \current$
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\State $P \gets \READ()$
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\If{$\forall j : (j, \PROVEtrace(r)) \not\in P$}
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\State \Return $\bot$
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\EndIf
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\State $m \leftarrow \text{to\_deliver.pop()}$
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% \State $to\_deliver \leftarrow to\_deliver \setminus \{m\}$
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\State $\APPEND(r)$; $P \gets \READ()$
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\State $W_r \gets \{ j : (j, \PROVEtrace(r)) \in P \}$
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\If{$\exists j \in W_r,\ \prop[r][j] = \bot$}
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\State \Return $\bot$
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\EndIf
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\State $M_r \gets \bigcup_{j \in W_r} \prop[r][j]$
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\State $m \gets \ordered(M_r \setminus \delivered)[0]$ \Comment{Set $m$ as the smaller message not already delivered}
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\State $\delivered \leftarrow \delivered \cup \{m\}$
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\If{$M_r \setminus \delivered = \emptyset$} \Comment{Check if all messages from round $r$ have been delivered}
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\State $\current \leftarrow \current + 1$
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\EndIf
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\State \textbf{return} $m$
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\EndFunction
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\end{algorithmic}
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\end{algorithm}
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% ------------------------------------------------------------------------------
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\section{Correctness}
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%attention au usage de "unique"
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%definition de APPEND* ssi r closed
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\begin{lemma}[Stable round closure]\label{lem:closure-stable}
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If a round $r$ is closed, then there exists a linearization point $t_0$ of $\APPEND(r)$ in the \DL, and from that point on, no $\PROVE(r)$ can be valid.
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@@ -238,13 +231,17 @@ Once closed, a round never becomes open again.
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\end{lemma}
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\begin{proof}
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By definition of closed round, some $\APPEND(r)$ occurs in the linearization $H$. \\
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$H$ is a total order of operations, the set of $\APPEND(r)$ operations is totally ordered, and hence there exists a smallest $\APPEND(r)$ in $H$. We denote this operation $\APPEND^\star(r)$ and $t_0$ its linearization point. \\
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By the validity property of \DL, a $\PROVE(r)$ is valid iff $\PROVE(r) \prec \APPEND^\star(r)$. Thus, after $t_0$, no $\PROVE(r)$ can be valid. \\
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By \Cref{def:closed-round}, some $\APPEND(r)$ occurs in the linearization $H$. \\
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$H$ is a total order of operations, the set of $\APPEND(r)$ operations is totally ordered, and hence there exists a smallest $\APPEND(r)$ in $H$. We denote this operation $\APPEND^{(\star)}(r)$ and $t_0$ its linearization point. \\
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By the validity property of \DL, a $\PROVE(r)$ is valid iff $\PROVE(r) \prec \APPEND^{(\star)}(r)$. Thus, after $t_0$, no $\PROVE(r)$ can be valid. \\
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$H$ is a immutable grow-only history, and hence once closed, a round never becomes open again. \\
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Hence there exists a linearization point $t_0$ of $\APPEND(r)$ in the \DL, and from that point on, no $\PROVE(r)$ can be valid and the closure is stable.
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\end{proof}
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\begin{definition}[First APPEND]\label{def:first-append}
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Given a \DL{} linearization $H$, for any closed round $r\in\mathcal{R}$, we denote by $\APPEND^{(\star)}(r)$ the earliest $\APPEND(r)$ in $H$.
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\end{definition}
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\begin{lemma}[Across rounds]\label{lem:across}
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If there exists a $r$ such that $r$ is closed, $\forall r'$ such that $r' < r$, r' is also closed.
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\end{lemma}
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@@ -255,30 +252,34 @@ If there exists a $r$ such that $r$ is closed, $\forall r'$ such that $r' < r$,
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\emph{Induction.} Assume the lemma is true for round $k\geq 0$, we prove it for round $k+1$.
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\smallskip
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Assume $k+1$ is closed and let $\APPEND^\star(k+1)$ be the earliest $\APPEND(k+1)$ in the DL linearization $H$.
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Assume $k+1$ is closed and let $\APPEND^{(\star)}(k+1)$ be the earliest $\APPEND(k+1)$ in the DL linearization $H$.
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By Lemma 1, after an $\APPEND(k)$ is in $H$, any later $\PROVE(k)$ is rejected also, a process’s program order is preserved in $H$.
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There are two possibilities for where $\APPEND^\star(k+1)$ is invoked.
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\paragraph{Case (B6).}
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Some process $p^\star$ executes the loop (lines B5–B11) and invokes $\APPEND^\star(k+1)$ at line B6.
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Immediately before line B6, line B5 sets $r\leftarrow r+1$, so the previous loop iteration (if any) targeted $k$. We consider two sub-cases.
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\emph{(i) $p^\star$ is not in its first loop iteration.}
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In the previous iteration, $p^\star$ executed $\PROVE^\star(k)$ at B6, followed (in program order) by $\APPEND^\star(k)$.
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If round $k$ wasn't closed when $p^\star$ execute $\PROVE^\star(k)$ at B9, then the condition at B8 would be true hence the tuple $(p^\star, \PROVEtrace(k))$ should be visible in P which implies that $p^\star$ should leave the loop at round $k$, contradicting the assumption that $p^\star$ is now executing another iteration.
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Since the tuple is not visible, the $\PROVE^\star(k)$ was rejected by the DL which implies by definition an $\APPEND(k)$ already exists in $H$. Thus in this case $k$ is closed.
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\emph{(ii) $p^\star$ is in its first loop iteration.}
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To compute the value $r_{max}$, $p^\star$ must have observed one or many $(\_ , \PROVEtrace(k+1))$ in $P$ at B2/B3, issued by some processes (possibly different from $p^\star$). Let's call $p_1$ the issuer of the first $\PROVE(k+1)$ in the linearization $H$. \\
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When $p_1$ executed $P \gets \READ()$ at B2 and compute $r_{max}$ at B3, he observed no tuple $(j,\PROVEtrace(k+1))$ in $P$ because he's the issuer of the first one. So when $p_1$ executed the loop (B5–B11), he run it for the round $k$, didn't seen any $(1,\PROVEtrace(k))$ in $P$ at B8, and then executed the first $\PROVE(k+1)$ at B6 in a second iteration. \\
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If round $k$ wasn't closed when $p_1$ execute $\PROVE(k)$ at B6, then the condition at B8 should be true which implies that $p_1$ sould leave the loop at round $k$, contradicting the assumption that $p_1$ is now executing $\PROVE(r+1)$. In this case $k$ is closed.
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\paragraph{Case (C9).}
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Some process invokes $\APPEND(k+1)$ at C9.
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Line C9 is guarded by the presence of $\PROVE(\textit{next})$ in $P$ with $\textit{next}=k+1$ (C6).
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Moreover, the local pointer $\textit{next}$ grow by increment of 1 and only advances after finishing the current round (C20), so if a process can reach $\textit{next}=k+1$ it implies that he has completed round $k$, which includes closing $k$ at C9 when $\PROVE(k)$ is observed.
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Hence $\APPEND^\star(k+1)$ implies a prior $\APPEND(k)$ in $H$, so $k$ is closed.
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There are two possibilities for where $\APPEND^{(\star)}(k+1)$ is invoked.
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\begin{itemize}
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\item \textbf{Case (B6) :}
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Some process $p^\star$ executes the loop (lines B5–B11) and invokes $\APPEND^{(\star)}(k+1)$ at line B6.
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Immediately before line B6, line B5 sets $r\leftarrow r+1$, so the previous loop iteration (if any) targeted $k$. We consider two sub-cases.
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\begin{itemize}
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\item \emph{(i) $p^\star$ is not in its first loop iteration.}
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In the previous iteration, $p^\star$ executed $\PROVE^{(\star)}(k)$ at B6, followed (in program order) by $\APPEND^{(\star)}(k)$.
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If round $k$ wasn't closed when $p^\star$ execute $\PROVE^{(\star)}(k)$ at B9, then the condition at B8 would be true hence the tuple $(p^\star, \PROVEtrace(k))$ should be visible in P which implies that $p^\star$ should leave the loop at round $k$, contradicting the assumption that $p^\star$ is now executing another iteration.
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Since the tuple is not visible, the $\PROVE^{(\star)}(k)$ was rejected by the DL which implies by definition an $\APPEND(k)$ already exists in $H$. Thus in this case $k$ is closed.
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\item \emph{(ii) $p^\star$ is in its first loop iteration.}
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To compute the value $r_{max}$, $p^\star$ must have observed one or many $(\_ , \PROVEtrace(k+1))$ in $P$ at B2/B3, issued by some processes (possibly different from $p^\star$). Let's call $p_1$ the issuer of the first $\PROVE^{(1)}(k+1)$ in the linearization $H$. \\
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When $p_1$ executed $P \gets \READ()$ at B2 and compute $r_{max}$ at B3, he observed no tuple $(\_,\PROVEtrace(k+1))$ in $P$ because he's the issuer of the first one. So when $p_1$ executed the loop (B5–B11), he run it for the round $k$, didn't seen any $(1,\PROVEtrace(k))$ in $P$ at B8, and then executed the first $\PROVE^{(1)}(k+1)$ at B6 in a second iteration. \\
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If round $k$ wasn't closed when $p_1$ execute $\PROVE^{(1)}(k)$ at B6, then the condition at B8 should be true which implies that $p_1$ sould leave the loop at round $k$, contradicting the assumption that $p_1$ is now executing $\PROVE^{(1)}(r+1)$. In this case $k$ is closed.
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\end{itemize}
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\item \textbf{Case (C8) :}
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Some process invokes $\APPEND(k+1)$ at C8.
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Line C8 is guarded by the presence of $\PROVE(\textit{next})$ in $P$ with $\textit{next}=k+1$ (C5).
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Moreover, the local pointer $\textit{next}$ grow by increment of 1 and only advances after finishing the current round (C17), so if a process can reach $\textit{next}=k+1$ it implies that he has completed round $k$, which includes closing $k$ at C8 when $\PROVE(k)$ is observed.
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Hence $\APPEND^\star(k+1)$ implies a prior $\APPEND(k)$ in $H$, so $k$ is closed.
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\end{itemize}
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\smallskip
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In all cases, $k+1$ closed implie $k$ closed. By induction on $k$, if the lemme is true for a closed $k$ then it is true for a closed $k+1$.
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@@ -286,11 +287,11 @@ If there exists a $r$ such that $r$ is closed, $\forall r'$ such that $r' < r$,
|
||||
\end{proof}
|
||||
|
||||
\begin{definition}[Winner Invariant]\label{def:winner-invariant}
|
||||
Let take a closed round $r$ (which implies by \Cref{def:closed-round} that some $\APPEND(r)$ occurs in the DL linearization $H$), there exist a unique and defined set of winners such as
|
||||
For any closed round $r$, define
|
||||
\[
|
||||
\Winners_r = \{ j : (j,\PROVEtrace(r)) \prec \APPEND^\star(r) \}
|
||||
\Winners_r \triangleq \{ j : \PROVE^{(j)}(r) \prec \APPEND^\star(r) \}
|
||||
\]
|
||||
with $\APPEND^\star(r)$ being the earliest \APPEND$(r)$ in the \DL linearization.
|
||||
as the unique set of winners of round $r$.
|
||||
\end{definition}
|
||||
|
||||
\begin{lemma}[Invariant view of closure]\label{lem:closure-view}
|
||||
@@ -298,7 +299,7 @@ If there exists a $r$ such that $r$ is closed, $\forall r'$ such that $r' < r$,
|
||||
\end{lemma}
|
||||
|
||||
\begin{proof}
|
||||
Fix a round $r$ such that some $\APPEND(r)$ occurs; hence $r$ is \emph{closed}. By \Cref{lem:closure-stable}, there exists a unique earliest $\APPEND(r)$ in the DL linearization, denoted $\APPEND^\star(r)$.
|
||||
Let's take a closed round $r$. By \Cref{def:first-append}, there exists a unique earliest $\APPEND(r)$ in the DL linearization, denoted $\APPEND^\star(r)$.
|
||||
|
||||
Consider any correct process $p$ that invokes $\READ()$ after $\APPEND^\star(r)$ in the DL linearization. Since $\APPEND^\star(r)$ invalidates all subsequent $\PROVE(r)$, the set of valid tuples $(\_,\PROVEtrace(r))$ observed by any correct process after $\APPEND^\star(r)$ is fixed and identical across all correct processes.
|
||||
|
||||
@@ -306,181 +307,178 @@ If there exists a $r$ such that $r$ is closed, $\forall r'$ such that $r' < r$,
|
||||
\end{proof}
|
||||
|
||||
\begin{lemma}[Well-defined winners]\label{lem:winners}
|
||||
In any execution, when a process computes $W_{r}$, $r$ is closed.
|
||||
For any correct process and round $r$, if the process computes $W_r$ at line C9, then :
|
||||
\begin{itemize}
|
||||
\item $\Winners_r$ is defined;
|
||||
\item the computed $W_r$ is exactly $\Winners_r$.
|
||||
\end{itemize}
|
||||
\end{lemma}
|
||||
|
||||
\begin{proof}
|
||||
Fix a process $p$ that reaches line C10 (computing $W_{\textit{next}}$). Immediately before C10, line C9 executes
|
||||
Let take a correct process $p_i$ that reach line C9 to compute $W_r$. \\
|
||||
By program order, $p_i$ must have executed $\APPEND^{(i)}(r)$ at C8 before, which implies by \Cref{def:closed-round} that round $r$ is closed. So by \Cref{def:winner-invariant}, $\Winners_r$ is defined. \\
|
||||
By \Cref{lem:closure-view}, all correct processes eventually observe the same set of valid tuples $(\_,\PROVEtrace(r))$ in their \DL view. Hence, when $p_i$ executes the $\READ()$ at C8 after the $\APPEND^{(i)}(r)$, it observes a set $P$ that includes all valid tuples $(\_,\PROVEtrace(r))$ such that
|
||||
\[
|
||||
\APPEND(\textit{next});\; P \gets \READ()
|
||||
W_r = \{ j : (j,\PROVEtrace(r)) \in P \} = \{j : \PROVE^{(j)}(r) \prec \APPEND^\star(r) \} = \Winners_r
|
||||
\]
|
||||
|
||||
Line C9 is guarded by the condition at line C6, which
|
||||
\end{proof}
|
||||
|
||||
\begin{lemma}[View-Invariant Winners]\label{lem:view-invariant}
|
||||
For any closed round $r$, there exists a unique set
|
||||
\[
|
||||
\Winners_r = \{ j : (j,\PROVEtrace(r)) \prec \APPEND^\star(r) \}
|
||||
\]
|
||||
with $\APPEND^\star(r)$ being the earliest \APPEND$(r)$ in the \DL linearization.\\
|
||||
Such that for any correct process $p$ that computes $W_r$, $W_r = \Winners_r$.
|
||||
\end{lemma}
|
||||
|
||||
\begin{proof}[Proof]
|
||||
Fix a round $r$ such that some $\APPEND(r)$ occurs; hence $r$ is \emph{closed}. By \Cref{lem:closure-stable}, there exists a unique earliest $\APPEND(r)$ in the DL linearization, denoted $\APPEND^\star(r)$.
|
||||
|
||||
Consider any correct process $p$ that computes $W_r$ at line C10. By \Cref{lem:winners}, $r$ is closed when $p$ computes $W_r$. By \Cref{lem:closure-stable}, the linearization point of $\APPEND^\star(r)$ precedes that of the subsequent $\READ()$ at C9, therefore the $P$ returned by this $\READ()$ is posterior to $\APPEND^\star(r)$ in the DL linearization.
|
||||
|
||||
Hence, at the moment C10 computes
|
||||
\[
|
||||
W_r \;\leftarrow\; \{\, j : (j,\PROVEtrace(r)) \in P \,\},
|
||||
\]
|
||||
the view $P$ is posterior to $\APPEND^\star(r)$, so
|
||||
\[
|
||||
W_r = \{\, j : (j,\PROVEtrace(r)) \prec \APPEND^\star(r) \,\} = \Winners_r.
|
||||
\]
|
||||
|
||||
Since this argument holds for any correct process computing $W_r$, all correct processes compute the same winner set $\Winners_r$ for closed round $r$.
|
||||
\end{proof}
|
||||
|
||||
\begin{lemma}[No APPEND without PROVE]\label{lem:append-prove}
|
||||
If some correct process invokes $\APPEND(r)$, then some correct process must have previously invoked $\PROVE(r)$.
|
||||
If some process invokes $\APPEND(r)$, then at least a process must have previously invoked $\PROVE(r)$.
|
||||
\end{lemma}
|
||||
|
||||
\begin{proof}[Proof]
|
||||
Consider the round $r$ such that some correct process invokes $\APPEND(r)$. There are two possible cases
|
||||
Consider the round $r$ such that some process invokes $\APPEND(r)$. There are two possible cases
|
||||
|
||||
\paragraph{Case (B6).}
|
||||
There exists a process $p^\star$ who's the issuer of the earliest $\APPEND^\star(r)$ in the DL linearization $H$. By program order, $p^\star$ must have previously invoked $\PROVE(r)$ before invoking $\APPEND^\star(r)$ at B6. In this case, there is at least one $\PROVE(r)$ valid in $H$ issued by a correct process before $\APPEND^\star(r)$.
|
||||
|
||||
\paragraph{Case (C9).}
|
||||
There exist a process $p^\star$ invokes $\APPEND^\star(r)$ at C9. Line C9 is guarded by the condition at C6, which ensures that $p$ observed some $(\_,\PROVEtrace(r))$ in $P$. In this case, there is at least one $\PROVE(r)$ valid in $H$ issued by some process before $\APPEND^\star(r)$.
|
||||
|
||||
In both cases, if some correct process invokes $\APPEND(r)$, then some correct process must have previously invoked $\PROVE(r)$.
|
||||
\begin{itemize}
|
||||
\item \textbf{Case (B6) :}
|
||||
There exists a process $p^\star$ who's the issuer of the earliest $\APPEND^{(\star)}(r)$ in the DL linearization $H$. By program order, $p^\star$ must have previously invoked $\PROVE^{(\star)}(r)$ before invoking $\APPEND^{(\star)}(r)$ at B6. In this case, there is at least one $\PROVE(r)$ valid in $H$ issued by a correct process before $\APPEND^{(\star)}(r)$.
|
||||
|
||||
\item \textbf{Case (C8) :}
|
||||
There exist a process $p^\star$ invokes $\APPEND^{(\star)}(r)$ at C8. Line C8 is guarded by the condition at C5, which ensures that $p$ observed some $(\_,\PROVEtrace(r))$ in $P$. In this case, there is at least one $\PROVE(r)$ valid in $H$ issued by some process before $\APPEND^{(\star)}(r)$.
|
||||
\end{itemize}
|
||||
|
||||
In both cases, if some process invokes $\APPEND(r)$, then some process must have previously invoked $\PROVE(r)$.
|
||||
\end{proof}
|
||||
|
||||
\begin{lemma}[No empty winners]\label{lem:nonempty}
|
||||
For any correct process that computes $W_r$, $W_r \neq \emptyset$.
|
||||
Let $r$ be a round, if $\Winners_r$ is defined, then $\Winners_r \neq \emptyset$.
|
||||
\end{lemma}
|
||||
|
||||
\begin{proof}[Proof]
|
||||
By \Cref{lem:winners}, any correct process that computes $W_r$ implies that round $r$ is closed. By \Cref{lem:closure-stable}, there exists a unique earliest $\APPEND(r)$ in the DL linearization, denoted $\APPEND^\star(r)$.
|
||||
If $\Winners_r$ is defined, then by \Cref{def:winner-invariant}, round $r$ is closed. By \Cref{def:closed-round}, some $\APPEND(r)$ occurs in the DL linearization. \\
|
||||
By \Cref{lem:append-prove}, at least a process must have invoked a valid $\PROVE(r)$ before $\APPEND^{(\star)}(r)$. Hence, there exists at least one $j$ such that $\{j: \PROVE^{(j)}(r) \prec \APPEND^\star(r)\}$, so $\Winners_r \neq \emptyset$.
|
||||
\end{proof}
|
||||
|
||||
By \Cref{lem:view-invariant}, any correct process computing $W_r$ obtains
|
||||
\begin{lemma}[Winners must propose]\label{lem:winners-propose}
|
||||
For any closed round $r$, $\forall j \in \Winners_r$, process $j$ must have invoked a $\RBcast(S^{(j)}, r, j)$
|
||||
\end{lemma}
|
||||
|
||||
\begin{proof}[Proof]
|
||||
Fix a closed round $r$. By \Cref{def:winner-invariant}, for any $j \in \Winners_r$, there exist a valid $\PROVE^{(j)}(r)$ such that $\PROVE^{(j)}(r) \prec \APPEND^\star(r)$ in the DL linearization. By program order, if $j$ invoked a valid $\PROVE^{(j)}(r)$ at line B6 he must have invoked $\RBcast(S^{(j)}, r, j)$ directly before.
|
||||
\end{proof}
|
||||
|
||||
\begin{definition}[Messages invariant]\label{def:messages-invariant}
|
||||
For any closed round $r$ and any correct process $p_i$ such that $\nexists j \in \Winners_r : prop^{[i)}[r][j] = \bot$, define
|
||||
\[
|
||||
W_r = \Winners_r = \{\, j : (j,\PROVEtrace(r)) \prec \APPEND^\star(r) \,\}.
|
||||
\Messages_r \triangleq \bigcup_{j\in\Winners_r} \prop^{(i)}[r][j]
|
||||
\]
|
||||
as the unique set of messages proposed by the winners of round $r$.
|
||||
\end{definition}
|
||||
|
||||
So
|
||||
\[
|
||||
W_r \neq \emptyset \implies \exists j : (j,\PROVEtrace(r)) \prec \APPEND^\star(r).
|
||||
\]
|
||||
\begin{lemma}[Non-empty winners proposal]\label{lem:winner-propose-nonbot}
|
||||
For any closed round $r$, $\forall j \in \Winners_r$, for any correct process $p_i$, eventually $\prop^{(i)}[r][j] \neq \bot$.
|
||||
\end{lemma}
|
||||
|
||||
Consider any correct process $p$ that computes $W_k$ at line C10. We show that $k$ is closed when $p$ executed the $\READ$ operation at C9 by \Cref{lem:winners}. This implies that some process must have invoked $\APPEND(k)$, which by \Cref{lem:append-prove} implies that some correct process must have previously invoked $\PROVE(k)$. Hence there exists at least one $j$ such that $(j,\PROVEtrace(k)) \prec \APPEND^\star(k)$, so $W_k \neq \emptyset$.
|
||||
\begin{proof}[Proof]
|
||||
Fix a closed round $r$. By \Cref{def:winner-invariant}, for any $j \in \Winners_r$, there exist a valid $\PROVE^{(j)}(r)$ such that $\PROVE^{(j)}(r) \prec \APPEND^\star(r)$ in the DL linearization. By \Cref{lem:winners-propose}, $j$ must have invoked $\RBcast(S^{(j)}, r, j)$.
|
||||
|
||||
Let take a process $p_i$, by \RB \emph{Validity}, every correct process eventually receives $j$'s \RB message for round $r$, which sets $\prop[r][j]$ to a non-$\bot$ value at line A3.
|
||||
\end{proof}
|
||||
|
||||
\begin{lemma}[Eventual proposal closure]\label{lem:eventual-closure}
|
||||
For any correct process $p_i$ that computes $M_r$ for round $r$, there exist no $j$ such that $j \in \Winners_r$ and $\prop^{(i)}[r][j] = \bot$.
|
||||
If a correct process $p_i$ define $M_r$ at line C13, then for every $j \in \Winners_r$, $\prop^{(i)}[r][j] \neq \bot$.
|
||||
\end{lemma}
|
||||
|
||||
\begin{proof}[Proof]
|
||||
Fix a correct process $p_i$ that computes $M_r$ at line C14. By \Cref{lem:winners}, $r$ is closed when $p_i$ executed the $\READ$ operation at C9. By \Cref{lem:view-invariant}, $p_i$ computes the unique winner set $\Winners_r$.
|
||||
Let take a correct process $p_i$ that computes $M_r$ at line C13. By \Cref{lem:winners}, $p_i$ computes the unique winner set $\Winners_r$.
|
||||
|
||||
By \Cref{lem:nonempty}, $\Winners_r \neq \emptyset$. Consider any $j \in \Winners_r$. Since $j$ is a winner, by definition of $\Winners_r$, $j$ must have invoked a valid $\PROVE(r)$ before the earliest $\APPEND(r)$ in the DL linearization. Since $j$ is correct and by program order, if he invoked a valid $\PROVE(r)$ he must have invoked $\RBcast(S^{(j)}, r, j)$ directly before. By \RB \emph{Validity}, every correct process eventually receives $j$'s \RB message for round $r$, which sets $\prop[r][j]$ to a non-$\bot$ value (A3). The instruction C14 where $p_i$ computes $M_r$ is guarded by the condition at C11, which ensures that $p_i$ has received all \RB messages from every winner $j \in \Winners_r$. Hence, when $p_i$ computes $M_r = \bigcup_{j\in\Winners_r} \prop[r][j]$, we have $\prop[r][j] \neq \bot$ for all $j \in \Winners_r$.
|
||||
By \Cref{lem:nonempty}, $\Winners_r \neq \emptyset$. The instruction at line C13 where $p_i$ computes $M_r$ is guarded by the condition at C10, which ensures that $p_i$ has received all \RB messages from every winner $j \in \Winners_r$. Hence, when $p_i$ computes $M_r = \bigcup_{j\in\Winners_r} \prop^{(i)}[r][j]$, we have $\prop^{(i)}[r][j] \neq \bot$ for all $j \in \Winners_r$.
|
||||
\end{proof}
|
||||
|
||||
\begin{lemma}[Unique proposal per sender per round]\label{lem:unique-proposal}
|
||||
For any round $r$ and any process $j$, $j$ invokes at most one $\RBcast(S, r, j)$.
|
||||
For any round $r$ and any process $p_i$, $p_i$ invokes at most one $\RBcast(S, r, i)$.
|
||||
\end{lemma}
|
||||
|
||||
\begin{proof}[Proof]
|
||||
By program order, any process $j$ invokes $\RBcast(S, r, j)$ at line B6 must be in the loop (B5–B11). No matter the number of iterations of the loop, line B5 always uses the current value of $r$ which is incremented by 1 at line B5. Hence, in any execution, any process $j$ invokes $\RBcast(S, r, j)$ at most once for any round $r$.
|
||||
By program order, any process $p_i$ invokes $\RBcast(S, r, i)$ at line B6 must be in the loop B5–B11. No matter the number of iterations of the loop, line B5 always uses the current value of $r$ which is incremented by 1 at each turn. Hence, in any execution, any process $p_i$ invokes $\RBcast(S, r, j)$ at most once for any round $r$.
|
||||
\end{proof}
|
||||
|
||||
\begin{lemma}[Proposal convergence]\label{lem:convergence}
|
||||
For any closed round $r$, after all \RB messages from $\Winners_r$ are received, every correct process computes the same $M_r = \bigcup_{j\in \Winners_r} \prop[r][j]$.
|
||||
For any round $r$, for any correct processes $p_i$ that define $M_r$ at line C13, we have
|
||||
\[
|
||||
M_r^{(i)} = \Messages_r
|
||||
\]
|
||||
\end{lemma}
|
||||
|
||||
\begin{proof}[Proof]
|
||||
Fix a closed round $r$. By \Cref{lem:view-invariant}, any correct process computing $W_r$ obtains the same winner set $\Winners_r$. By \Cref{lem:eventual-closure}, every correct process that computes $M_r$ has received all \RB messages from every winner $j \in \Winners_r$, so $\prop[r][j]$ is non-$\bot$ for all $j \in \Winners_r$. By \Cref{lem:unique-proposal}, each winner $j$ invokes at most one $\RBcast(S^{(j)}, r, j)$, so $\prop[r][j] = S^{(j)}$ is uniquely defined. Hence, every correct process computes the same
|
||||
Let take a correct process $p_i$ that define $M_r$ at line C13. That implies that $p_i$ has defined $W_r$ at line C9. It implies that, by \Cref{lem:winners}, $r$ is closed and $W_r = \Winners_r$. \\
|
||||
By \Cref{lem:eventual-closure}, for every $j \in \Winners_r$, $\prop^{(i)}[r][j] \neq \bot$. By \Cref{lem:unique-proposal}, each winner $j$ invokes at most one $\RBcast(S^{(j)}, r, j)$, so $\prop^{(i)}[r][j] = S^{(j)}$ is uniquely defined. Hence, when $p_i$ computes
|
||||
\[
|
||||
M_r = \bigcup_{j\in\Winners_r} \prop[r][j] = \bigcup_{j\in\Winners_r} S^{(j)}.
|
||||
M_r^{(i)} = \bigcup_{j\in\Winners_r} \prop^{(i)}[r][j] = \bigcup_{j\in\Winners_r} S^{(j)} = \Messages_r.
|
||||
\]
|
||||
\end{proof}
|
||||
|
||||
\begin{lemma}[Inclusion]\label{lem:inclusion}
|
||||
If some correct process invokes $\ABbroadcast(m)$, then there exist a round $r$ and a process $j\in\Winners_r$ such that $p_j$ invokes
|
||||
\[
|
||||
\RBcast(S, r, j)\quad\text{with}\quad m\in S.
|
||||
\]
|
||||
\end{lemma}
|
||||
|
||||
\begin{proof}
|
||||
Fix a correct process $p_i$ that invokes $\ABbroadcast(m)$ and eventually exits the loop (B5–B11) at some round $r$. There are two possible cases.
|
||||
|
||||
\begin{itemize}
|
||||
\item \textbf{Case 1:} $p_i$ exits the loop because $(i, \PROVEtrace(r)) \in P$. In this case, by \Cref{def:winner-invariant}, $p_i$ is a winner in round $r$. By program order, $p_i$ must have invoked $\RBcast(S, r, i)$ at B6 before invoking $\PROVE^{(i)}(r)$ at B7. By line B4, $m \in S$. Hence there exist a closed round $r$ and a correct process $j=i\in\Winners_r$ such that $j$ invokes $\RBcast(S, r, j)$ with $m\in S$.
|
||||
|
||||
\item \textbf{Case 2:} $p_i$ exits the loop because $\exists j, r': (j, \PROVEtrace(r')) \in P \wedge m \in \prop[r'][j]$. In this case, by \Cref{lem:winners-propose} and \Cref{lem:unique-proposal} $j$ must have invoked a unique $\RBcast(S, r', j)$. Which set $\prop^{(i)}[r'][j] = S$ with $m \in S$.
|
||||
\end{itemize}
|
||||
|
||||
In both cases, if some correct process invokes $\ABbroadcast(m)$, then there exist a round $r$ and a correct process $j\in\Winners_r$ such that $j$ invokes
|
||||
\[
|
||||
\RBcast(S, r, j)\quad\text{with}\quad m\in S.
|
||||
\]
|
||||
\end{proof}
|
||||
|
||||
\begin{lemma}[Broadcast Termination]\label{lem:bcast-termination}
|
||||
If a correct process $p$ invokes $\ABbroadcast(m)$, then $p$ eventually quit the function and returns.
|
||||
If a correct process invokes $\ABbroadcast(m)$, then he eventually exit the function and returns.
|
||||
\end{lemma}
|
||||
|
||||
\begin{proof}[Proof]
|
||||
Fix a correct process $p$ that invokes $\ABbroadcast(m)$. The lemma is true iff $(i, \PROVEtrace(r)) \in P$ or $\exists j, r': (j, \PROVEtrace(r')) \in P \wedge m \in \prop[r'][j]$ (like guarded at B8).
|
||||
Let a correct process $p_i$ that invokes $\ABbroadcast(m)$. The lemma is true if $\exists r_1$ such that $r_1 \geq r_{max}$ and if $(i, \PROVEtrace(r_1)) \in P$; or if $\exists r_2$ such that $r_2 \geq r_{max}$ and if $\exists j: (j, \PROVEtrace(r_2)) \in P \wedge m \in \prop[r_2][j]$ (like guarded at B8).
|
||||
|
||||
\begin{itemize}
|
||||
\item \textbf{Case 1:} there exists a round $r$ such that $p$ invokes $\PROVE(r)$ and this $\PROVE(r)$ is valid. Hence eventually $(i, \PROVEtrace(r)) \in P$ and $p$ exits the loop at B8.
|
||||
\item \textbf{Case 2:} there exists no round $r$ such that $p$ invokes $\PROVE(r)$ and this $\PROVE(r)$ is valid. In this case $p$ invokes infinitely many $\RBcast(S, r, i)$ at B6 with $m \in S$ (line B4).\\
|
||||
The assumption that no $\PROVE(r)$ invoked by $p$ is valid implies by \Cref{lem:nonempty} that for every possible round $r$ there at least one winner. Because there is an infinite number of rounds, and a finite number of processes, there exists at least one correct process $j$ that invokes infinitely many valid $\PROVE(r)$ and by extension infinitely many $\ABbroadcast(\_)$. By \RB \emph{Validity}, $j$ eventually receives $p$ 's \RB messages. Let call $t$ the time when $j$ receives $p$ 's \RB message \\
|
||||
For the first invocation of $\ABbroadcast(m)$ by $j$ after $t$, $j$ must include $m$ in his proposal $S$ at B4 for round $r'$ (because $m$ is pending in $j$ 's $\received \setminus \delivered$). Because $j \in \Winners_{r'}$ and by \RB \emph{Validity}, every correct process eventually receives $j$ 's \RB message for round $r'$, including $p$. When $p$ receives $j$ 's \RB message, $\prop[r'][j]$ is set to a non-$\bot$ value (A3) which includes $m$. Hence eventually $\exists j, r': (j, \PROVEtrace(r')) \in P \wedge m \in \prop[r'][j]$ and $p$ exits the loop at B8.
|
||||
\end{itemize}
|
||||
The first case explicit the case where $p$ is a winner and also covers the condition $(i, \PROVEtrace(r)) \in P$. And in the second case, we show that if the first condition is never satisfied, the second one will eventually be satisfied. Hence either the first or the second condition will eventually be satisfied, and $p$ eventually exits the loop and returns from $\ABbroadcast(m)$.
|
||||
Let admit that there exists no round $r_1$ such that $p_i$ invokes a valid $\PROVE(r_1)$. In this case $p_i$ invokes infinitely many $\RBcast(S, \_, i)$ at B6 with $m \in S$ (line B4).\\
|
||||
The assumption that no $\PROVE(r_1)$ invoked by $p$ is valid implies by \DL \emph{Validity} that for every round $r' \geq r_{max}$, there exists at least one $\APPEND(r')$ in the DL linearization, hence by \Cref{lem:nonempty} for every possible round $r'$ there at least a winner. Because there is an infinite number of rounds, and a finite number of processes, there exists at least a correct process $p_k$ that invokes infinitely many valid $\PROVE(r')$ and by extension infinitely many $\ABbroadcast(\_)$. By \RB \emph{Validity}, $p_k$ eventually receives $p_i$ 's \RB messages. Let call $t_0$ the time when $p_k$ receives $p_i$ 's \RB message. \\
|
||||
At $t_0$, $p_k$ execute \Cref{alg:rb-handler} and do $\received \leftarrow \received \cup \{S\}$ with $m \in S$ (line A2).
|
||||
For the first invocation of $\ABbroadcast(\_)$ by $p_k$ after the execution of \Cref{alg:rb-handler}, $p_k$ must include $m$ in his proposal $S$ at line B4 (because $m$ is pending in $j$'s $\received \setminus \delivered$ set). There exists a minimum round $r_2$ such that $p_k \in \Winners_{r_2}$ after $t_0$. By \Cref{lem:winner-propose-nonbot}, eventually $\prop^{(i)}[r_2][k] \neq \bot$. By \Cref{lem:unique-proposal}, $\prop^{(i)}[r_2][k]$ is uniquely defined as the set $S$ proposed by $p_k$ at B6, which by \Cref{lem:inclusion} includes $m$. Hence eventually $m \in \prop^{(i)}[r_2][k]$ and $k \in \Winners_{r_2}$.
|
||||
|
||||
So if $p_i$ is a winner he cover the condition $(i, \PROVEtrace(r_1)) \in P$. And we show that if the first condition is never satisfied, the second one will eventually be satisfied. Hence either the first or the second condition will eventually be satisfied, and $p_i$ eventually exits the loop and returns from $\ABbroadcast(m)$.
|
||||
\end{proof}
|
||||
|
||||
|
||||
\begin{lemma}[Inclusion]\label{lem:inclusion}
|
||||
If some correct process invokes $\ABbroadcast(m)$, then there exist a (closed) round $r$ and a corrcet process $j\in\Winners_r$ such that $j$ invokes
|
||||
\[
|
||||
\RBcast(S, r, j)\quad\text{with}\quad m\in S.
|
||||
\]
|
||||
\end{lemma}
|
||||
|
||||
\begin{proof}
|
||||
Fix a correct process $p$ that invokes $\ABbroadcast(m)$. By \Cref{lem:bcast-termination}, $p$ eventually exits the loop (B5–B11) at some round $r$. There are two possible cases.
|
||||
|
||||
\paragraph{Case 1:} $p$ exits the loop because $(i, \PROVEtrace(r)) \in P$. In this case, $p$ is a winner in round $r$ by definition of $\Winners_r$. By program order, $p$ must have invoked $\RBcast(S, r, i)$ at B6 before invoking $\PROVE(r)$ at B7. By line B4, $m \in S$. Hence there exist a closed round $r$ and a correct process $j=i\in\Winners_r$ such that $j$ invokes $\RBcast(S, r, j)$ with $m\in S$.
|
||||
|
||||
\paragraph{Case 2:} $p$ exits the loop because $\exists j, r': (j, \PROVEtrace(r')) \in P \wedge m \in \prop[r'][j]$. In this case, by \Cref{lem:view-invariant}, any correct process computing $W_{r'}$ obtains the unique winner set $\Winners_{r'}$, so $j\in\Winners_{r'}$. By program order, $j$ must have invoked $\RBcast(S, r', j)$ at B6 before invoking $\PROVE(r')$ at B7. By line B4, $m \in S$. Hence there exist a closed round $r'=r$ and a correct process $j\in\Winners_{r'}$ such that $j$ invokes $\RBcast(S, r', j)$ with $m\in S$.
|
||||
|
||||
In both cases, if some correct process invokes $\ABbroadcast(m)$, then there exist a (closed) round $r$ and a correct process $j\in\Winners_r$ such that $j$ invokes
|
||||
\[
|
||||
\RBcast(S, r, j)\quad\text{with}\quad m\in S.
|
||||
\]
|
||||
\end{proof}
|
||||
|
||||
|
||||
% \begin{proof}[Proof]
|
||||
% 2 cas, p_i est winner et cest gagné.
|
||||
% Si p_i n'est pas winner on fait par l'absurde comme quoi personne ne propose m. Et on montre que c'est impossible.
|
||||
|
||||
% S est modif a chaque iteration de la boucle. On montre que y a un plus petit S bloqué
|
||||
|
||||
% Pour sortir de la boucle on doit avoir A ou B vrai, et il faut montrer que non A implique B
|
||||
% \end{proof}
|
||||
|
||||
\begin{lemma}[Validity]\label{lem:validity}
|
||||
If a correct process $p$ invokes $\ABbroadcast(m)$, then every correct process who eventually invokes $\ABdeliver()$ delivers $m$.
|
||||
If a correct process $p$ invokes $\ABbroadcast(m)$, then every correct process that invokes a infinitely often times $\ABdeliver()$ eventually delivers $m$.
|
||||
\end{lemma}
|
||||
|
||||
\begin{proof}[Proof]
|
||||
Fix a correct process $p$ that invokes $\ABbroadcast(m)$. By \Cref{lem:inclusion}, there exist a closed round $r$ and a correct process $j\in\Winners_r$ such that $j$ invokes
|
||||
Let $p_i$ a correct process that invokes $\ABbroadcast(m)$ and $p_q$ a correct process that infinitely invokes $\ABdeliver()$. By \Cref{lem:inclusion}, there exist a closed round $r$ and a correct process $j\in\Winners_r$ such that $p_j$ invokes
|
||||
\[
|
||||
\RBcast(S, r, j)\quad\text{with}\quad m\in S.
|
||||
\]
|
||||
|
||||
Consider any correct process $q$ that eventually invokes $\ABdeliver()$. By \Cref{lem:eventual-closure}, when $q$ computes $M_r$ at line C14, $\prop[r][j]$ is non-$\bot$ because $j \in \Winners_r$. By \Cref{lem:unique-proposal}, $j$ invokes at most one $\RBcast(S, r, j)$, so $\prop[r][j] = S$ is uniquely defined. Hence, when $q$ computes
|
||||
By \Cref{lem:eventual-closure}, when $p_q$ computes $M_r$ at line C13, $\prop[r][j]$ is non-$\bot$ because $j \in \Winners_r$. By \Cref{lem:unique-proposal}, $p_j$ invokes at most one $\RBcast(S, r, j)$, so $\prop[r][j]$ is uniquely defined. Hence, when $p_q$ computes
|
||||
\[
|
||||
M_r = \bigcup_{k\in\Winners_r} \prop[r][k],
|
||||
\]
|
||||
we have $m \in \prop[r][j] = S$, so $m \in M_r$. By line C16–C18, $m$ is enqueued into $to\_deliver$ unless it has already been delivered. Therefore, when $q$ eventually invokes $\ABdeliver()$, it eventually delivers $m$.
|
||||
we have $m \in \prop[r][j] = S$, so $m \in M_r$. By \Cref{lem:convergence}, $M_r$ is invariant so each computation of $M_r$ by any correct process that defines it includes $m$. At each invocation of $\ABdeliver()$ which deliver $m'$, $m'$ is add to $\delivered$ until $M_r \subseteq \delivered$. Once this append we're assured that there exist an invocation of $\ABdeliver()$ which return $m$. Hence $m$ is well delivered.
|
||||
\end{proof}
|
||||
|
||||
\begin{lemma}[Total Order]\label{lem:total-order}
|
||||
For any two messages $m_1$ and $m_2$ broadcast by correct processes such that $m_1$ is broadcast before $m_2$, any correct process $p_j$ that delivers both $m_1$ and $m_2$ delivers $m_1$ before $m_2$.
|
||||
\begin{lemma}[No duplication]\label{lem:no-duplication}
|
||||
No correct process delivers the same message more than once.
|
||||
\end{lemma}
|
||||
|
||||
\begin{proof}[Proof]
|
||||
Fix two messages $m_1$ and $m_2$ broadcast by correct processes. Consider any correct process $p_i$ that delivers $m_1$ before $m_2$. By program order, $p_i$ must have invoked $\ABbroadcast(m_1)$ before $\ABbroadcast(m_2)$. Let $r_1$ and $r_2$ be the rounds at which $p_i$ exits the loop (B5–B11) when invoking $\ABbroadcast(m_1)$ and $\ABbroadcast(m_2)$ respectively. By program order, $r_1 < r_2$.
|
||||
\begin{proof}
|
||||
Let consider two invokations of $\ABdeliver()$ made by the same correct process which returns $m$. Let call these two invocations respectively $\ABdeliver^{(A)}()$ and $\ABdeliver^{(B)}()$.
|
||||
|
||||
Consider any correct process $p_j$ that delivers both $m_1$ and $m_2$. By \Cref{lem:validity}, there exist closed rounds $r'_1$ and $r'_2$ and correct processes $k_1 \in \Winners_{r'_1}$ and $k_2 \in \Winners_{r'_2}$ such that
|
||||
When $\ABdeliver^{(A)}()$ occur, by program order and because it reach line C19 to return $m$, the process must have add $m$ to $\delivered$. Hence when $\ABdeliver^{(B)}()$ reach line C14 to extract the next message to deliver, it can't be $m$ because $m \not\in (M_r \setminus \{..., m, ...\})$. So a $\ABdeliver^{(B)}()$ which deliver $m$ can't occur.
|
||||
\end{proof}
|
||||
|
||||
\begin{lemma}[Total order]\label{lem:total-order}
|
||||
For any two messages $m_1$ and $m_2$ delivered by correct processes, if a correct process $p_i$ delivers $m_1$ before $m_2$, then any correct process $p_j$ that delivers both $m_1$ and $m_2$ delivers $m_1$ before $m_2$.
|
||||
\end{lemma}
|
||||
|
||||
\begin{proof}
|
||||
Consider any correct process that delivers both $m_1$ and $m_2$. By \Cref{lem:validity}, there exist closed rounds $r'_1$ and $r'_2$ and correct processes $k_1 \in \Winners_{r'_1}$ and $k_2 \in \Winners_{r'_2}$ such that
|
||||
\[
|
||||
\RBcast(S_1, r'_1, k_1)\quad\text{with}\quad m_1\in S_1,
|
||||
\]
|
||||
@@ -488,16 +486,138 @@ If some correct process invokes $\ABbroadcast(m)$, then there exist a (closed) r
|
||||
\RBcast(S_2, r'_2, k_2)\quad\text{with}\quad m_2\in S_2.
|
||||
\]
|
||||
|
||||
By program order, $p_i$ must have waited for the loop (B5–B11) to exit at round $r_1$ before invoking $\ABbroadcast(m_2)$. Hence, $r_1 \leq r'_1 < r_2 \leq r'_2$, so $r'_1 < r'_2$. Because $p_j$ delivers messages in round order (C5–C20), $p_j$ delivers all messages from round $r'_1$ (including $m_1$) before delivering any message from round $r'_2$ (including $m_2$). Therefore, $p_j$ delivers $m_1$ before $m_2$.
|
||||
Let consider three cases :
|
||||
\begin{itemize}
|
||||
\item \textbf{Case 1:} $r_1 < r_2$. By program order, any correct process must have waited to append in $\delivered$ every messages in $M_{r_1}$ (which contains $m_1$) to increment $\current$ and eventually set $\current = r_2$ to compute $M_{r_2}$ and then invoke the $\ABdeliver()$ that returns $m_2$. Hence, for any correct process that delivers both $m_1$ and $m_2$, it delivers $m_1$ before $m_2$.
|
||||
|
||||
\item \textbf{Case 2:} $r_1 = r_2$. By \Cref{lem:convergence}, any correct process that computes $M_{r_1}$ at line C13 computes the same set of messages $\Messages_{r_1}$. By line C14 the messages are pull in a deterministic order defined by $\ordered(\_)$. Hence, for any correct process that delivers both $m_1$ and $m_2$, it delivers $m_1$ and $m_2$ in the deterministic order defined by $\ordered(\_)$.
|
||||
\end{itemize}
|
||||
|
||||
In all possible cases, any correct process that delivers both $m_1$ and $m_2$ delivers $m_1$ and $m_2$ in the same order.
|
||||
\end{proof}
|
||||
|
||||
\begin{theorem}[\ARB]
|
||||
Under the assumed \DL synchrony and \RB reliability, the algorithm implements Atomic Reliable Broadcast.
|
||||
\begin{lemma}[Fifo Order]\label{lem:fifo-order}
|
||||
For any two messages $m_1$ and $m_2$ broadcast by the same correct process $p_i$, if $p_i$ invokes $\ABbroadcast(m_1)$ before $\ABbroadcast(m_2)$, then any correct process $p_j$ that delivers both $m_1$ and $m_2$ delivers $m_1$ before $m_2$.
|
||||
\end{lemma}
|
||||
|
||||
\begin{proof}
|
||||
Let take two messages $m_1$ and $m_2$ broadcast by the same correct process $p_i$, with $p_i$ invoking $\ABbroadcast(m_1)$ before $\ABbroadcast(m_2)$. By \Cref{lem:validity}, there exist closed rounds $r_1$ and $r_2$ and correct processes $k_1 \in \Winners_{r_1}$ and $k_2 \in \Winners_{r_2}$ such that
|
||||
\[
|
||||
\RBcast(S_1, r_1, k_1)\quad\text{with}\quad m_1\in S_1,
|
||||
\]
|
||||
\[
|
||||
\RBcast(S_2, r_2, k_2)\quad\text{with}\quad m_2\in S_2.
|
||||
\]
|
||||
|
||||
By program order, $p_i$ must have invoked $\RBcast(S_1, r_1, i)$ before $\RBcast(S_2, r_2, i)$. By \Cref{lem:unique-proposal}, any process invokes at most one $\RBcast(S, r, i)$ per round, hence $r_1 < r_2$. By \Cref{lem:total-order}, any correct process that delivers both $m_1$ and $m_2$ delivers them in a deterministic order.
|
||||
|
||||
In all possible cases, any correct process that delivers both $m_1$ and $m_2$ delivers $m_1$ before $m_2$.
|
||||
\end{proof}
|
||||
|
||||
\begin{theorem}[FIFO-\ARB]
|
||||
Under the assumed \DL synchrony and \RB reliability, the algorithm implements FIFO Atomic Reliable Broadcast.
|
||||
\end{theorem}
|
||||
|
||||
\begin{proof}
|
||||
We show that the algorithm satisfies the properties of FIFO Atomic Reliable Broadcast under the assumed \DL synchrony and \RB reliability.
|
||||
|
||||
First, by \Cref{lem:bcast-termination}, if a correct process invokes \ABbroadcast$(m)$, then it eventually returns from this invocation.
|
||||
Moreover, \Cref{lem:validity} states that if a correct process invokes \ABbroadcast$(m)$, then every correct process that invokes \ABdeliver() infinitely often eventually delivers $m$.
|
||||
This gives the usual Validity property of \ARB.
|
||||
|
||||
Concerning Integrity and No-duplicates, the construction only ever delivers messages that have been obtained from the underlying \RB primitive.
|
||||
By the Integrity property of \RB, every such message was previously \RBcast by some process, so no spurious messages are delivered.
|
||||
In addition, \Cref{lem:no-duplication} states that no correct process delivers the same message more than once.
|
||||
Together, these arguments yield the Integrity and No-duplicates properties required by \ARB.
|
||||
|
||||
For the ordering guarantees, \Cref{lem:total-order} shows that for any two messages $m_1$ and $m_2$ delivered by correct processes, every correct process that delivers both $m_1$ and $m_2$ delivers them in the same order.
|
||||
Hence all correct processes share a common total order on delivered messages.
|
||||
Furthermore, \Cref{lem:fifo-order} states that for any two messages $m_1$ and $m_2$ broadcast by the same correct process, any correct process that delivers both messages delivers $m_1$ before $m_2$ whenever $m_1$ was broadcast before $m_2$.
|
||||
Thus the global total order extends the per-sender FIFO order of \ABbroadcast.
|
||||
|
||||
All the above lemmas are proved under the assumptions that \DL satisfies the required synchrony properties and that the underlying primitive is a Reliable Broadcast (\RB) with Integrity, No-duplicates and Validity.
|
||||
Therefore, under these assumptions, the algorithm satisfies Validity, Integrity/No-duplicates, total order and per-sender FIFO order, and hence implements FIFO Atomic Reliable Broadcast, as claimed.
|
||||
\end{proof}
|
||||
|
||||
\section{Reciprocity}
|
||||
% ------------------------------------------------------------------------------
|
||||
|
||||
So far, we assumed the existence of a synchronous DenyList (\DL) object and
|
||||
showed how to upgrade a Reliable Broadcast (\RB) primitive into FIFO Atomic
|
||||
Reliable Broadcast (\ARB). We now briefly argue that, conversely, an \ARB{}
|
||||
primitive is strong enough to implement a synchronous \DL object (ignoring the
|
||||
anonymity property).
|
||||
|
||||
\paragraph{DenyList as a deterministic state machine.}
|
||||
Without anonymity, the \DL specification defines a
|
||||
deterministic abstract object: given a sequence $\Seq$ of operations
|
||||
$\APPEND(x)$, $\PROVE(x)$, and $\READ()$, the resulting sequence of return
|
||||
values and the evolution of the abstract state (set of appended elements,
|
||||
history of operations) are uniquely determined.
|
||||
|
||||
\paragraph{State machine replication over \ARB.}
|
||||
Assume a system that exports a FIFO-\ARB primitive with the guarantees that if a correct process invokes $\ABbroadcast(m)$, then every correct process eventually $\ABdeliver(m)$ and the invocation eventually returns.
|
||||
Following the classical \emph{state machine replication} approach
|
||||
such as described in Schneider~\cite{Schneider90}, we can implement a fault-tolerant service by ensuring the following properties:
|
||||
\begin{quote}
|
||||
\textbf{Agreement.} Every nonfaulty state machine replica receives every request. \\
|
||||
\textbf{Order.} Every nonfaulty state machine replica processes the requests it receives in
|
||||
the same relative order.
|
||||
\end{quote}
|
||||
|
||||
Which are cover by our FIFO-\ARB specification.
|
||||
|
||||
\paragraph{Correctness.}
|
||||
|
||||
|
||||
\begin{theorem}[From \ARB to synchronous \DL]\label{thm:arb-to-dl}
|
||||
In an asynchronous message-passing system with crash failures, assume a
|
||||
FIFO Atomic Reliable Broadcast primitive with Integrity, No-duplicates,
|
||||
Validity, and the liveness of $\ABbroadcast$. Then, ignoring anonymity, there
|
||||
exists an implementation of a synchronous DenyList object that satisfies the
|
||||
Termination, Validity, and Anti-flickering properties.
|
||||
\end{theorem}
|
||||
|
||||
\begin{proof}
|
||||
Because the \DL object is deterministic, all correct processes see the same
|
||||
sequence of operations and compute the same sequence of states and return
|
||||
values. We obtain:
|
||||
|
||||
\begin{itemize}[leftmargin=*]
|
||||
\item \textbf{Termination.} The liveness of \ARB ensures that each
|
||||
$\ABbroadcast$ invocation by a correct process eventually returns, and
|
||||
the corresponding operation is eventually delivered and applied at all
|
||||
correct processes. Thus every $\APPEND$, $\PROVE$, and $\READ$ operation invoked by a correct process
|
||||
eventually returns.
|
||||
\item \textbf{APPEND/PROVE/READ Validity.} The local code that forms
|
||||
\ABbroadcast requests can achieve the same preconditions as in the
|
||||
abstract \DL specification (e.g., $p\in\Pi_M$, $x\in S$ for
|
||||
$\APPEND(x)$). Once an operation is delivered, its effect and return
|
||||
value are exactly those of the sequential \DL specification applied in
|
||||
the common order.
|
||||
\item \textbf{PROVE Anti-Flickering.} In the sequential \DL specification,
|
||||
once an element $x$ has been appended, all subsequent $\PROVE(x)$ are
|
||||
invalid forever. Since all replicas apply operations in the same order,
|
||||
this property holds in every execution of the replicated implementation:
|
||||
after the first linearization point of $\APPEND(x)$, no later
|
||||
$\PROVE(x)$ can return ``valid'' at any correct process.
|
||||
\end{itemize}
|
||||
|
||||
Formally, we can describe the \DL object with the state machine approach for
|
||||
crash-fault, asynchronous message-passing systems with a total order broadcast
|
||||
layer~\cite{Schneider90}.
|
||||
\end{proof}
|
||||
|
||||
|
||||
|
||||
\bibliographystyle{plain}
|
||||
\begin{thebibliography}{9}
|
||||
% (left intentionally blank)
|
||||
\bibitem{Schneider90}
|
||||
Fred B.~Schneider.
|
||||
\newblock Implementing fault-tolerant services using the state machine
|
||||
approach: a tutorial.
|
||||
\newblock {\em ACM Computing Surveys}, 22(4):299--319, 1990.
|
||||
\end{thebibliography}
|
||||
|
||||
\end{document}
|
||||
|
||||
40
Recherche/BFT-ARBover/.latexmkrc
Normal file
40
Recherche/BFT-ARBover/.latexmkrc
Normal file
@@ -0,0 +1,40 @@
|
||||
$pdf_mode = 1; # latexmk -pdf par défaut
|
||||
$pdflatex = 'pdflatex -interaction=nonstopmode -synctex=1 %O %S';
|
||||
|
||||
# --- Config PlantUML ----------------------------------------------------
|
||||
# Si plantuml est dans le PATH :
|
||||
# $plantuml = 'plantuml';
|
||||
|
||||
# Si tu utilises un JAR :
|
||||
$plantuml = 'java -jar -Djava.awt.headless=true /usr/share/plantuml/plantuml.jar';
|
||||
|
||||
# Options PlantUML : sortie LaTeX/TikZ
|
||||
$plantuml_opts = '-tlatex:nopreamble';
|
||||
|
||||
# --- Dépendance personnalisée .puml -> .tex -----------------------------
|
||||
# Quand latexmk a besoin de "truc.tex" et que "truc.puml" existe,
|
||||
# il appelle la fonction puml2tex pour le générer.
|
||||
|
||||
add_cus_dep( 'puml', 'tex', 0, 'puml2tex' );
|
||||
|
||||
sub puml2tex {
|
||||
my ($base_name) = @_; # base du fichier cible, sans extension
|
||||
|
||||
# Exemple : $base_name = 'diagrams/login'
|
||||
my $puml = "$base_name.puml";
|
||||
my $tex = "$base_name.tex";
|
||||
|
||||
# Message dans le log latexmk
|
||||
print "PlantUML: génération de $tex à partir de $puml\n";
|
||||
|
||||
# Commande PlantUML
|
||||
my $cmd = "$plantuml $plantuml_opts $puml ";
|
||||
my $ret = system($cmd);
|
||||
|
||||
# 0 = succès, 1 = erreur pour latexmk
|
||||
return $ret ? 1 : 0;
|
||||
}
|
||||
|
||||
# --- Confort ------------------------------------------------------------
|
||||
# Compilation continue (latexmk -pvc)
|
||||
$preview_continuous = 1;
|
||||
54
Recherche/BFT-ARBover/diagrams/BFT_behaviour.puml
Normal file
54
Recherche/BFT-ARBover/diagrams/BFT_behaviour.puml
Normal file
@@ -0,0 +1,54 @@
|
||||
@startuml
|
||||
!pragma teoz true
|
||||
|
||||
database DL
|
||||
actor P1
|
||||
actor P2
|
||||
P1 -> DL : <latex>READ()</latex>
|
||||
DL --> P1 : <latex>P</latex>
|
||||
P1 -> P1 : <latex>r_{max} = max\{r : (\_, prove(r)) \in P\}</latex>
|
||||
|
||||
loop <latex>\textbf{foreach } r \in \{r_{max} + 1, \dots\}</latex>
|
||||
|
||||
' P1 ->(05) P2 : <latex>RBcast(prop, S, r, 1)</latex>
|
||||
|
||||
P1 -> DL : <latex>PROVE(r)</latex>
|
||||
P1 -> DL : <latex>APPEND(r)</latex>
|
||||
|
||||
P1 -> DL : <latex>READ()</latex>
|
||||
DL --> P1 : <latex>P</latex>
|
||||
|
||||
alt <latex>(1, \text{prove(}r\text{)}) \in P</latex>
|
||||
note over P1 : break
|
||||
end
|
||||
end
|
||||
|
||||
P2 -> P2 : <latex>ABdeliver()</latex>
|
||||
P2 -> DL : <latex>READ()</latex>
|
||||
DL --> P2 : <latex>P</latex>
|
||||
note over P2
|
||||
line(C4)
|
||||
process P2 check locally if
|
||||
<latex>\forall j : (j, prove(r)) \not\in P</latex>
|
||||
which is false since P1 correctly
|
||||
PROVE(r) and APPEND(r)
|
||||
|
||||
<latex>\text{P1 is next include in } W_r</latex>
|
||||
end note
|
||||
|
||||
P2 -> DL : <latex>APPEND(r)</latex>
|
||||
P2 -> DL : <latex>READ()</latex>
|
||||
DL --> P2 : <latex>P</latex>
|
||||
|
||||
note over P2
|
||||
line(C9)
|
||||
process P2 check locally if
|
||||
<latex>\forall j \in W_r : prop[r][j] = \bot</latex>
|
||||
which can't be false since P1 didn't
|
||||
execute <latex>RBcast(prop, S, r, 1)</latex>
|
||||
|
||||
P2 will never progress and
|
||||
deliver any futur messages
|
||||
end note
|
||||
hide footbox
|
||||
@enduml
|
||||
113
Recherche/BFT-ARBover/diagrams/BFT_behaviour.tex
Normal file
113
Recherche/BFT-ARBover/diagrams/BFT_behaviour.tex
Normal file
@@ -0,0 +1,113 @@
|
||||
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|
||||
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|
||||
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|
||||
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|
||||
\node at (268.4883pt,65pt)[below right,color=black,inner sep=0]{P2};
|
||||
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|
||||
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|
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|
||||
\draw[pstyle5] (29pt,99pt) -- (108.3255pt,99pt);
|
||||
\node at (41pt,87pt)[below right,inner sep=0]{$READ()$};
|
||||
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|
||||
\draw[pstyle6] (24pt,119.8333pt) -- (103.3255pt,119.8333pt);
|
||||
\node at (31pt,111pt)[below right,inner sep=0]{$P$};
|
||||
\draw[pstyle5] (109.3255pt,143.8333pt) -- (151.3255pt,143.8333pt);
|
||||
\draw[pstyle5] (151.3255pt,143.8333pt) -- (151.3255pt,156.8333pt);
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|
||||
\node at (116.3255pt,131.8333pt)[below right,inner sep=0]{$r_{max} = max\{r : (\_, prove(r)) \in P\}$};
|
||||
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|
||||
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||||
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|
||||
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|
||||
\node at (92.59pt,170.8333pt)[below right,inner sep=0]{$\textbf{foreach } r \in \{r_{max} + 1, \dots\}$};
|
||||
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|
||||
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||||
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|
||||
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|
||||
\draw[pstyle4] (35pt,226.8333pt) -- (25pt,230.8333pt) -- (35pt,234.8333pt) -- (31pt,230.8333pt) -- cycle;
|
||||
\draw[pstyle5] (29pt,230.8333pt) -- (108.3255pt,230.8333pt);
|
||||
\node at (41pt,218.8333pt)[below right,inner sep=0]{$APPEND(r)$};
|
||||
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||||
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|
||||
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|
||||
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|
||||
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||||
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||||
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||||
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|
||||
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|
||||
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|
||||
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|
||||
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|
||||
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|
||||
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||||
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|
||||
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|
||||
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|
||||
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|
||||
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|
||||
\node at (284.3933pt,371.6666pt)[below right,inner sep=0]{$ABdeliver()$};
|
||||
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|
||||
\draw[pstyle5] (29pt,420.6666pt) -- (276.3933pt,420.6666pt);
|
||||
\node at (41pt,408.6666pt)[below right,inner sep=0]{$READ()$};
|
||||
\draw[pstyle4] (265.3933pt,437.4999pt) -- (275.3933pt,441.4999pt) -- (265.3933pt,445.4999pt) -- (269.3933pt,441.4999pt) -- cycle;
|
||||
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|
||||
\node at (31pt,432.6666pt)[below right,inner sep=0]{$P$};
|
||||
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||||
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|
||||
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|
||||
\node at (201.7533pt,470.4999pt)[below right,color=black,inner sep=0]{process P2 check locally if~};
|
||||
\node at (201.7533pt,480.4999pt)[below right,inner sep=0]{$\forall j : (j, prove(r)) \not\in P$};
|
||||
\node at (201.7533pt,491.4999pt)[below right,color=black,inner sep=0]{which is false since P1 correctly~};
|
||||
\node at (201.7533pt,501.4999pt)[below right,color=black,inner sep=0]{PROVE(r) and APPEND(r)};
|
||||
\node at (201.7533pt,511.4999pt)[below right,color=black,inner sep=0]{~};
|
||||
\node at (201.7533pt,521.4999pt)[below right,inner sep=0]{$\text{P1 is next include in } W_r$};
|
||||
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|
||||
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|
||||
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|
||||
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|
||||
\draw[pstyle5] (29pt,579.9399pt) -- (276.3933pt,579.9399pt);
|
||||
\node at (41pt,567.9399pt)[below right,inner sep=0]{$READ()$};
|
||||
\draw[pstyle4] (265.3933pt,596.7732pt) -- (275.3933pt,600.7732pt) -- (265.3933pt,604.7732pt) -- (269.3933pt,600.7732pt) -- cycle;
|
||||
\draw[pstyle6] (24pt,600.7732pt) -- (271.3933pt,600.7732pt);
|
||||
\node at (31pt,591.9399pt)[below right,inner sep=0]{$P$};
|
||||
\draw[pstyle9] (189.1283pt,613.7732pt) -- (189.1283pt,706.7732pt) -- (365.1283pt,706.7732pt) -- (365.1283pt,623.7732pt) -- (355.1283pt,613.7732pt) -- (189.1283pt,613.7732pt);
|
||||
\draw[pstyle9] (355.1283pt,613.7732pt) -- (355.1283pt,623.7732pt) -- (365.1283pt,623.7732pt) -- (355.1283pt,613.7732pt);
|
||||
\node at (195.1283pt,618.7732pt)[below right,color=black,inner sep=0]{line(C9)};
|
||||
\node at (195.1283pt,629.7732pt)[below right,color=black,inner sep=0]{process P2 check locally if};
|
||||
\node at (195.1283pt,639.7732pt)[below right,inner sep=0]{$\forall j \in W_r : prop[r][j] = \bot$};
|
||||
\node at (195.1283pt,650.7732pt)[below right,color=black,inner sep=0]{which can't be false since P1 didn't};
|
||||
\node at (195.1283pt,662.6982pt)[below right,color=black,inner sep=0]{execute~};
|
||||
\node at (230.9483pt,660.7732pt)[below right,inner sep=0]{$RBcast(prop, S, r, 1)$};
|
||||
\node at (195.1283pt,672.6982pt)[below right,color=black,inner sep=0]{~};
|
||||
\node at (195.1283pt,682.6982pt)[below right,color=black,inner sep=0]{P2 will never progress and};
|
||||
\node at (195.1283pt,692.6982pt)[below right,color=black,inner sep=0]{deliver any futur messages};
|
||||
\end{tikzpicture}
|
||||
37
Recherche/BFT-ARBover/diagrams/classic_seq.puml
Normal file
37
Recherche/BFT-ARBover/diagrams/classic_seq.puml
Normal file
@@ -0,0 +1,37 @@
|
||||
@startuml
|
||||
!pragma teoz true
|
||||
|
||||
database DL
|
||||
actor P1
|
||||
actor P2
|
||||
actor Pt
|
||||
actor Pn
|
||||
|
||||
P1 ->(05) P2: <latex>RBcast(prop, S, r, 1)</latex>
|
||||
& P1 ->(25) Pt : <latex>RBcast(prop, S, r, 1)</latex>
|
||||
& P1 ->(50) Pn : <latex>RBcast(prop, S, r, 1)</latex>
|
||||
P2 -> P2 : <latex>S'(sk_2, r)</latex>
|
||||
P2 -> P1 : <latex>send(\sigma_2)</latex>
|
||||
... <latex>\text{Wait until P1 received }\sigma \text{ t times}</latex> ...
|
||||
Pt -> Pt : <latex>S'(sk_t, r)</latex>
|
||||
Pt -> P1 : <latex>send(\sigma_t)</latex>
|
||||
|
||||
P1 -> P1 : <latex>C'(pkc, r, J, \{\sigma_r^j\}_{j\in J})</latex>
|
||||
|
||||
P1 -> DL : <latex>PROVE(\sigma)</latex>
|
||||
P1 -> DL : <latex>APPEND(\sigma)</latex>
|
||||
P2 -> Pt
|
||||
|
||||
P1 ->(05) P2: <latex>RBcast(submit, S, r, 1, \sigma)</latex>
|
||||
& P1 ->(25) Pt : <latex>RBcast(submit, S, r, 1, \sigma)</latex>
|
||||
& P1 ->(50) Pn : <latex>RBcast(submit, S, r, 1, \sigma)</latex>
|
||||
|
||||
P2 -> DL : <latex>P \gets READ()</latex>
|
||||
& Pt -> DL
|
||||
& Pn -> DL
|
||||
P2 -> P2 : <latex>V'(pk, r, \sigma)</latex>
|
||||
& Pt -> Pt : <latex>V'(pk, r, \sigma)</latex>
|
||||
& Pn -> Pn : <latex>V'(pk, r, \sigma)</latex>
|
||||
|
||||
hide footbox
|
||||
@enduml
|
||||
127
Recherche/BFT-ARBover/diagrams/classic_seq.tex
Normal file
127
Recherche/BFT-ARBover/diagrams/classic_seq.tex
Normal file
@@ -0,0 +1,127 @@
|
||||
% generated by Plantuml 1.2025.10
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|
||||
\definecolor{plantucolor0003}{RGB}{226,226,240}
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|
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||||
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||||
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|
||||
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||||
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||||
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||||
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|
||||
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|
||||
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|
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||||
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||||
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||||
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|
||||
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|
||||
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|
||||
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|
||||
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||||
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||||
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|
||||
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|
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||||
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||||
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||||
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|
||||
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|
||||
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||||
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||||
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||||
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||||
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|
||||
\end{tikzpicture}
|
||||
32
Recherche/BFT-ARBover/diagrams/nonBFT_behaviour.puml
Normal file
32
Recherche/BFT-ARBover/diagrams/nonBFT_behaviour.puml
Normal file
@@ -0,0 +1,32 @@
|
||||
@startuml
|
||||
!pragma teoz true
|
||||
|
||||
database DL
|
||||
actor P1
|
||||
actor Pi
|
||||
|
||||
P1 -> P1 : <latex>ABcast(m)</latex>
|
||||
P1 -> P1 : <latex>m \in S</latex>
|
||||
|
||||
P1 -> DL : <latex>READ()</latex>
|
||||
DL --> P1 : <latex>P</latex>
|
||||
P1 -> P1 : <latex>r_{max} = max\{r : (\_, prove(r)) \in P\}</latex>
|
||||
|
||||
loop <latex>\textbf{foreach } r \in \{r_{max} + 1, \dots\}</latex>
|
||||
|
||||
P1 ->(05) Pi : <latex>RBcast(prop, S, r, 1)</latex>
|
||||
|
||||
P1 -> DL : <latex>PROVE(r)</latex>
|
||||
P1 -> DL : <latex>APPEND(r)</latex>
|
||||
|
||||
P1 -> DL : <latex>READ()</latex>
|
||||
DL --> P1 : <latex>P</latex>
|
||||
|
||||
alt <latex>(1, \text{prove(}r\text{)}) \in P</latex>
|
||||
note over P1 : break
|
||||
else <latex>(\exists j, r' : (j, prove(r')) \in P \land m \in prop[r'][j])</latex>
|
||||
note over P1 : break
|
||||
end
|
||||
end
|
||||
hide footbox
|
||||
@enduml
|
||||
93
Recherche/BFT-ARBover/diagrams/nonBFT_behaviour.tex
Normal file
93
Recherche/BFT-ARBover/diagrams/nonBFT_behaviour.tex
Normal file
@@ -0,0 +1,93 @@
|
||||
% generated by Plantuml 1.2025.10
|
||||
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||||
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||||
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|
||||
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|
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|
||||
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|
||||
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|
||||
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|
||||
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|
||||
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|
||||
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|
||||
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|
||||
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|
||||
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|
||||
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|
||||
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||||
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||||
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|
||||
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|
||||
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|
||||
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|
||||
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|
||||
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|
||||
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|
||||
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|
||||
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|
||||
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|
||||
\draw[pstyle4] (110.3255pt,228.0576pt) -- (151.3255pt,228.0576pt);
|
||||
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|
||||
\node at (116.3255pt,203.0576pt)[below right,inner sep=0]{$r_{max} = max\{r : (\_, prove(r)) \in P\}$};
|
||||
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|
||||
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|
||||
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|
||||
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|
||||
\node at (92.59pt,242.0576pt)[below right,inner sep=0]{$\textbf{foreach } r \in \{r_{max} + 1, \dots\}$};
|
||||
\node at (215.4514pt,243.3076pt)[below right,color=black,inner sep=0]{\textbf{]}};
|
||||
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|
||||
\draw[pstyle4] (109.3255pt,278.0576pt) -- (275.3933pt,283.0576pt);
|
||||
\node at (116.3255pt,266.0576pt)[below right,inner sep=0]{$RBcast(prop, S, r, 1)$};
|
||||
\draw[pstyle5] (35pt,303.0576pt) -- (25pt,307.0576pt) -- (35pt,311.0576pt) -- (31pt,307.0576pt) -- cycle;
|
||||
\draw[pstyle4] (29pt,307.0576pt) -- (108.3255pt,307.0576pt);
|
||||
\node at (41pt,295.0576pt)[below right,inner sep=0]{$PROVE(r)$};
|
||||
\draw[pstyle5] (35pt,327.0576pt) -- (25pt,331.0576pt) -- (35pt,335.0576pt) -- (31pt,331.0576pt) -- cycle;
|
||||
\draw[pstyle4] (29pt,331.0576pt) -- (108.3255pt,331.0576pt);
|
||||
\node at (41pt,319.0576pt)[below right,inner sep=0]{$APPEND(r)$};
|
||||
\draw[pstyle5] (35pt,351.0576pt) -- (25pt,355.0576pt) -- (35pt,359.0576pt) -- (31pt,355.0576pt) -- cycle;
|
||||
\draw[pstyle4] (29pt,355.0576pt) -- (108.3255pt,355.0576pt);
|
||||
\node at (41pt,343.0576pt)[below right,inner sep=0]{$READ()$};
|
||||
\draw[pstyle5] (97.3255pt,371.8909pt) -- (107.3255pt,375.8909pt) -- (97.3255pt,379.8909pt) -- (101.3255pt,375.8909pt) -- cycle;
|
||||
\draw[pstyle6] (24pt,375.8909pt) -- (103.3255pt,375.8909pt);
|
||||
\node at (31pt,367.0576pt)[below right,inner sep=0]{$P$};
|
||||
\draw[pstyle7] (65.7255pt,387.8909pt) -- (123.9755pt,387.8909pt) -- (123.9755pt,389.8909pt) -- (113.9755pt,399.8909pt) -- (65.7255pt,399.8909pt) -- (65.7255pt,387.8909pt);
|
||||
\draw[pstyle8] (65.7255pt,387.8909pt) rectangle (268.4464pt,499.1498pt);
|
||||
\node at (80.7255pt,388.8909pt)[below right,color=black,inner sep=0]{\textbf{alt}};
|
||||
\node at (138.9755pt,391.1409pt)[below right,color=black,inner sep=0]{\textbf{[}};
|
||||
\node at (142.1655pt,389.8909pt)[below right,inner sep=0]{$(1, \text{prove(}r\text{)}) \in P$};
|
||||
\node at (215.6282pt,391.1409pt)[below right,color=black,inner sep=0]{\textbf{]}};
|
||||
\draw[color=black,line width=1.0pt,dash pattern=on 2.0pt off 2.0pt] (65.7255pt,450.8909pt) -- (268.4464pt,450.8909pt);
|
||||
\node at (70.7255pt,454.1498pt)[below right,color=black,inner sep=0]{\textbf{[}};
|
||||
\node at (73.9155pt,452.8909pt)[below right,inner sep=0]{$(\exists j, r' : (j, prove(r')) \in P \land m \in prop[r'][j])$};
|
||||
\node at (263.2564pt,454.1498pt)[below right,color=black,inner sep=0]{\textbf{]}};
|
||||
\draw[pstyle10] (86.7255pt,414.8909pt) -- (86.7255pt,434.8909pt) -- (131.7255pt,434.8909pt) -- (131.7255pt,424.8909pt) -- (121.7255pt,414.8909pt) -- (86.7255pt,414.8909pt);
|
||||
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|
||||
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|
||||
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|
||||
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|
||||
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|
||||
\end{tikzpicture}
|
||||
BIN
Recherche/BFT-ARBover/main.pdf
Normal file
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Recherche/BFT-ARBover/main.pdf
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933
Recherche/BFT-ARBover/main.tex
Normal file
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Recherche/BFT-ARBover/main.tex
Normal file
@@ -0,0 +1,933 @@
|
||||
\documentclass[11pt]{article}
|
||||
|
||||
\usepackage[margin=1in]{geometry}
|
||||
\usepackage[T1]{fontenc}
|
||||
\usepackage[utf8]{inputenc}
|
||||
\usepackage{lmodern}
|
||||
\usepackage{microtype}
|
||||
\usepackage{amsmath,amssymb,amsthm,mathtools}
|
||||
\usepackage{thmtools}
|
||||
\usepackage{enumitem}
|
||||
\usepackage{csquotes}
|
||||
\usepackage[hidelinks]{hyperref}
|
||||
\usepackage[nameinlink,noabbrev]{cleveref}
|
||||
\usepackage{algorithm}
|
||||
\usepackage{algpseudocode}
|
||||
\usepackage{graphicx}
|
||||
% Line-number prefix configuration (A/B/C)
|
||||
\renewcommand{\thealgorithm}{\Alph{algorithm}} % Float labels: Algorithm A, B, C
|
||||
\newcommand{\algletter}{}
|
||||
\algrenewcommand\alglinenumber[1]{\scriptsize\textbf{\algletter}#1}
|
||||
\usepackage{tikz}
|
||||
\graphicspath{{diagrams/out}}
|
||||
\usepackage{xspace}
|
||||
% \usepackage{plantuml}
|
||||
|
||||
\usepackage[fr-FR]{datetime2}
|
||||
|
||||
\usepackage{fancyhdr}
|
||||
\pagestyle{fancy}
|
||||
\fancyhf{}
|
||||
\fancyfoot[L]{Compilé le \DTMnow}
|
||||
\fancyfoot[C]{\thepage}
|
||||
\renewcommand{\headrulewidth}{0pt}
|
||||
\renewcommand{\footrulewidth}{0pt}
|
||||
|
||||
\theoremstyle{plain}
|
||||
\newtheorem{theorem}{Theorem}
|
||||
\newtheorem{lemma}[theorem]{Lemma}
|
||||
\newtheorem{corollary}[theorem]{Corollary}
|
||||
\theoremstyle{definition}
|
||||
\newtheorem{definition}{Definition}
|
||||
\theoremstyle{remark}
|
||||
\newtheorem{remark}{Remark}
|
||||
|
||||
\newcommand{\RB}{\textsf{RB}\xspace}
|
||||
\newcommand{\res}{\mathsf{res}}
|
||||
\newcommand{\ARB}{\textsf{ARB}\xspace}
|
||||
\newcommand{\DL}{\textsf{DL}\xspace}
|
||||
\newcommand{\APPEND}{\textsf{APPEND}}
|
||||
\newcommand{\PROVE}{\textsf{PROVE}}
|
||||
\newcommand{\PROVEtrace}{\textsf{prove}}
|
||||
\newcommand{\READ}{\textsf{READ}}
|
||||
\newcommand{\ABbroadcast}{\textsf{AB-broadcast}}
|
||||
\newcommand{\ABdeliver}{\textsf{AB-deliver}}
|
||||
\newcommand{\RBcast}{\textsf{RB-cast}}
|
||||
\newcommand{\RBreceived}{\textsf{RB-received}}
|
||||
\newcommand{\ordered}{\textsf{ordered}}
|
||||
\newcommand{\Winners}{\mathsf{Winners}}
|
||||
\newcommand{\Messages}{\mathsf{Messages}}
|
||||
\newcommand{\ABlisten}{\textsf{AB-listen}}
|
||||
|
||||
\newcommand{\SHARE}{\mathsf{SHARE}}
|
||||
\newcommand{\COMBINE}{\mathsf{COMBINE}}
|
||||
\newcommand{\VERIFY}{\mathsf{VERIFY}}
|
||||
|
||||
\newcommand{\delivered}{\mathsf{delivered}}
|
||||
\newcommand{\received}{\mathsf{received}}
|
||||
\newcommand{\prop}{\mathsf{prop}}
|
||||
\newcommand{\resolved}{\mathsf{resolved}}
|
||||
\newcommand{\current}{\mathsf{current}}
|
||||
|
||||
\newcommand{\Seq}{\mathsf{Seq}}
|
||||
|
||||
\crefname{theorem}{Theorem}{Theorems}
|
||||
\crefname{lemma}{Lemma}{Lemmas}
|
||||
\crefname{definition}{Definition}{Definitions}
|
||||
\crefname{algorithm}{Algorithm}{Algorithms}
|
||||
% Code exécuté par tout processus p_i
|
||||
|
||||
\begin{document}
|
||||
|
||||
\section{Model}
|
||||
We consider a static set of $n$ processes with known identities, communicating by reliable point-to-point channels, in a complete graph. Messages are uniquely identifiable.
|
||||
|
||||
\paragraph{Synchrony.} The network is asynchronous. Processes may crash; at most $f$ crashes occur.
|
||||
|
||||
\paragraph{Communication.} Processes can exchange through a Reliable Broadcast (\RB) primitive (defined below) which's invoked with the functions \RBcast$(m)$ and \RBreceived$(m)$. There exists a shared object called DenyList (\DL) (defined below) that is interfaced with the functions \APPEND$(x)$, \PROVE$(x)$ and \READ$()$.
|
||||
|
||||
\paragraph{Notation.} Let $\Pi$ be the finite set of process identifiers and let $n \triangleq |\Pi|$. Two authorization subsets are $\Pi_M \subseteq \Pi$ (processes allowed to issue \APPEND) and $\Pi_V \subseteq \Pi$ (processes allowed to issue \PROVE). Indices $i,j \in \Pi$ refer to processes, and $p_i$ denotes the process with identifier $i$. Let $\mathcal{M}$ denote the universe of uniquely identifiable messages, with $m \in \mathcal{M}$. Let $\mathcal{R} \subseteq \mathbb{N}$ be the set of round identifiers; we write $r \in \mathcal{R}$ for a round. We use the precedence relation $\prec$ for the \DL{} linearization: $x \prec y$ means that operation $x$ appears strictly before $y$ in the linearized history of \DL. For any finite set $A \subseteq \mathcal{M}$, \ordered$(A)$ returns a deterministic total order over $A$ (e.g., lexicographic order on $(\textit{senderId},\textit{messageId})$ or on message hashes). For any round $r \in \mathcal{R}$, define $\Winners_r \triangleq \{\, j \in \Pi \mid (j,\PROVEtrace(r)) \prec \APPEND(r) \,\}$, i.e., the set of processes whose $\PROVE(r)$ appears before the first $\APPEND(r)$ in the \DL{} linearization.
|
||||
We denoted by $\PROVE^{(j)}(r)$ or $\APPEND^{(j)}(r)$ the operation $\PROVE(r)$ or $\APPEND(r)$ invoked by process $j$.
|
||||
|
||||
% ------------------------------------------------------------------------------
|
||||
\section{Primitives}
|
||||
\subsection{Reliable Broadcast (RB)}
|
||||
|
||||
\RB provides the following properties in the model.
|
||||
\begin{itemize}[leftmargin=*]
|
||||
\item \textbf{Integrity}: Every message received was previously sent. $\forall p_i:\ \RBreceived_i(m) \Rightarrow \exists p_j:\ \RBcast_j(m)$.
|
||||
\item \textbf{No-duplicates}: No message is received more than once at any process.
|
||||
\item \textbf{Validity}: If a correct process broadcasts $m$, every correct process eventually receives $m$.
|
||||
\end{itemize}
|
||||
|
||||
\subsection{DenyList (DL)}
|
||||
We assume a synchronous DenyList (\DL) object with the following properties.
|
||||
|
||||
The DenyList object type supports three operations: $\APPEND$, $\PROVE$, and $\READ$. These operations appear as if executed in a sequence $\Seq$ such that:
|
||||
\begin{itemize}
|
||||
\item \textbf{Termination.} A $\PROVE$, an $\APPEND$, or a $\READ$ operation invoked by a correct process always returns.
|
||||
\item \textbf{APPEND Validity.} The invocation of $\APPEND(x)$ by a process $p$ is valid if:
|
||||
\begin{itemize}
|
||||
\item $p \in \Pi_M \subseteq \Pi$; \textbf{and}
|
||||
\item $x \in S$, where $S$ is a predefined set.
|
||||
\end{itemize}
|
||||
Otherwise, the operation is invalid.
|
||||
\item \textbf{PROVE Validity.} If the invocation of a $op = \PROVE(x)$ by a correct process $p$ is not valid, then:
|
||||
\begin{itemize}
|
||||
\item $p \not\in \Pi_V \subseteq \Pi$; \textbf{or}
|
||||
\item A valid $\APPEND(x)$ appears before $op$ in $\Seq$.
|
||||
\end{itemize}
|
||||
Otherwise, the operation is valid.
|
||||
\item \textbf{PROVE Anti-Flickering.} If the invocation of a operation $op = \PROVE(x)$ by a correct process $p \in \Pi_V$ is invalid, then any $\PROVE(x)$ operation that appears after $op$ in $\Seq$ is invalid.
|
||||
\item \textbf{READ Validity.} The invocation of $op = \READ()$ by a process $p \in \pi_V$ returns the list of valid invocations of $\PROVE$ that appears before $op$ in $\Seq$ along with the names of the processes that invoked each operation.
|
||||
\item \textbf{Anonymity.} Let us assume the process $p$ invokes a $\PROVE(v)$ operation. If the process $p'$ invokes a $\READ()$ operation, then $p'$ cannot learn the value $v$ unless $p$ leaks additional information.
|
||||
\end{itemize}
|
||||
|
||||
% ------------------------------------------------------------------------------
|
||||
\section{Target Abstraction: Atomic Reliable Broadcast (ARB)}
|
||||
Processes export \ABbroadcast$(m)$ and \ABdeliver$(m)$. \ARB requires total order:
|
||||
\begin{equation*}
|
||||
\forall m_1,m_2,\ \forall p_i,p_j:\ \ \ABdeliver_i(m_1) < \ABdeliver_i(m_2) \Rightarrow \ABdeliver_j(m_1) < \ABdeliver_j(m_2),
|
||||
\end{equation*}
|
||||
plus Integrity/No-duplicates/Validity (inherited from \RB and the construction).
|
||||
|
||||
|
||||
\section{ARB over RB and DL}
|
||||
|
||||
% ------------------------------------------------------------------------------
|
||||
|
||||
\subsection{Algorithm}
|
||||
% granularité diff commentaire de code et paragraphe pre algo
|
||||
|
||||
\begin{definition}[Closed round]\label{def:closed-round}
|
||||
Given a \DL{} linearization $H$, a round $r\in\mathcal{R}$ is \emph{closed} in $H$ iff $H$ contains an operation $\APPEND(r)$.
|
||||
Equivalently, there exists a time after which every $\PROVE(r)$ is invalid in $H$.
|
||||
\end{definition}
|
||||
|
||||
\subsubsection{Variables}
|
||||
Each process $p_i$ maintains:
|
||||
|
||||
%on met toutes les variables locales ici
|
||||
\begin{algorithmic}
|
||||
\State $\received \gets \emptyset$ \Comment{Messages received via \RB but not yet delivered}
|
||||
\State $\delivered \gets \emptyset$ \Comment{Messages already delivered}
|
||||
\State $\prop[r][j] \gets \bot,\ \forall r,j$ \Comment{Proposal from process $j$ for round $r$}
|
||||
\State $\current \gets 0$
|
||||
\end{algorithmic}
|
||||
|
||||
\paragraph{DenyList.} The \DL is initialized empty. We assume $\Pi_M = \Pi_V = \Pi$ (all processes can invoke \APPEND and \PROVE).
|
||||
|
||||
\subsubsection{Handlers and Procedures}
|
||||
|
||||
\renewcommand{\algletter}{A}
|
||||
\begin{algorithm}[H]
|
||||
\caption{\RB handler (at process $p_i$)}\label{alg:rb-handler}
|
||||
\begin{algorithmic}[1]
|
||||
\Function{RBreceived}{$S, r, j$}
|
||||
% \State \textbf{on} $\RBreceived(S, r, j)$ \textbf{do}
|
||||
\State $\received \leftarrow \received \cup \{S\}$
|
||||
\State $\prop[r][j] \leftarrow S$ \Comment{Record sender $j$'s proposal $S$ for round $r$}
|
||||
\EndFunction
|
||||
\end{algorithmic}
|
||||
\end{algorithm}
|
||||
|
||||
% \paragraph{} An \ABbroadcast$(m)$ chooses the next open round from the \DL view, proposes all pending messages together with the new $m$, disseminates this proposal via \RB, then executes $\PROVE(r)$ followed by $\APPEND(r)$ to freeze the winners of the round. The loop polls \DL until (i) some winner’s proposal includes $m$ in a \emph{closed} round and (ii) all winners' proposals for closed rounds are known locally, ensuring eventual inclusion and delivery.
|
||||
|
||||
\renewcommand{\algletter}{B}
|
||||
\begin{algorithm}[H]
|
||||
\caption{\ABbroadcast$(m)$ (at process $p_i$)}\label{alg:ab-bcast}
|
||||
\begin{algorithmic}[1]
|
||||
\Function{ABbroadcast}{$m$}
|
||||
\State $P \leftarrow \READ()$ \Comment{Fetch latest \DL state to learn recent $\PROVE$ operations}
|
||||
\State $r_{max} \leftarrow \max(\{ r' : \exists j,\ (j,\PROVE(r')) \in P \})$ \Comment{Pick current open round}
|
||||
\State $S \leftarrow (\received \setminus \delivered) \cup \{m\}$ \Comment{Propose all pending messages plus the new $m$}
|
||||
|
||||
\vspace{1em}
|
||||
|
||||
\For{\textbf{each}\ $r \in \{r_{max}, r_{max}+1, \cdots \}$}
|
||||
\State $\RBcast(S, r, i)$; $\PROVE(r)$; $\APPEND(r)$;
|
||||
\State $P \leftarrow \READ()$ \Comment{Refresh local view of \DL}
|
||||
\If{($\big((i, \PROVEtrace(r)) \in P\ \vee\ (\exists j, r': (j, \PROVEtrace(r')) \in P \wedge \ m \in \prop[r'][j]))$)}
|
||||
\State \textbf{break} \Comment{Exit loop once $m$ is included in some closed round}
|
||||
\EndIf
|
||||
\EndFor
|
||||
\EndFunction
|
||||
\end{algorithmic}
|
||||
\end{algorithm}
|
||||
|
||||
\renewcommand{\algletter}{C}
|
||||
\begin{algorithm}[H]
|
||||
\caption{\ABdeliver() at process $p_i$}\label{alg:delivery}
|
||||
\begin{algorithmic}[1]
|
||||
\Function{ABdeliver}{}
|
||||
\State $r \gets \current$
|
||||
\State $P \gets \READ()$
|
||||
\If{$\forall j : (j, \PROVEtrace(r)) \not\in P$}
|
||||
\State \Return $\bot$
|
||||
\EndIf
|
||||
\State $\APPEND(r)$; $P \gets \READ()$
|
||||
\State $W_r \gets \{ j : (j, \PROVEtrace(r)) \in P \}$
|
||||
\If{$\exists j \in W_r,\ \prop[r][j] = \bot$}
|
||||
\State \Return $\bot$
|
||||
\EndIf
|
||||
\State $M_r \gets \bigcup_{j \in W_r} \prop[r][j]$
|
||||
\State $m \gets \ordered(M_r \setminus \delivered)[0]$ \Comment{Set $m$ as the smaller message not already delivered}
|
||||
\State $\delivered \leftarrow \delivered \cup \{m\}$
|
||||
\If{$M_r \setminus \delivered = \emptyset$} \Comment{Check if all messages from round $r$ have been delivered}
|
||||
\State $\current \leftarrow \current + 1$
|
||||
\EndIf
|
||||
\State \textbf{return} $m$
|
||||
\EndFunction
|
||||
\end{algorithmic}
|
||||
\end{algorithm}
|
||||
|
||||
% ------------------------------------------------------------------------------
|
||||
\subsection{Correctness}
|
||||
|
||||
\begin{lemma}[Stable round closure]\label{lem:closure-stable}
|
||||
If a round $r$ is closed, then there exists a linearization point $t_0$ of $\APPEND(r)$ in the \DL, and from that point on, no $\PROVE(r)$ can be valid.
|
||||
Once closed, a round never becomes open again.
|
||||
\end{lemma}
|
||||
|
||||
\begin{proof}
|
||||
By \Cref{def:closed-round}, some $\APPEND(r)$ occurs in the linearization $H$. \\
|
||||
$H$ is a total order of operations, the set of $\APPEND(r)$ operations is totally ordered, and hence there exists a smallest $\APPEND(r)$ in $H$. We denote this operation $\APPEND^{(\star)}(r)$ and $t_0$ its linearization point. \\
|
||||
By the validity property of \DL, a $\PROVE(r)$ is valid iff $\PROVE(r) \prec \APPEND^{(\star)}(r)$. Thus, after $t_0$, no $\PROVE(r)$ can be valid. \\
|
||||
$H$ is a immutable grow-only history, and hence once closed, a round never becomes open again. \\
|
||||
Hence there exists a linearization point $t_0$ of $\APPEND(r)$ in the \DL, and from that point on, no $\PROVE(r)$ can be valid and the closure is stable.
|
||||
\end{proof}
|
||||
|
||||
\begin{definition}[First APPEND]\label{def:first-append}
|
||||
Given a \DL{} linearization $H$, for any closed round $r\in\mathcal{R}$, we denote by $\APPEND^{(\star)}(r)$ the earliest $\APPEND(r)$ in $H$.
|
||||
\end{definition}
|
||||
|
||||
\begin{lemma}[Across rounds]\label{lem:across}
|
||||
If there exists a $r$ such that $r$ is closed, $\forall r'$ such that $r' < r$, r' is also closed.
|
||||
\end{lemma}
|
||||
|
||||
\begin{proof}
|
||||
\emph{Base.} For a closed round $k=0$, the set $\{k' \in \mathcal{R},\ k' < k\}$ is empty, hence the lemma is true.
|
||||
|
||||
\emph{Induction.} Assume the lemma is true for round $k\geq 0$, we prove it for round $k+1$.
|
||||
|
||||
\smallskip
|
||||
Assume $k+1$ is closed and let $\APPEND^{(\star)}(k+1)$ be the earliest $\APPEND(k+1)$ in the DL linearization $H$.
|
||||
By Lemma 1, after an $\APPEND(k)$ is in $H$, any later $\PROVE(k)$ is rejected also, a process’s program order is preserved in $H$.
|
||||
|
||||
There are two possibilities for where $\APPEND^{(\star)}(k+1)$ is invoked.
|
||||
|
||||
\begin{itemize}
|
||||
\item \textbf{Case (B6) :}
|
||||
Some process $p^\star$ executes the loop (lines B5–B11) and invokes $\APPEND^{(\star)}(k+1)$ at line B6.
|
||||
Immediately before line B6, line B5 sets $r\leftarrow r+1$, so the previous loop iteration (if any) targeted $k$. We consider two sub-cases.
|
||||
|
||||
\begin{itemize}
|
||||
\item \emph{(i) $p^\star$ is not in its first loop iteration.}
|
||||
In the previous iteration, $p^\star$ executed $\PROVE^{(\star)}(k)$ at B6, followed (in program order) by $\APPEND^{(\star)}(k)$.
|
||||
If round $k$ wasn't closed when $p^\star$ execute $\PROVE^{(\star)}(k)$ at B9, then the condition at B8 would be true hence the tuple $(p^\star, \PROVEtrace(k))$ should be visible in P which implies that $p^\star$ should leave the loop at round $k$, contradicting the assumption that $p^\star$ is now executing another iteration.
|
||||
Since the tuple is not visible, the $\PROVE^{(\star)}(k)$ was rejected by the DL which implies by definition an $\APPEND(k)$ already exists in $H$. Thus in this case $k$ is closed.
|
||||
|
||||
\item \emph{(ii) $p^\star$ is in its first loop iteration.}
|
||||
To compute the value $r_{max}$, $p^\star$ must have observed one or many $(\_ , \PROVEtrace(k+1))$ in $P$ at B2/B3, issued by some processes (possibly different from $p^\star$). Let's call $p_1$ the issuer of the first $\PROVE^{(1)}(k+1)$ in the linearization $H$. \\
|
||||
When $p_1$ executed $P \gets \READ()$ at B2 and compute $r_{max}$ at B3, he observed no tuple $(\_,\PROVEtrace(k+1))$ in $P$ because he's the issuer of the first one. So when $p_1$ executed the loop (B5–B11), he run it for the round $k$, didn't seen any $(1,\PROVEtrace(k))$ in $P$ at B8, and then executed the first $\PROVE^{(1)}(k+1)$ at B6 in a second iteration. \\
|
||||
If round $k$ wasn't closed when $p_1$ execute $\PROVE^{(1)}(k)$ at B6, then the condition at B8 should be true which implies that $p_1$ sould leave the loop at round $k$, contradicting the assumption that $p_1$ is now executing $\PROVE^{(1)}(r+1)$. In this case $k$ is closed.
|
||||
\end{itemize}
|
||||
|
||||
\item \textbf{Case (C8) :}
|
||||
Some process invokes $\APPEND(k+1)$ at C8.
|
||||
Line C8 is guarded by the presence of $\PROVE(\textit{next})$ in $P$ with $\textit{next}=k+1$ (C5).
|
||||
Moreover, the local pointer $\textit{next}$ grow by increment of 1 and only advances after finishing the current round (C17), so if a process can reach $\textit{next}=k+1$ it implies that he has completed round $k$, which includes closing $k$ at C8 when $\PROVE(k)$ is observed.
|
||||
Hence $\APPEND^\star(k+1)$ implies a prior $\APPEND(k)$ in $H$, so $k$ is closed.
|
||||
\end{itemize}
|
||||
|
||||
\smallskip
|
||||
In all cases, $k+1$ closed implie $k$ closed. By induction on $k$, if the lemme is true for a closed $k$ then it is true for a closed $k+1$.
|
||||
Therefore, the lemma is true for all closed rounds $r$.
|
||||
\end{proof}
|
||||
|
||||
\begin{definition}[Winner Invariant]\label{def:winner-invariant}
|
||||
For any closed round $r$, define
|
||||
\[
|
||||
\Winners_r \triangleq \{ j : \PROVE^{(j)}(r) \prec \APPEND^\star(r) \}
|
||||
\]
|
||||
as the unique set of winners of round $r$.
|
||||
\end{definition}
|
||||
|
||||
\begin{lemma}[Invariant view of closure]\label{lem:closure-view}
|
||||
For any closed round $r$, all correct processes eventually observe the same set of valid tuples $(\_,\PROVEtrace(r))$ in their \DL view.
|
||||
\end{lemma}
|
||||
|
||||
\begin{proof}
|
||||
Let's take a closed round $r$. By \Cref{def:first-append}, there exists a unique earliest $\APPEND(r)$ in the DL linearization, denoted $\APPEND^\star(r)$.
|
||||
|
||||
Consider any correct process $p$ that invokes $\READ()$ after $\APPEND^\star(r)$ in the DL linearization. Since $\APPEND^\star(r)$ invalidates all subsequent $\PROVE(r)$, the set of valid tuples $(\_,\PROVEtrace(r))$ observed by any correct process after $\APPEND^\star(r)$ is fixed and identical across all correct processes.
|
||||
|
||||
Therefore, for any closed round $r$, all correct processes eventually observe the same set of valid tuples $(\_,\PROVEtrace(r))$ in their \DL view.
|
||||
\end{proof}
|
||||
|
||||
\begin{lemma}[Well-defined winners]\label{lem:winners}
|
||||
For any correct process and round $r$, if the process computes $W_r$ at line C9, then :
|
||||
\begin{itemize}
|
||||
\item $\Winners_r$ is defined;
|
||||
\item the computed $W_r$ is exactly $\Winners_r$.
|
||||
\end{itemize}
|
||||
\end{lemma}
|
||||
|
||||
\begin{proof}
|
||||
Let take a correct process $p_i$ that reach line C9 to compute $W_r$. \\
|
||||
By program order, $p_i$ must have executed $\APPEND^{(i)}(r)$ at C8 before, which implies by \Cref{def:closed-round} that round $r$ is closed. So by \Cref{def:winner-invariant}, $\Winners_r$ is defined. \\
|
||||
By \Cref{lem:closure-view}, all correct processes eventually observe the same set of valid tuples $(\_,\PROVEtrace(r))$ in their \DL view. Hence, when $p_i$ executes the $\READ()$ at C8 after the $\APPEND^{(i)}(r)$, it observes a set $P$ that includes all valid tuples $(\_,\PROVEtrace(r))$ such that
|
||||
\[
|
||||
W_r = \{ j : (j,\PROVEtrace(r)) \in P \} = \{j : \PROVE^{(j)}(r) \prec \APPEND^\star(r) \} = \Winners_r
|
||||
\]
|
||||
\end{proof}
|
||||
|
||||
\begin{lemma}[No APPEND without PROVE]\label{lem:append-prove}
|
||||
If some process invokes $\APPEND(r)$, then at least a process must have previously invoked $\PROVE(r)$.
|
||||
\end{lemma}
|
||||
|
||||
\begin{proof}[Proof]
|
||||
Consider the round $r$ such that some process invokes $\APPEND(r)$. There are two possible cases
|
||||
|
||||
\begin{itemize}
|
||||
\item \textbf{Case (B6) :}
|
||||
There exists a process $p^\star$ who's the issuer of the earliest $\APPEND^{(\star)}(r)$ in the DL linearization $H$. By program order, $p^\star$ must have previously invoked $\PROVE^{(\star)}(r)$ before invoking $\APPEND^{(\star)}(r)$ at B6. In this case, there is at least one $\PROVE(r)$ valid in $H$ issued by a correct process before $\APPEND^{(\star)}(r)$.
|
||||
|
||||
\item \textbf{Case (C8) :}
|
||||
There exist a process $p^\star$ invokes $\APPEND^{(\star)}(r)$ at C8. Line C8 is guarded by the condition at C5, which ensures that $p$ observed some $(\_,\PROVEtrace(r))$ in $P$. In this case, there is at least one $\PROVE(r)$ valid in $H$ issued by some process before $\APPEND^{(\star)}(r)$.
|
||||
\end{itemize}
|
||||
|
||||
In both cases, if some process invokes $\APPEND(r)$, then some process must have previously invoked $\PROVE(r)$.
|
||||
\end{proof}
|
||||
|
||||
\begin{lemma}[No empty winners]\label{lem:nonempty}
|
||||
Let $r$ be a round, if $\Winners_r$ is defined, then $\Winners_r \neq \emptyset$.
|
||||
\end{lemma}
|
||||
|
||||
\begin{proof}[Proof]
|
||||
If $\Winners_r$ is defined, then by \Cref{def:winner-invariant}, round $r$ is closed. By \Cref{def:closed-round}, some $\APPEND(r)$ occurs in the DL linearization. \\
|
||||
By \Cref{lem:append-prove}, at least a process must have invoked a valid $\PROVE(r)$ before $\APPEND^{(\star)}(r)$. Hence, there exists at least one $j$ such that $\{j: \PROVE^{(j)}(r) \prec \APPEND^\star(r)\}$, so $\Winners_r \neq \emptyset$.
|
||||
\end{proof}
|
||||
|
||||
\begin{lemma}[Winners must propose]\label{lem:winners-propose}
|
||||
For any closed round $r$, $\forall j \in \Winners_r$, process $j$ must have invoked a $\RBcast(S^{(j)}, r, j)$
|
||||
\end{lemma}
|
||||
|
||||
\begin{proof}[Proof]
|
||||
Fix a closed round $r$. By \Cref{def:winner-invariant}, for any $j \in \Winners_r$, there exist a valid $\PROVE^{(j)}(r)$ such that $\PROVE^{(j)}(r) \prec \APPEND^\star(r)$ in the DL linearization. By program order, if $j$ invoked a valid $\PROVE^{(j)}(r)$ at line B6 he must have invoked $\RBcast(S^{(j)}, r, j)$ directly before.
|
||||
\end{proof}
|
||||
|
||||
\begin{definition}[Messages invariant]\label{def:messages-invariant}
|
||||
For any closed round $r$ and any correct process $p_i$ such that $\nexists j \in \Winners_r : prop^{[i)}[r][j] = \bot$, define
|
||||
\[
|
||||
\Messages_r \triangleq \bigcup_{j\in\Winners_r} \prop^{(i)}[r][j]
|
||||
\]
|
||||
as the unique set of messages proposed by the winners of round $r$.
|
||||
\end{definition}
|
||||
|
||||
\begin{lemma}[Non-empty winners proposal]\label{lem:winner-propose-nonbot}
|
||||
For any closed round $r$, $\forall j \in \Winners_r$, for any correct process $p_i$, eventually $\prop^{(i)}[r][j] \neq \bot$.
|
||||
\end{lemma}
|
||||
|
||||
\begin{proof}[Proof]
|
||||
Fix a closed round $r$. By \Cref{def:winner-invariant}, for any $j \in \Winners_r$, there exist a valid $\PROVE^{(j)}(r)$ such that $\PROVE^{(j)}(r) \prec \APPEND^\star(r)$ in the DL linearization. By \Cref{lem:winners-propose}, $j$ must have invoked $\RBcast(S^{(j)}, r, j)$.
|
||||
|
||||
Let take a process $p_i$, by \RB \emph{Validity}, every correct process eventually receives $j$'s \RB message for round $r$, which sets $\prop[r][j]$ to a non-$\bot$ value at line A3.
|
||||
\end{proof}
|
||||
|
||||
\begin{lemma}[Eventual proposal closure]\label{lem:eventual-closure}
|
||||
If a correct process $p_i$ define $M_r$ at line C13, then for every $j \in \Winners_r$, $\prop^{(i)}[r][j] \neq \bot$.
|
||||
\end{lemma}
|
||||
|
||||
\begin{proof}[Proof]
|
||||
Let take a correct process $p_i$ that computes $M_r$ at line C13. By \Cref{lem:winners}, $p_i$ computes the unique winner set $\Winners_r$.
|
||||
|
||||
By \Cref{lem:nonempty}, $\Winners_r \neq \emptyset$. The instruction at line C13 where $p_i$ computes $M_r$ is guarded by the condition at C10, which ensures that $p_i$ has received all \RB messages from every winner $j \in \Winners_r$. Hence, when $p_i$ computes $M_r = \bigcup_{j\in\Winners_r} \prop^{(i)}[r][j]$, we have $\prop^{(i)}[r][j] \neq \bot$ for all $j \in \Winners_r$.
|
||||
\end{proof}
|
||||
|
||||
\begin{lemma}[Unique proposal per sender per round]\label{lem:unique-proposal}
|
||||
For any round $r$ and any process $p_i$, $p_i$ invokes at most one $\RBcast(S, r, i)$.
|
||||
\end{lemma}
|
||||
|
||||
\begin{proof}[Proof]
|
||||
By program order, any process $p_i$ invokes $\RBcast(S, r, i)$ at line B6 must be in the loop B5–B11. No matter the number of iterations of the loop, line B5 always uses the current value of $r$ which is incremented by 1 at each turn. Hence, in any execution, any process $p_i$ invokes $\RBcast(S, r, j)$ at most once for any round $r$.
|
||||
\end{proof}
|
||||
|
||||
\begin{lemma}[Proposal convergence]\label{lem:convergence}
|
||||
For any round $r$, for any correct processes $p_i$ that define $M_r$ at line C13, we have
|
||||
\[
|
||||
M_r^{(i)} = \Messages_r
|
||||
\]
|
||||
\end{lemma}
|
||||
|
||||
\begin{proof}[Proof]
|
||||
Let take a correct process $p_i$ that define $M_r$ at line C13. That implies that $p_i$ has defined $W_r$ at line C9. It implies that, by \Cref{lem:winners}, $r$ is closed and $W_r = \Winners_r$. \\
|
||||
By \Cref{lem:eventual-closure}, for every $j \in \Winners_r$, $\prop^{(i)}[r][j] \neq \bot$. By \Cref{lem:unique-proposal}, each winner $j$ invokes at most one $\RBcast(S^{(j)}, r, j)$, so $\prop^{(i)}[r][j] = S^{(j)}$ is uniquely defined. Hence, when $p_i$ computes
|
||||
\[
|
||||
M_r^{(i)} = \bigcup_{j\in\Winners_r} \prop^{(i)}[r][j] = \bigcup_{j\in\Winners_r} S^{(j)} = \Messages_r.
|
||||
\]
|
||||
\end{proof}
|
||||
|
||||
\begin{lemma}[Inclusion]\label{lem:inclusion}
|
||||
If some correct process invokes $\ABbroadcast(m)$, then there exist a round $r$ and a process $j\in\Winners_r$ such that $p_j$ invokes
|
||||
\[
|
||||
\RBcast(S, r, j)\quad\text{with}\quad m\in S.
|
||||
\]
|
||||
\end{lemma}
|
||||
|
||||
\begin{proof}
|
||||
Fix a correct process $p_i$ that invokes $\ABbroadcast(m)$ and eventually exits the loop (B5–B11) at some round $r$. There are two possible cases.
|
||||
|
||||
\begin{itemize}
|
||||
\item \textbf{Case 1:} $p_i$ exits the loop because $(i, \PROVEtrace(r)) \in P$. In this case, by \Cref{def:winner-invariant}, $p_i$ is a winner in round $r$. By program order, $p_i$ must have invoked $\RBcast(S, r, i)$ at B6 before invoking $\PROVE^{(i)}(r)$ at B7. By line B4, $m \in S$. Hence there exist a closed round $r$ and a correct process $j=i\in\Winners_r$ such that $j$ invokes $\RBcast(S, r, j)$ with $m\in S$.
|
||||
|
||||
\item \textbf{Case 2:} $p_i$ exits the loop because $\exists j, r': (j, \PROVEtrace(r')) \in P \wedge m \in \prop[r'][j]$. In this case, by \Cref{lem:winners-propose} and \Cref{lem:unique-proposal} $j$ must have invoked a unique $\RBcast(S, r', j)$. Which set $\prop^{(i)}[r'][j] = S$ with $m \in S$.
|
||||
\end{itemize}
|
||||
|
||||
In both cases, if some correct process invokes $\ABbroadcast(m)$, then there exist a round $r$ and a correct process $j\in\Winners_r$ such that $j$ invokes
|
||||
\[
|
||||
\RBcast(S, r, j)\quad\text{with}\quad m\in S.
|
||||
\]
|
||||
\end{proof}
|
||||
|
||||
\begin{lemma}[Broadcast Termination]\label{lem:bcast-termination}
|
||||
If a correct process invokes $\ABbroadcast(m)$, then he eventually exit the function and returns.
|
||||
\end{lemma}
|
||||
|
||||
\begin{proof}[Proof]
|
||||
Let a correct process $p_i$ that invokes $\ABbroadcast(m)$. The lemma is true if $\exists r_1$ such that $r_1 \geq r_{max}$ and if $(i, \PROVEtrace(r_1)) \in P$; or if $\exists r_2$ such that $r_2 \geq r_{max}$ and if $\exists j: (j, \PROVEtrace(r_2)) \in P \wedge m \in \prop[r_2][j]$ (like guarded at B8).
|
||||
|
||||
Let admit that there exists no round $r_1$ such that $p_i$ invokes a valid $\PROVE(r_1)$. In this case $p_i$ invokes infinitely many $\RBcast(S, \_, i)$ at B6 with $m \in S$ (line B4).\\
|
||||
The assumption that no $\PROVE(r_1)$ invoked by $p$ is valid implies by \DL \emph{Validity} that for every round $r' \geq r_{max}$, there exists at least one $\APPEND(r')$ in the DL linearization, hence by \Cref{lem:nonempty} for every possible round $r'$ there at least a winner. Because there is an infinite number of rounds, and a finite number of processes, there exists at least a correct process $p_k$ that invokes infinitely many valid $\PROVE(r')$ and by extension infinitely many $\ABbroadcast(\_)$. By \RB \emph{Validity}, $p_k$ eventually receives $p_i$ 's \RB messages. Let call $t_0$ the time when $p_k$ receives $p_i$ 's \RB message. \\
|
||||
At $t_0$, $p_k$ execute \Cref{alg:rb-handler} and do $\received \leftarrow \received \cup \{S\}$ with $m \in S$ (line A2).
|
||||
For the first invocation of $\ABbroadcast(\_)$ by $p_k$ after the execution of \Cref{alg:rb-handler}, $p_k$ must include $m$ in his proposal $S$ at line B4 (because $m$ is pending in $j$'s $\received \setminus \delivered$ set). There exists a minimum round $r_2$ such that $p_k \in \Winners_{r_2}$ after $t_0$. By \Cref{lem:winner-propose-nonbot}, eventually $\prop^{(i)}[r_2][k] \neq \bot$. By \Cref{lem:unique-proposal}, $\prop^{(i)}[r_2][k]$ is uniquely defined as the set $S$ proposed by $p_k$ at B6, which by \Cref{lem:inclusion} includes $m$. Hence eventually $m \in \prop^{(i)}[r_2][k]$ and $k \in \Winners_{r_2}$.
|
||||
|
||||
So if $p_i$ is a winner he cover the condition $(i, \PROVEtrace(r_1)) \in P$. And we show that if the first condition is never satisfied, the second one will eventually be satisfied. Hence either the first or the second condition will eventually be satisfied, and $p_i$ eventually exits the loop and returns from $\ABbroadcast(m)$.
|
||||
\end{proof}
|
||||
|
||||
\begin{lemma}[Validity]\label{lem:validity}
|
||||
If a correct process $p$ invokes $\ABbroadcast(m)$, then every correct process that invokes a infinitely often times $\ABdeliver()$ eventually delivers $m$.
|
||||
\end{lemma}
|
||||
|
||||
\begin{proof}[Proof]
|
||||
Let $p_i$ a correct process that invokes $\ABbroadcast(m)$ and $p_q$ a correct process that infinitely invokes $\ABdeliver()$. By \Cref{lem:inclusion}, there exist a closed round $r$ and a correct process $j\in\Winners_r$ such that $p_j$ invokes
|
||||
\[
|
||||
\RBcast(S, r, j)\quad\text{with}\quad m\in S.
|
||||
\]
|
||||
|
||||
By \Cref{lem:eventual-closure}, when $p_q$ computes $M_r$ at line C13, $\prop[r][j]$ is non-$\bot$ because $j \in \Winners_r$. By \Cref{lem:unique-proposal}, $p_j$ invokes at most one $\RBcast(S, r, j)$, so $\prop[r][j]$ is uniquely defined. Hence, when $p_q$ computes
|
||||
\[
|
||||
M_r = \bigcup_{k\in\Winners_r} \prop[r][k],
|
||||
\]
|
||||
we have $m \in \prop[r][j] = S$, so $m \in M_r$. By \Cref{lem:convergence}, $M_r$ is invariant so each computation of $M_r$ by any correct process that defines it includes $m$. At each invocation of $\ABdeliver()$ which deliver $m'$, $m'$ is add to $\delivered$ until $M_r \subseteq \delivered$. Once this append we're assured that there exist an invocation of $\ABdeliver()$ which return $m$. Hence $m$ is well delivered.
|
||||
\end{proof}
|
||||
|
||||
\begin{lemma}[No duplication]\label{lem:no-duplication}
|
||||
No correct process delivers the same message more than once.
|
||||
\end{lemma}
|
||||
|
||||
\begin{proof}
|
||||
Let consider two invokations of $\ABdeliver()$ made by the same correct process which returns $m$. Let call these two invocations respectively $\ABdeliver^{(A)}()$ and $\ABdeliver^{(B)}()$.
|
||||
|
||||
When $\ABdeliver^{(A)}()$ occur, by program order and because it reach line C19 to return $m$, the process must have add $m$ to $\delivered$. Hence when $\ABdeliver^{(B)}()$ reach line C14 to extract the next message to deliver, it can't be $m$ because $m \not\in (M_r \setminus \{..., m, ...\})$. So a $\ABdeliver^{(B)}()$ which deliver $m$ can't occur.
|
||||
\end{proof}
|
||||
|
||||
\begin{lemma}[Total order]\label{lem:total-order}
|
||||
For any two messages $m_1$ and $m_2$ delivered by correct processes, if a correct process $p_i$ delivers $m_1$ before $m_2$, then any correct process $p_j$ that delivers both $m_1$ and $m_2$ delivers $m_1$ before $m_2$.
|
||||
\end{lemma}
|
||||
|
||||
\begin{proof}
|
||||
Consider any correct process that delivers both $m_1$ and $m_2$. By \Cref{lem:validity}, there exist closed rounds $r'_1$ and $r'_2$ and correct processes $k_1 \in \Winners_{r'_1}$ and $k_2 \in \Winners_{r'_2}$ such that
|
||||
\[
|
||||
\RBcast(S_1, r'_1, k_1)\quad\text{with}\quad m_1\in S_1,
|
||||
\]
|
||||
\[
|
||||
\RBcast(S_2, r'_2, k_2)\quad\text{with}\quad m_2\in S_2.
|
||||
\]
|
||||
|
||||
Let consider three cases :
|
||||
\begin{itemize}
|
||||
\item \textbf{Case 1:} $r_1 < r_2$. By program order, any correct process must have waited to append in $\delivered$ every messages in $M_{r_1}$ (which contains $m_1$) to increment $\current$ and eventually set $\current = r_2$ to compute $M_{r_2}$ and then invoke the $\ABdeliver()$ that returns $m_2$. Hence, for any correct process that delivers both $m_1$ and $m_2$, it delivers $m_1$ before $m_2$.
|
||||
|
||||
\item \textbf{Case 2:} $r_1 = r_2$. By \Cref{lem:convergence}, any correct process that computes $M_{r_1}$ at line C13 computes the same set of messages $\Messages_{r_1}$. By line C14 the messages are pull in a deterministic order defined by $\ordered(\_)$. Hence, for any correct process that delivers both $m_1$ and $m_2$, it delivers $m_1$ and $m_2$ in the deterministic order defined by $\ordered(\_)$.
|
||||
\end{itemize}
|
||||
|
||||
In all possible cases, any correct process that delivers both $m_1$ and $m_2$ delivers $m_1$ and $m_2$ in the same order.
|
||||
\end{proof}
|
||||
|
||||
\begin{lemma}[Fifo Order]\label{lem:fifo-order}
|
||||
For any two messages $m_1$ and $m_2$ broadcast by the same correct process $p_i$, if $p_i$ invokes $\ABbroadcast(m_1)$ before $\ABbroadcast(m_2)$, then any correct process $p_j$ that delivers both $m_1$ and $m_2$ delivers $m_1$ before $m_2$.
|
||||
\end{lemma}
|
||||
|
||||
\begin{proof}
|
||||
Let take two messages $m_1$ and $m_2$ broadcast by the same correct process $p_i$, with $p_i$ invoking $\ABbroadcast(m_1)$ before $\ABbroadcast(m_2)$. By \Cref{lem:validity}, there exist closed rounds $r_1$ and $r_2$ and correct processes $k_1 \in \Winners_{r_1}$ and $k_2 \in \Winners_{r_2}$ such that
|
||||
\[
|
||||
\RBcast(S_1, r_1, k_1)\quad\text{with}\quad m_1\in S_1,
|
||||
\]
|
||||
\[
|
||||
\RBcast(S_2, r_2, k_2)\quad\text{with}\quad m_2\in S_2.
|
||||
\]
|
||||
|
||||
By program order, $p_i$ must have invoked $\RBcast(S_1, r_1, i)$ before $\RBcast(S_2, r_2, i)$. By \Cref{lem:unique-proposal}, any process invokes at most one $\RBcast(S, r, i)$ per round, hence $r_1 < r_2$. By \Cref{lem:total-order}, any correct process that delivers both $m_1$ and $m_2$ delivers them in a deterministic order.
|
||||
|
||||
In all possible cases, any correct process that delivers both $m_1$ and $m_2$ delivers $m_1$ before $m_2$.
|
||||
\end{proof}
|
||||
|
||||
\begin{theorem}[FIFO-\ARB]
|
||||
Under the assumed \DL synchrony and \RB reliability, the algorithm implements FIFO Atomic Reliable Broadcast.
|
||||
\end{theorem}
|
||||
|
||||
\begin{proof}
|
||||
We show that the algorithm satisfies the properties of FIFO Atomic Reliable Broadcast under the assumed \DL synchrony and \RB reliability.
|
||||
|
||||
First, by \Cref{lem:bcast-termination}, if a correct process invokes \ABbroadcast$(m)$, then it eventually returns from this invocation.
|
||||
Moreover, \Cref{lem:validity} states that if a correct process invokes \ABbroadcast$(m)$, then every correct process that invokes \ABdeliver() infinitely often eventually delivers $m$.
|
||||
This gives the usual Validity property of \ARB.
|
||||
|
||||
Concerning Integrity and No-duplicates, the construction only ever delivers messages that have been obtained from the underlying \RB primitive.
|
||||
By the Integrity property of \RB, every such message was previously \RBcast by some process, so no spurious messages are delivered.
|
||||
In addition, \Cref{lem:no-duplication} states that no correct process delivers the same message more than once.
|
||||
Together, these arguments yield the Integrity and No-duplicates properties required by \ARB.
|
||||
|
||||
For the ordering guarantees, \Cref{lem:total-order} shows that for any two messages $m_1$ and $m_2$ delivered by correct processes, every correct process that delivers both $m_1$ and $m_2$ delivers them in the same order.
|
||||
Hence all correct processes share a common total order on delivered messages.
|
||||
Furthermore, \Cref{lem:fifo-order} states that for any two messages $m_1$ and $m_2$ broadcast by the same correct process, any correct process that delivers both messages delivers $m_1$ before $m_2$ whenever $m_1$ was broadcast before $m_2$.
|
||||
Thus the global total order extends the per-sender FIFO order of \ABbroadcast.
|
||||
|
||||
All the above lemmas are proved under the assumptions that \DL satisfies the required synchrony properties and that the underlying primitive is a Reliable Broadcast (\RB) with Integrity, No-duplicates and Validity.
|
||||
Therefore, under these assumptions, the algorithm satisfies Validity, Integrity/No-duplicates, total order and per-sender FIFO order, and hence implements FIFO Atomic Reliable Broadcast, as claimed.
|
||||
\end{proof}
|
||||
|
||||
\subsection{Reciprocity}
|
||||
% ------------------------------------------------------------------------------
|
||||
|
||||
So far, we assumed the existence of a synchronous DenyList (\DL) object and
|
||||
showed how to upgrade a Reliable Broadcast (\RB) primitive into FIFO Atomic
|
||||
Reliable Broadcast (\ARB). We now briefly argue that, conversely, an \ARB{}
|
||||
primitive is strong enough to implement a synchronous \DL object (ignoring the
|
||||
anonymity property).
|
||||
|
||||
\paragraph{DenyList as a deterministic state machine.}
|
||||
Without anonymity, the \DL specification defines a
|
||||
deterministic abstract object: given a sequence $\Seq$ of operations
|
||||
$\APPEND(x)$, $\PROVE(x)$, and $\READ()$, the resulting sequence of return
|
||||
values and the evolution of the abstract state (set of appended elements,
|
||||
history of operations) are uniquely determined.
|
||||
|
||||
\paragraph{State machine replication over \ARB.}
|
||||
Assume a system that exports a FIFO-\ARB primitive with the guarantees that if a correct process invokes $\ABbroadcast(m)$, then every correct process eventually $\ABdeliver(m)$ and the invocation eventually returns.
|
||||
Following the classical \emph{state machine replication} approach
|
||||
such as described in Schneider~\cite{Schneider90}, we can implement a fault-tolerant service by ensuring the following properties:
|
||||
\begin{quote}
|
||||
\textbf{Agreement.} Every nonfaulty state machine replica receives every request. \\
|
||||
\textbf{Order.} Every nonfaulty state machine replica processes the requests it receives in
|
||||
the same relative order.
|
||||
\end{quote}
|
||||
|
||||
Which are cover by our FIFO-\ARB specification.
|
||||
|
||||
\paragraph{Correctness.}
|
||||
|
||||
|
||||
\begin{theorem}[From \ARB to synchronous \DL]\label{thm:arb-to-dl}
|
||||
In an asynchronous message-passing system with crash failures, assume a
|
||||
FIFO Atomic Reliable Broadcast primitive with Integrity, No-duplicates,
|
||||
Validity, and the liveness of $\ABbroadcast$. Then, ignoring anonymity, there
|
||||
exists an implementation of a synchronous DenyList object that satisfies the
|
||||
Termination, Validity, and Anti-flickering properties.
|
||||
\end{theorem}
|
||||
|
||||
\begin{proof}
|
||||
Because the \DL object is deterministic, all correct processes see the same
|
||||
sequence of operations and compute the same sequence of states and return
|
||||
values. We obtain:
|
||||
|
||||
\begin{itemize}[leftmargin=*]
|
||||
\item \textbf{Termination.} The liveness of \ARB ensures that each
|
||||
$\ABbroadcast$ invocation by a correct process eventually returns, and
|
||||
the corresponding operation is eventually delivered and applied at all
|
||||
correct processes. Thus every $\APPEND$, $\PROVE$, and $\READ$ operation invoked by a correct process
|
||||
eventually returns.
|
||||
\item \textbf{APPEND/PROVE/READ Validity.} The local code that forms
|
||||
\ABbroadcast requests can achieve the same preconditions as in the
|
||||
abstract \DL specification (e.g., $p\in\Pi_M$, $x\in S$ for
|
||||
$\APPEND(x)$). Once an operation is delivered, its effect and return
|
||||
value are exactly those of the sequential \DL specification applied in
|
||||
the common order.
|
||||
\item \textbf{PROVE Anti-Flickering.} In the sequential \DL specification,
|
||||
once an element $x$ has been appended, all subsequent $\PROVE(x)$ are
|
||||
invalid forever. Since all replicas apply operations in the same order,
|
||||
this property holds in every execution of the replicated implementation:
|
||||
after the first linearization point of $\APPEND(x)$, no later
|
||||
$\PROVE(x)$ can return ``valid'' at any correct process.
|
||||
\end{itemize}
|
||||
|
||||
Formally, we can describe the \DL object with the state machine approach for
|
||||
crash-fault, asynchronous message-passing systems with a total order broadcast
|
||||
layer~\cite{Schneider90}.
|
||||
\end{proof}
|
||||
|
||||
\subsubsection{Example executions}
|
||||
|
||||
\begin{figure}[H]
|
||||
\centering
|
||||
\resizebox{0.4\textwidth}{!}{
|
||||
\input{diagrams/nonBFT_behaviour.tex}
|
||||
}
|
||||
\caption{Example execution of the ARB algorithm in a non-BFT setting}
|
||||
\end{figure}
|
||||
|
||||
|
||||
\begin{figure}
|
||||
\centering
|
||||
\resizebox{0.4\textwidth}{!}{
|
||||
\input{diagrams/BFT_behaviour.tex}
|
||||
}
|
||||
\caption{Example execution of the ARB algorithm with a byzantine process}
|
||||
\end{figure}
|
||||
|
||||
\section{BFT-ARB over RB and DL}
|
||||
|
||||
|
||||
\subsection{Model extension}
|
||||
We consider a static set of $n$ processes with known identities, communicating by reliable point-to-point channels, in a complete graph. Messages are uniquely identifiable.
|
||||
|
||||
\paragraph{Synchrony.} The network is asynchronous. Processes may crash or be byzantine; at most $f = \frac{n}{2} - 1$ processes can be faulty.
|
||||
|
||||
\paragraph{Communication.} Processes can exchange through a Reliable Broadcast (\RB) primitive (defined below) which's invoked with the functions \RBcast$(m)$ and \RBreceived$(m)$. There exists a shared object called DenyList (\DL) (defined below) that is interfaced with the functions \APPEND$(x)$, \PROVE$(x)$ and \READ$()$.
|
||||
|
||||
\paragraph{Byzantine behaviour}
|
||||
A process exhibits Byzantine behavior if it deviates arbitrarily from the specified algorithm. This includes, but is not limited to, the following actions:
|
||||
|
||||
\begin{itemize}
|
||||
\item Invoking primitives (\RBcast, \APPEND, \PROVE, etc.) with invalid or maliciously crafted inputs.
|
||||
\item Colluding with other Byzantine processes to manipulate the system's state or violate its guarantees.
|
||||
\item Delaying or accelerating message delivery to specific nodes to disrupt the expected timing of operations.
|
||||
\item Withholding messages or responses to create inconsistencies in the system's state.
|
||||
\end{itemize}
|
||||
|
||||
Byzantine processes are constrained by the following:
|
||||
\begin{itemize}
|
||||
\item They cannot forge valid cryptographic signatures or threshold shares without the corresponding private keys.
|
||||
\item They cannot violate the termination, validity, or anti-flickering properties of the \DL{} object.
|
||||
\item They cannot break the integrity, no-duplicates, or validity properties of the \RB{} primitive.
|
||||
\end{itemize}
|
||||
|
||||
\paragraph{Notation.} Let $\Pi$ be the finite set of process identifiers and let $n \triangleq |\Pi|$. Two authorization subsets are $M \subseteq \Pi$ (processes allowed to issue \APPEND) and $V \subseteq \Pi$ (processes allowed to issue \PROVE). Indices $i,j \in \Pi$ refer to processes, and $p_i$ denotes the process with identifier $i$. Let $\mathcal{M}$ denote the universe of uniquely identifiable messages, with $m \in \mathcal{M}$. Let $\mathcal{R} \subseteq \mathbb{N}$ be the set of round identifiers; we write $r \in \mathcal{R}$ for a round. We use the precedence relation $\prec$ for the \DL{} linearization: $x \prec y$ means that operation $x$ appears strictly before $y$ in the linearized history of \DL. For any finite set $A \subseteq \mathcal{M}$, \ordered$(A)$ returns a deterministic total order over $A$ (e.g., lexicographic order on $(\textit{senderId},\textit{messageId})$ or on message hashes). For any round $r \in \mathcal{R}$, define $\Winners_r \triangleq \{\, j \in \Pi \mid (j,\PROVEtrace(r)) \prec \APPEND(r) \,\}$, i.e., the set of processes whose $\PROVE(r)$ appears before the first $\APPEND(r)$ in the \DL{} linearization.
|
||||
We denoted by $\PROVE^{(j)}(r)$ or $\APPEND^{(j)}(r)$ the operation $\PROVE(r)$ or $\APPEND(r)$ invoked by process $j$.
|
||||
|
||||
% ------------------------------------------------------------------------------
|
||||
\subsection{Primitives}
|
||||
|
||||
\subsubsection{BFT DenyList}
|
||||
We consider a \DL object that satisfies the following properties despite the presence of up to $f$ byzantine processes:
|
||||
\begin{itemize}
|
||||
\item \textbf{Termination.} Every operation $\APPEND(x)$, $\PROVE(x)$, and $\READ()$ invoked by a correct process eventually returns.
|
||||
\item \textbf{APPEND/PROVE/READ Validity.} The preconditions for invoking $\APPEND(x)$, $\PROVE(x)$, and $\READ()$ are respected (cf. \#2.2). The return values of these operations conform to the sequential specification of the \DL object.
|
||||
\item \textbf{APPEND Justification.} For any element $x$, if an operation $\APPEND(x)$ invoked by a correct process completes successfully, then there exists at least one valid $\PROVE(x)$ operation that that precedes this $\APPEND(x)$ in the \DL linearization.
|
||||
\item \textbf{PROVE Anti-Flickering.} Once an element $x$ has been appended to the \DL by any process, all subsequent invocations of $\PROVE(x)$ by any process return ``invalid''.
|
||||
\end{itemize}
|
||||
|
||||
\subsubsection{t-out-of-n Threshold Random Number Generator}
|
||||
|
||||
We consider a function that with t out of n shares any process can reconstruct a deterministic random number. The function is defined as follows:
|
||||
|
||||
\begin{itemize}
|
||||
\item \textbf{$t$-reconstruction.} Given any subset $S$ of at least $t$ valid shares derived from the same value $r$, there exists a unique value $\sigma$ consistent with all shares in~$S$, and $\sigma$ can be efficiently reconstructed from~$S$.
|
||||
\item \textbf{$(t-1)$-non-reconstructibility.} Given any subset $S$ of at most $t-1$ valid shares derived from the same value $r$, there exist two distinct values $\sigma$ and $\sigma'$ that are both consistent with all shares in~$S$. In particular, no algorithm that only sees the shares in $S$ can always distinguish whether the underlying value is $\sigma$ or $\sigma'$. \item \textbf{Per-process non-equivocation.} For any process $p$ and value $r$, there is at most one valid share that $p$ can derive from $r$. Formally, if $\sigma$ and $\sigma'$ are two valid shares output by process $p$ from the same value $r$, then $\sigma = \sigma'$. In particular, a single process cannot emit two different valid shares for the same underlying value~$r$.
|
||||
\end{itemize}
|
||||
|
||||
\paragraph{Interface.}
|
||||
For algorithmic purposes, we model the $t$-out-of-$n$ threshold random number generator
|
||||
as providing the following interface to each process $p \in \Pi$.
|
||||
|
||||
\begin{itemize}
|
||||
\item{$\mathsf{SHARE}_{p_i}(r)$:} On input a value $r$, run locally by process $p_i$, returns a valid share $\sigma_r^i$. By per-process share uniqueness, for any fixed $p_i$ and $r$ the value $\sigma_r^i$ is uniquely determined.
|
||||
|
||||
\item{$\mathsf{COMBINE}(S)$:} On a set $S$ of at least $t$ pairs $(p_i,\sigma_r^i)$, returns the reconstructed value $\sigma_r$. By $t$-reconstruction, this value is well defined and independent of the particular set $S$ of valid shares of size at least $t$.
|
||||
|
||||
\item{$\mathsf{VERIFY}(r,\sigma_{r'})$:} On input a value $r$ and a candidate value $\sigma_{r'}$, returns \textsf{true} if and only if there exists a set $S$ of at least $t$ valid shares for $r$ such that $\mathsf{Combine}(r,S) = \sigma_{r'}$, and \textsf{false} otherwise. We say that $\sigma_{r'}$ is \emph{valid for $r$} if $\mathsf{Verify}(r,\sigma_{r'})=\textsf{true}$.
|
||||
\end{itemize}
|
||||
|
||||
|
||||
\subsection{Algorithm}
|
||||
|
||||
\subsubsection{Variables}
|
||||
Each process $p_i$ maintains the following local variables:
|
||||
\begin{algorithmic}
|
||||
\State $\current \gets 0$
|
||||
\State $\received \gets \emptyset$
|
||||
\State $\delivered \gets \emptyset$
|
||||
\State $\prop[r][j] \gets \bot, \forall r, j$
|
||||
\State $X_r \gets \bot, \forall r$
|
||||
\State $\resolved[r] \gets \bot, \forall r$
|
||||
\end{algorithmic}
|
||||
|
||||
\renewcommand{\algletter}{D}
|
||||
\begin{algorithm}[H]
|
||||
\caption{\ABbroadcast}\label{alg:ab-cast}
|
||||
\begin{algorithmic}[1]
|
||||
\Function{ABcast}{$m$}
|
||||
\State $S \gets (\received \setminus \delivered) \cup \{m\}$
|
||||
\State $\RBcast(prop, S, r, i)$
|
||||
\State \textbf{wait until} $|X_r| \geq f+1$
|
||||
\State $\sigma_r \gets \COMBINE(X_r)$
|
||||
\State $\PROVE(\sigma_r); \APPEND(\sigma_r);$
|
||||
\State $\RBcast(submit, S, \sigma_r, r, i)$
|
||||
\EndFunction
|
||||
\end{algorithmic}
|
||||
\end{algorithm}
|
||||
|
||||
\renewcommand{\algletter}{E}
|
||||
\begin{algorithm}[H]
|
||||
\caption{\ABdeliver}\label{alg:ab-deliver}
|
||||
\begin{algorithmic}[1]
|
||||
\Function{$\ABdeliver$}{}
|
||||
\State $r \gets \current; \sigma_r \gets \resolved[r];$
|
||||
\If{$\sigma_r == \bot$}
|
||||
\State \Return $\bot$
|
||||
\EndIf
|
||||
\State $P \gets \READ()$
|
||||
\If{$\forall j : (j,prove(\sigma_r)) \not\in P$}
|
||||
\State \Return $\bot$
|
||||
\EndIf
|
||||
\State $\APPEND(\sigma_r); P \gets \READ();$
|
||||
\State $W_r \gets \{j : (j, \PROVEtrace(\sigma_r)) \in P\}$
|
||||
\If{$\exists j \in W_r : \prop[r][j] == \bot$}
|
||||
\State \Return $\bot$
|
||||
\EndIf
|
||||
\State $M_r \gets \bigcup_{j \in W_r} \prop[r][j];$
|
||||
\State $m \gets \ordered(M_r)[0]$
|
||||
\State $\delivered \gets \delivered \cup \{m\};$
|
||||
\If{$M_r \setminus \delivered = \emptyset$}
|
||||
\State $\current \gets \current + 1;$
|
||||
\EndIf
|
||||
\State \Return $m$
|
||||
\EndFunction
|
||||
\end{algorithmic}
|
||||
|
||||
\end{algorithm}
|
||||
|
||||
\renewcommand{\algletter}{F}
|
||||
\begin{algorithm}[H]
|
||||
\caption{RBreceived handler}\label{alg:rb-handler}
|
||||
\begin{algorithmic}[1]
|
||||
\Function{RBrcvd}{$prop, S_j, r_j, j$}
|
||||
\If{$r_j \geq r$}
|
||||
\State $\prop[r_j][j] = S_j$
|
||||
\State $\sigma^i_{r_j} \gets \SHARE(r_j)$
|
||||
\State $send_j(r, \sigma^i_{r_j})$
|
||||
\EndIf
|
||||
\EndFunction
|
||||
\end{algorithmic}
|
||||
\end{algorithm}
|
||||
|
||||
|
||||
\renewcommand{\algletter}{G}
|
||||
\begin{algorithm}[H]
|
||||
\caption{RBreceived handler}\label{alg:rb-handler-2}
|
||||
\begin{algorithmic}[1]
|
||||
\Function{RBrcvd}{$submit, S_j, \sigma_{r_j}, r_j, j$}
|
||||
\If{$\VERIFY(r_j, \sigma_{r_j})$}
|
||||
\State $\resolved[r_j] \gets \sigma_{r_j}$
|
||||
\EndIf
|
||||
\EndFunction
|
||||
\end{algorithmic}
|
||||
\end{algorithm}
|
||||
|
||||
|
||||
\renewcommand{\algletter}{H}
|
||||
\begin{algorithm}[H]
|
||||
\caption{Share received handler}\label{alg:share-handler}
|
||||
\begin{algorithmic}[1]
|
||||
\Function{received}{$r_j, \sigma^j_{r_j}, j$}
|
||||
\If{$r_j == r$}
|
||||
\State $X_r \gets X_r \cup \sigma^j_{r}$
|
||||
\EndIf
|
||||
\EndFunction
|
||||
\end{algorithmic}
|
||||
\end{algorithm}
|
||||
|
||||
\subsection{Example execution}
|
||||
|
||||
\begin{figure}[H]
|
||||
\centering
|
||||
\input{diagrams/classic_seq.tex}
|
||||
\caption{Expected Executions of P1 willing to send a message at round r}
|
||||
\end{figure}
|
||||
|
||||
\section{Implementation of BFT-DenyList and Threshold Cryptography}
|
||||
|
||||
\subsection{DenyList}
|
||||
|
||||
\paragraph{BFT-DenyList}
|
||||
In our algorithm we use multiple DenyList as follows:
|
||||
|
||||
\begin{itemize}
|
||||
\item Let $\mathcal{DL} = \{DL_1, \dots, DL_k\}$ be the set of DenyList used by the algorithm.
|
||||
\item We set $k = \binom{n}{f}$.
|
||||
\item For each $i \in \{1,\dots,k\}$, let $M_i$ be the set of moderators associated with $DL_i$ according to the DenyList definition, so that $|M_i| = n-f$.
|
||||
\item Let $\mathcal{M} = \{M_1, \dots, M_k\}$. We require that the $M_i$ are pairwise distinct:
|
||||
\[
|
||||
\forall i,j \in \{1,\dots,k\},\ i \neq j \implies M_i \neq M_j.
|
||||
\]
|
||||
\end{itemize}
|
||||
|
||||
|
||||
\begin{lemma}
|
||||
$\exists M_i \in M : \forall p \in M_i$ $p$ is correct.
|
||||
\end{lemma}
|
||||
|
||||
\begin{proof}
|
||||
Let consider the set $F$ of faulty processes, with $|F| = f$. We can construct the set $M_i = \Pi \setminus F$ such that $|M_i| = n - |F| = n - f$. By construction, $\forall p \in M_i$ $p$ is correct.
|
||||
\end{proof}
|
||||
|
||||
\begin{lemma}
|
||||
$\forall M_i \in M, \exists p \in M_i$ such that $p$ is correct.
|
||||
\end{lemma}
|
||||
|
||||
\begin{proof}
|
||||
$\forall i \in \{1, \dots, k\}, |M_i| = n-f$ with $n \geq 2f+1$. We can say that $|M_i| \geq 2f+1-f = f+1 > f$
|
||||
\end{proof}
|
||||
|
||||
Each process can invoke the following functions :
|
||||
|
||||
\begin{itemize}
|
||||
\item $\READ' : () \rightarrow \mathcal{L}(\mathbb{R} \times \PROVEtrace(\mathbb{R}))$
|
||||
\item $\APPEND' : \mathbb{R} \rightarrow ()$
|
||||
\item $\PROVE' : \mathbb{R} \rightarrow \{0, 1\}$
|
||||
\end{itemize}
|
||||
|
||||
Such that :
|
||||
|
||||
\begin{algorithm}[H]
|
||||
\caption{$\READ'() \rightarrow \mathcal{L}(\mathbb{R} \times \PROVEtrace(\mathbb{R}))$}
|
||||
\begin{algorithmic}
|
||||
\Function{READ'}{}
|
||||
\State $j \gets$ the process invoking $\READ'()$
|
||||
\State $res \gets \emptyset$
|
||||
\ForAll{$i \in \{1, \dots, k\}$}
|
||||
\State $res \gets res \cup DL_i.\READ()$
|
||||
\EndFor
|
||||
\State \Return $res$
|
||||
\EndFunction
|
||||
\end{algorithmic}
|
||||
\end{algorithm}
|
||||
|
||||
\begin{algorithm}[H]
|
||||
\caption{$\APPEND'(\sigma) \rightarrow ()$}
|
||||
\begin{algorithmic}
|
||||
\Function{APPEND'}{$\sigma$}
|
||||
\State $j \gets$ the process invoking $\APPEND'(\sigma)$
|
||||
\ForAll{$M_i \in \{M_k \in M : j \in M_k\}$}
|
||||
\State $DL_i.\APPEND(\sigma)$
|
||||
\EndFor
|
||||
\EndFunction
|
||||
\end{algorithmic}
|
||||
\end{algorithm}
|
||||
|
||||
\begin{algorithm}[H]
|
||||
\caption{$\PROVE'(\sigma) \rightarrow \{0, 1\}$}
|
||||
\begin{algorithmic}
|
||||
\Function{PROVE'}{$\sigma$}
|
||||
\State $j \gets$ the process invoking $\PROVE'(\sigma)$
|
||||
\State $flag \gets false$
|
||||
\ForAll{$i \in \{1, \dots, k\}$}
|
||||
\State $flag \gets flag$ OR $DL_i.\PROVE(\sigma)$
|
||||
\EndFor
|
||||
\State \Return $flag$
|
||||
\EndFunction
|
||||
\end{algorithmic}
|
||||
\end{algorithm}
|
||||
|
||||
\subsection{Threshold Cryptography}
|
||||
|
||||
We are using the Boneh-Lynn-Shacham scheme as cryptography primitive to our threshold signature scheme.
|
||||
With :
|
||||
|
||||
\begin{itemize}
|
||||
\item $G : \mathbb{R} \rightarrow \mathbb{R} \times \mathbb{R} $
|
||||
\item $S : \mathbb{R} \times \mathcal{R} \rightarrow \mathbb{R} $
|
||||
\item $V : \mathbb{R} \times \mathcal{R} \times \mathbb{R} \rightarrow \{0, 1\} $
|
||||
\end{itemize}
|
||||
|
||||
Such that :
|
||||
|
||||
\begin{itemize}
|
||||
\item $G(x) \rightarrow (pk, sk)$ : where $x$ is a random value such that $\nexists x_1, x_2: x_1 \neq x_2, G(x_1) = G(x_2)$
|
||||
\item $S(sk, m) \rightarrow \sigma_m$
|
||||
\item $V(pk, m_1, \sigma_{m_2}) \rightarrow k$ : with $k = 1$ iff $m_1 == m_2$ and $\exists x \in \mathbb{R}$ such that $G(x) \rightarrow (pk, sk)$; otherwise $k = 0$
|
||||
\end{itemize}
|
||||
|
||||
\paragraph{threshold Scheme}
|
||||
|
||||
In our algorithm we are only using the following functions :
|
||||
|
||||
\begin{itemize}
|
||||
\item $G' : \mathbb{R} \times \mathbb{N} \times \mathbb{N} \rightarrow \mathbb{R} \times (\mathbb{R} \times \mathbb{R})^n$ : with $n \triangleq |\Pi|$
|
||||
\item $S' : \mathbb{R} \times \mathcal{R} \rightarrow \mathbb{R}$
|
||||
\item $C' : \mathbb{R}^n \times \mathcal{R} \times \mathbb{R} \times \mathbb{R}^t \rightarrow \{\mathbb{R}, \bot\}$ : with $t \leq n$
|
||||
\item $V' : \mathbb{R} \times \mathcal{R} \times \mathbb{R} \rightarrow \{0, 1\}$
|
||||
\end{itemize}
|
||||
|
||||
Such that :
|
||||
|
||||
\begin{itemize}
|
||||
\item $G'(x, n, t) \rightarrow (pk, pk_1, sk_1, \dots, pk_n, sk_n)$ : let define $pkc = {pk_1, \dots, pk_n}$
|
||||
\item $S'(sk_i, m) \rightarrow \sigma_m^i$
|
||||
\item $C'(pkc, m_1, J, \{\sigma_{m_2}^j\}_{j \in J}) \rightarrow \sigma$ : with $J \subseteq \Pi$; and $\sigma = \sigma_{m_1}$ iff $|J| \geq t, \forall j \in J: V(pk_j, m_1, \sigma_{m_2}^j) == 1$; otherwise $\sigma = \bot$.
|
||||
\item $V'(pk, m_1, \sigma_{m_2}) \rightarrow V(pk, m_1, \sigma_{m_2})$
|
||||
\end{itemize}
|
||||
|
||||
|
||||
\bibliographystyle{plain}
|
||||
\begin{thebibliography}{9}
|
||||
% (left intentionally blank)
|
||||
\bibitem{Schneider90}
|
||||
Fred B.~Schneider.
|
||||
\newblock Implementing fault-tolerant services using the state machine
|
||||
approach: a tutorial.
|
||||
\newblock {\em ACM Computing Surveys}, 22(4):299--319, 1990.
|
||||
\end{thebibliography}
|
||||
|
||||
\end{document}
|
||||
647
docs/presentations/autre/capitoledulibre/main.tex
Normal file
647
docs/presentations/autre/capitoledulibre/main.tex
Normal file
@@ -0,0 +1,647 @@
|
||||
\documentclass[aspectratio=43]{beamer}
|
||||
|
||||
\usetheme{Berlin}
|
||||
%\usefonttheme{professionalfonts}
|
||||
|
||||
\usepackage[utf8]{inputenc}
|
||||
\usepackage[T1]{fontenc}
|
||||
\usepackage[french]{babel}
|
||||
\usepackage{mathtools,amssymb}
|
||||
\usepackage{microtype}
|
||||
\usepackage{pgfgantt}
|
||||
\usepackage{adjustbox}
|
||||
\usetikzlibrary{positioning}
|
||||
% \usepackage{qrcode}
|
||||
|
||||
\title[]{Systèmes centralisés, fédérés et pair à pair : vu par la théorie des graphs}
|
||||
\author{Amaury JOLY}
|
||||
\institute[AMU - Skeptikon - Parsec]{Aix-Marseille université, Skeptikon, Parsec.cloud}
|
||||
|
||||
\begin{document}
|
||||
|
||||
\begin{frame}{}
|
||||
\titlepage
|
||||
\end{frame}
|
||||
|
||||
\note{test}
|
||||
\begin{frame}{Qui suis-je}
|
||||
\begin{itemize}
|
||||
\item \textbf{Doctorant} en informatique distribué
|
||||
\begin{itemize}
|
||||
\item \textit{Étude de la cohérence des données dans les systèmes distribués}
|
||||
\end{itemize}
|
||||
\item Ingénieur R\&D chez \textbf{Parsec.cloud}
|
||||
\item Membre de l'association \textbf{Skeptikon}
|
||||
\end{itemize}
|
||||
\end{frame}
|
||||
|
||||
\section{Introduction a l'algorithmie distribué}
|
||||
|
||||
\begin{frame}{Un système distribué c'est quoi}
|
||||
\begin{quote}
|
||||
On définit un système comme un ensemble de processus. On considère qu'un système est dit distribué lorsque plusieurs processus distincts travail à la réalisation d'une tâche commune.
|
||||
\end{quote}
|
||||
|
||||
\vspace{1em}
|
||||
Par exemple la rédaction de documents collaboratifs (ex: cryptpad, grist, colabora), le chat (ex: matrix, signal, XMPP, IRC, IMAP), le partage d'annuaire de vidéo (ex: PeerTube), la résolution des noms de domaines (DNS).
|
||||
\end{frame}
|
||||
|
||||
\begin{frame}{Définir un système distribué: le modèle de communication}
|
||||
La recherche traite globalement deux modèles de communication
|
||||
\begin{itemize}
|
||||
\item Shared-Register
|
||||
|
||||
\pause
|
||||
|
||||
\resizebox{0.5\linewidth}{!}{%
|
||||
\begin{tikzpicture}[
|
||||
node distance=1.8cm and 2.4cm,
|
||||
process/.style={draw, circle, thick, minimum size=1cm, font=\small},
|
||||
mem/.style={draw, thick, rounded corners, minimum width=5cm, minimum height=1.8cm, fill=gray!10},
|
||||
reg/.style={draw, thin, minimum width=1.1cm, minimum height=0.7cm, font=\scriptsize},
|
||||
arrow/.style={-Latex, thick}
|
||||
]
|
||||
|
||||
% --- Mémoire partagée (abstraction) ---
|
||||
\node[mem] (mem) {};
|
||||
|
||||
% Quelques registres logiques dans la mémoire
|
||||
\node[reg] (r1) at ($(mem.west)!0.2!(mem.east)$) {$R_1$};
|
||||
\node[reg] (r2) at ($(mem.west)!0.5!(mem.east)$) {$R_\dots$};
|
||||
\node[reg] (r3) at ($(mem.west)!0.8!(mem.east)$) {$R_n$};
|
||||
|
||||
\node[font=\scriptsize, below=0.15cm of mem]{Tous les processus voient la même mémoire logique};
|
||||
|
||||
% --- Processus concurrents ---
|
||||
\node[process] (p1) [left=of p2] {$P_1$};
|
||||
\node[process] (p2) [above=1.8cm of mem] {$P_\dots$};
|
||||
\node[process] (p3) [right=of p2] {$P_n$};
|
||||
|
||||
% --- Accès mémoire : opérations read/write ---
|
||||
\draw[arrow] (p1.south) -- (r1.north);
|
||||
\draw[arrow] (p1.south) -- (r2.north);
|
||||
\draw[arrow] (p1.south) -- (r3.north);
|
||||
\draw[arrow] (p2.south) -- (r2.north);
|
||||
\draw[arrow] (p2.south) -- (r1.north);
|
||||
\draw[arrow] (p2.south) -- (r3.north);
|
||||
\draw[arrow] (p3.south) -- (r1.north);
|
||||
\draw[arrow] (p3.south) -- (r2.north);
|
||||
\draw[arrow] (p3.south) -- (r3.north);
|
||||
|
||||
\end{tikzpicture}
|
||||
}
|
||||
\end{itemize}
|
||||
\end{frame}
|
||||
|
||||
\begin{frame}{Définir un système distribué: le modèle de communication}
|
||||
\begin{itemize}
|
||||
\item Shared-Register
|
||||
\item \textbf{Message-passing}
|
||||
|
||||
\centering
|
||||
\resizebox{0.2\linewidth}{!}{%
|
||||
\begin{tikzpicture}[
|
||||
node distance=2.2cm and 1.4cm,
|
||||
process/.style={draw, circle, thick, minimum size=1cm, font=\small},
|
||||
state/.style={draw, thin, rounded corners, minimum width=1.6cm,
|
||||
minimum height=0.7cm, font=\scriptsize, fill=gray!10},
|
||||
msg/.style={-Latex, thick},
|
||||
netbox/.style={draw, dashed, rounded corners, thick, inner sep=0.4cm}
|
||||
]
|
||||
|
||||
% --- Processus avec état local ---
|
||||
\node[process] (p1) {$P_1$};
|
||||
\node[process] (p2) [above=of p1] {$P_2$};
|
||||
\node[process] (p3) [right=of p1] {$P_3$};
|
||||
|
||||
\node[state] (s1) [below=0.7cm of p1] {état local $S_1$};
|
||||
\node[state] (s2) [above=0.7cm of p2] {état local $S_2$};
|
||||
\node[state] (s3) [below=0.7cm of p3] {état local $S_3$};
|
||||
|
||||
% --- Messages entre processus ---
|
||||
\draw[msg] (p1) to[bend left=15] node[above, font=\scriptsize]{message $m_1$} (p2);
|
||||
\draw[msg] (p2) to[bend left=15] node[above, font=\scriptsize]{message $m_2$} (p3);
|
||||
\draw[msg] (p3) to[bend left=15] node[above, font=\scriptsize]{message $m_3$} (p1);
|
||||
|
||||
\end{tikzpicture}
|
||||
}%
|
||||
|
||||
\end{itemize}
|
||||
\end{frame}
|
||||
|
||||
\begin{frame}{Définir un système distribué: le modèle de synchronisme}
|
||||
La recherche distingue deux grandes hypothèses de synchronisme :
|
||||
\begin{itemize}
|
||||
\item \textbf{Synchrone} (temps découpé en tours / rounds)
|
||||
|
||||
\pause
|
||||
|
||||
\centering
|
||||
\resizebox{0.7\linewidth}{!}{%
|
||||
\begin{tikzpicture}[
|
||||
x=1.4cm, y=0.9cm,
|
||||
msg/.style={-Latex, thick},
|
||||
proc/.style={font=\small},
|
||||
roundline/.style={thick},
|
||||
sep/.style={dashed, gray}
|
||||
]
|
||||
|
||||
% --- Processus ---
|
||||
\node[proc, anchor=east] at (0,0) {$P_1$};
|
||||
\node[proc, anchor=east] at (0,-1) {$P_2$};
|
||||
\node[proc, anchor=east] at (0,-2) {$P_3$};
|
||||
|
||||
% --- Lignes de temps pour chaque processus ---
|
||||
\foreach \y in {0,-1,-2}{
|
||||
\draw[roundline] (0.2,\y) -- (3.0,\y);
|
||||
}
|
||||
|
||||
% --- Séparation en tours synchrones ---
|
||||
\foreach \x/\t in {0.8/1,1.6/2,2.4/3}{
|
||||
\draw[sep] (\x,0.3) -- (\x,-2.3);
|
||||
\node[font=\scriptsize] at (\x+0.35,0.35) {tour \t};
|
||||
}
|
||||
|
||||
% --- Messages livrés dans le tour suivant ---
|
||||
\draw[msg] (0.35,0) .. controls (0.55,-0.2) .. (0.75,-1); % P1 -> P2 dans le tour 1
|
||||
\draw[msg] (1.0,-1) .. controls (1.2,-1.8) .. (1.4,-2); % P2 -> P3 dans le tour 2
|
||||
\draw[msg] (1.7,-2) .. controls (1.9,-1.2) .. (2.1,0); % P3 -> P1 dans le tour 3
|
||||
|
||||
% --- Légende ---
|
||||
\node[font=\scriptsize, align=center] at (1.6,-2.6){Bornes connues sur les durées de calcul et de communication,\\
|
||||
tous les processus avancent par tours synchrones};
|
||||
|
||||
\end{tikzpicture}
|
||||
}
|
||||
\end{itemize}
|
||||
\end{frame}
|
||||
|
||||
\begin{frame}{Définir un système distribué: les modèles de synchronisme}
|
||||
\begin{itemize}
|
||||
\item Synchrone
|
||||
\item \textbf{Asynchrone}
|
||||
\end{itemize}
|
||||
|
||||
\pause
|
||||
|
||||
\centering
|
||||
\resizebox{0.7\linewidth}{!}{%
|
||||
\begin{tikzpicture}[
|
||||
x=1.4cm, y=0.9cm,
|
||||
msg/.style={-Latex, thick},
|
||||
proc/.style={font=\small},
|
||||
opbox/.style={draw, rounded corners, thick}
|
||||
]
|
||||
|
||||
% --- Processus ---
|
||||
\node[proc, anchor=east] at (0,0) {$P_1$};
|
||||
\node[proc, anchor=east] at (0,-1) {$P_2$};
|
||||
\node[proc, anchor=east] at (0,-2) {$P_3$};
|
||||
|
||||
\pause
|
||||
|
||||
% --- Lignes de temps pour chaque processus ---
|
||||
\foreach \y in {0,-1,-2}{
|
||||
\draw[thick] (0.2,\y) -- (3.0,\y);
|
||||
}
|
||||
|
||||
\pause
|
||||
|
||||
% --- "Ticks" locaux non alignés (horloges indépendantes) ---
|
||||
\foreach \x in {0.5,1.2,2.3}{
|
||||
\draw (\x,0.1) -- (\x,-0.1); % P1
|
||||
}
|
||||
\foreach \x in {0.4,1.5,2.6}{
|
||||
\draw (\x,-0.9) -- (\x,-1.1); % P2
|
||||
}
|
||||
\foreach \x in {0.7,1.8,2.1}{
|
||||
\draw (\x,-1.9) -- (\x,-2.1); % P3
|
||||
}
|
||||
|
||||
\pause
|
||||
|
||||
% --- Messages ---
|
||||
\draw[msg] (0.35,0) -- (0.8,-1);
|
||||
\draw[msg] (0.35,0) -- (2.5,-2);
|
||||
|
||||
\pause
|
||||
|
||||
\draw[msg] (1.0,-1) -- (1.8,0);
|
||||
\draw[msg] (1.0,-1) -- (2.2,-2);
|
||||
|
||||
\pause
|
||||
|
||||
\draw[msg] (1.4,-2) -- (2.1,0);
|
||||
\draw[msg] (1.4,-2) -- (2.5,-1);
|
||||
|
||||
\pause
|
||||
% op_1 : diffusion de P1
|
||||
% commence à x = 0.35 (émission) et finit à x = 2.5 (dernier reçu : P3)
|
||||
\draw[opbox,densely dotted] (0.35,0.25) rectangle (2.5,-0.25);
|
||||
% ligne de vie de P1 en pointillé dans l'intervalle de op_1
|
||||
\draw[thick,densely dotted] (0.35,0) -- (2.5,0);
|
||||
\node[font=\scriptsize, anchor=west] at (2.55,-0.2) {$op_1$};
|
||||
|
||||
\pause
|
||||
% op_2 : diffusion de P2
|
||||
% commence à x = 1.0 (émission) et finit à x = 2.2 (dernier reçu : P3)
|
||||
\draw[opbox,densely dotted] (1.0,-0.75) rectangle (2.2,-1.25);
|
||||
% ligne de vie de P2 en pointillé dans l'intervalle de op_2
|
||||
\draw[thick,densely dotted] (1.0,-1) -- (2.2,-1);
|
||||
\node[font=\scriptsize, anchor=west] at (2.25,-1.2) {$op_2$};
|
||||
|
||||
\pause
|
||||
% op_3 : diffusion de P3
|
||||
% commence à x = 1.4 (émission) et finit à x = 2.5 (dernier reçu : P2)
|
||||
\draw[opbox,densely dotted] (1.4,-1.75) rectangle (2.5,-2.25);
|
||||
% ligne de vie de P3 en pointillé dans l'intervalle de op_3
|
||||
\draw[thick,densely dotted] (1.4,-2) -- (2.5,-2);
|
||||
\node[font=\scriptsize, anchor=west] at (2.55,-2.2) {$op_3$};
|
||||
|
||||
\end{tikzpicture}
|
||||
}
|
||||
\end{frame}
|
||||
|
||||
\begin{frame}{Définir un système distribué: les modèles de synchronisme}
|
||||
\begin{itemize}
|
||||
\item Synchrone
|
||||
\item \textbf{Asynchrone}
|
||||
\end{itemize}
|
||||
|
||||
\centering
|
||||
\resizebox{0.7\linewidth}{!}{%
|
||||
\begin{tikzpicture}[
|
||||
x=1.4cm, y=0.9cm,
|
||||
msg/.style={-Latex, thick},
|
||||
proc/.style={font=\small},
|
||||
opbox/.style={draw, rounded corners, thick}
|
||||
]
|
||||
|
||||
% --- Processus ---
|
||||
\node[proc, anchor=east] at (0,0) {$P_1$};
|
||||
\node[proc, anchor=east] at (0,-1) {$P_2$};
|
||||
\node[proc, anchor=east] at (0,-2) {$P_3$};
|
||||
|
||||
% --- Lignes de temps pour chaque processus ---
|
||||
\foreach \y in {0,-1,-2}{
|
||||
\draw[thick] (0.2,\y) -- (3.0,\y);
|
||||
}
|
||||
|
||||
% op_1 : diffusion de P1
|
||||
% commence à x = 0.35 (émission) et finit à x = 2.5 (dernier reçu : P3)
|
||||
\draw[opbox,densely dotted] (0.35,0.25) rectangle (2.5,-0.25);
|
||||
% ligne de vie de P1 en pointillé dans l'intervalle de op_1
|
||||
\draw[thick,densely dotted] (0.35,0) -- (2.5,0);
|
||||
\node[font=\scriptsize, anchor=west] at (2.55,-0.2) {$op_1$};
|
||||
|
||||
% op_2 : diffusion de P2
|
||||
% commence à x = 1.0 (émission) et finit à x = 2.2 (dernier reçu : P3)
|
||||
\draw[opbox,densely dotted] (1.0,-0.75) rectangle (2.2,-1.25);
|
||||
% ligne de vie de P2 en pointillé dans l'intervalle de op_2
|
||||
\draw[thick,densely dotted] (1.0,-1) -- (2.2,-1);
|
||||
\node[font=\scriptsize, anchor=west] at (2.25,-1.2) {$op_2$};
|
||||
|
||||
% op_3 : diffusion de P3
|
||||
% commence à x = 1.4 (émission) et finit à x = 2.5 (dernier reçu : P2)
|
||||
\draw[opbox,densely dotted] (1.4,-1.75) rectangle (2.5,-2.25);
|
||||
% ligne de vie de P3 en pointillé dans l'intervalle de op_3
|
||||
\draw[thick,densely dotted] (1.4,-2) -- (2.5,-2);
|
||||
\node[font=\scriptsize, anchor=west] at (2.55,-2.2) {$op_3$};
|
||||
|
||||
\end{tikzpicture}
|
||||
}
|
||||
\end{frame}
|
||||
|
||||
\section{Les différentes architectures de systèmes}
|
||||
|
||||
\subsection{Les systèmes centralisés}
|
||||
|
||||
\begin{frame}{Système centralisé}
|
||||
\begin{itemize}
|
||||
\item Architecture client--serveur
|
||||
\item Un seul serveur central maintient l'état
|
||||
\end{itemize}
|
||||
|
||||
\vspace{0.3cm}
|
||||
\centering
|
||||
\resizebox{0.6\linewidth}{!}{%
|
||||
\begin{tikzpicture}[
|
||||
node distance=1.8cm and 2.5cm,
|
||||
client/.style={draw, circle, thick, minimum size=1cm, font=\small},
|
||||
server/.style={draw, thick, rounded corners, minimum width=2.8cm,
|
||||
minimum height=1.2cm, font=\small, fill=gray!10},
|
||||
msg/.style={-Latex, thick},
|
||||
resp/.style={Latex-, thick, dashed}
|
||||
]
|
||||
|
||||
% --- Serveur central ---
|
||||
\node[server] (srv) at (0,0) {Serveur central};
|
||||
|
||||
% --- Clients ---
|
||||
\node[client] (c1) at (-3,0) {$P_1$};
|
||||
\node[client] (c2) at (0,-2) {$P_2$};
|
||||
\node[client] (c3) at (3,-0) {$P_3$};
|
||||
|
||||
% --- Messages requête / réponse ---
|
||||
\draw[msg] (c1.east) -- (srv.west);
|
||||
\draw[msg] (c2.north) -- (srv.south);
|
||||
\draw[msg] (c3.west) -- (srv.east);
|
||||
|
||||
\end{tikzpicture}
|
||||
}
|
||||
\end{frame}
|
||||
|
||||
|
||||
\begin{frame}{Système centralisé : exécution parallèle des clients}
|
||||
\begin{itemize}
|
||||
\item Requêtes concurrentes côté clients
|
||||
\item Traitement séquentiel côté serveur (ordre total)
|
||||
\end{itemize}
|
||||
|
||||
\vspace{0.3cm}
|
||||
\centering
|
||||
\resizebox{0.6\linewidth}{!}{%
|
||||
\begin{tikzpicture}[
|
||||
x=1.4cm, y=0.9cm,
|
||||
msg/.style={-Latex, thick},
|
||||
proc/.style={font=\small},
|
||||
opbox/.style={draw, rounded corners, thick, fill=gray!10}
|
||||
]
|
||||
|
||||
% --- Processus : serveur + clients (lignes de vie) ---
|
||||
\node[proc, anchor=east] at (0,0) {$Serveur$};
|
||||
\node[proc, anchor=east] at (0,-1) {$P_1$};
|
||||
\node[proc, anchor=east] at (0,-2) {$P_2$};
|
||||
\node[proc, anchor=east] at (0,-3) {$P_3$};
|
||||
|
||||
\pause
|
||||
|
||||
% --- Lignes de temps ---
|
||||
\foreach \y in {0,-1,-2,-3}{
|
||||
\draw[thick] (0.2,\y) -- (4.0,\y);
|
||||
}
|
||||
|
||||
\pause
|
||||
% --- Requêtes des clients vers le serveur (légèrement décalées) ---
|
||||
\draw[msg] (0.8,-1) -- (1.2,0)
|
||||
node[midway, left, font=\scriptsize] {$op_1$};
|
||||
\draw[msg] (0.9,-2) -- (1.3,0)
|
||||
node[midway, left, font=\scriptsize] {$op_2$};
|
||||
\draw[msg] (1.0,-3) -- (1.4,0)
|
||||
node[midway, left, font=\scriptsize] {$op_3$};
|
||||
|
||||
% (On suggère qu'elles sont émises "en parallèle" côté clients.)
|
||||
|
||||
\pause
|
||||
% --- Traitement séquentiel au serveur : op_1, op_2, op_3 ---
|
||||
% op_1
|
||||
\node[opbox, minimum width=0.2cm, minimum height=0.5cm] (o1) at (1.8,0) {$op_1$};
|
||||
% op_2
|
||||
\node[opbox, minimum width=0.2cm, minimum height=0.5cm, right=0.2cm of o1] (o2) {$op_2$};
|
||||
% op_3
|
||||
\node[opbox, minimum width=0.2cm, minimum height=0.5cm, right=0.2cm of o2] (o3) {$op_3$};
|
||||
|
||||
% Relier dans l'ordre pour bien montrer la séquentialisation
|
||||
\draw[->, thick] (o1.east) -- (o2.west);
|
||||
\draw[->, thick] (o2.east) -- (o3.west);
|
||||
|
||||
% --- Réponses (optionnelles, pour fermer la boucle) ---
|
||||
\pause
|
||||
\draw[msg] (o1.south) -- (2,-1);
|
||||
\draw[msg] (o1.south) -- (2,-2);
|
||||
\draw[msg] (o1.south) -- (2,-3);
|
||||
\draw[msg] (o2.south) -- (2.8,-1);
|
||||
\draw[msg] (o2.south) -- (2.8,-2);
|
||||
\draw[msg] (o2.south) -- (2.8,-3);
|
||||
\draw[msg] (o3.south) -- (3.8,-1);
|
||||
\draw[msg] (o3.south) -- (3.8,-2);
|
||||
\draw[msg] (o3.south) -- (3.8,-3);
|
||||
|
||||
% --- Évolution de l'état (optionnel mais parlant) ---
|
||||
\node[font=\scriptsize, align=left] at (2.7,0.9)
|
||||
{$S_0 \xrightarrow{op_1} S_1 \xrightarrow{op_2} S_2 \xrightarrow{op_3} S_3$};
|
||||
\end{tikzpicture}
|
||||
}
|
||||
\end{frame}
|
||||
|
||||
\subsection{Les systèmes pair à pair}
|
||||
|
||||
\begin{frame}{Système pair à pair}
|
||||
\begin{itemize}
|
||||
\item Pas de serveur central
|
||||
\item Chaque processus entretient son état local en fonction des messages recu.
|
||||
\end{itemize}
|
||||
|
||||
\vspace{0.3cm}
|
||||
\centering
|
||||
\resizebox{0.6\linewidth}{!}{%
|
||||
\begin{tikzpicture}[
|
||||
node distance=1.8cm and 2.5cm,
|
||||
peer/.style={draw, circle, thick, minimum size=1cm, font=\small},
|
||||
link/.style={<->, thick}
|
||||
]
|
||||
|
||||
% --- Pairs ---
|
||||
\node[peer] (p1) at (-2,0) {$P_1$};
|
||||
\node[peer] (p2) at ( 2,0) {$P_2$};
|
||||
\node[peer] (p3) at ( 0,-2) {$P_3$};
|
||||
|
||||
% --- Connexions pair à pair ---
|
||||
\draw[link] (p1) -- (p2);
|
||||
\draw[link] (p1) -- (p3);
|
||||
\draw[link] (p2) -- (p3);
|
||||
|
||||
\end{tikzpicture}
|
||||
}
|
||||
\end{frame}
|
||||
|
||||
\begin{frame}{Système pair à pair : exécution parallèle}
|
||||
\begin{itemize}
|
||||
\item Chaque pair exécute des opérations concurrentes
|
||||
\item Pas de serveur central pour imposer un ordre total
|
||||
\end{itemize}
|
||||
|
||||
\centering
|
||||
\resizebox{0.7\linewidth}{!}{%
|
||||
\begin{tikzpicture}[
|
||||
x=1.4cm, y=0.9cm,
|
||||
proc/.style={font=\small},
|
||||
opbox/.style={draw, rounded corners, thick}
|
||||
]
|
||||
|
||||
% --- Processus ---
|
||||
\node[proc, anchor=east] at (0,0) {$P_1$};
|
||||
\node[proc, anchor=east] at (0,-1) {$P_2$};
|
||||
\node[proc, anchor=east] at (0,-2) {$P_3$};
|
||||
|
||||
% --- Lignes de temps pour chaque processus ---
|
||||
\foreach \y in {0,-1,-2}{
|
||||
\draw[thick] (0.2,\y) -- (3.0,\y);
|
||||
}
|
||||
|
||||
% op_1 : opération de P1
|
||||
% commence à x = 0.35 et finit à x = 2.5
|
||||
\draw[opbox,densely dotted] (0.35,0.25) rectangle (2.5,-0.25);
|
||||
% ligne de vie de P_1 en pointillé dans l'intervalle de op_1
|
||||
\draw[thick,densely dotted] (0.35,0) -- (2.5,0);
|
||||
\node[font=\scriptsize, anchor=west] at (2.55,-0.2) {$op_1$};
|
||||
|
||||
% op_2 : opération de P2
|
||||
% commence à x = 1.0 et finit à x = 2.2
|
||||
\draw[opbox,densely dotted] (0.45,-0.75) rectangle (1.2,-1.25);
|
||||
% ligne de vie de P_2 en pointillé dans l'intervalle de op_2
|
||||
\draw[thick,densely dotted] (1.0,-1) -- (2.2,-1);
|
||||
\node[font=\scriptsize, anchor=west] at (2.25,-1.2) {$op_2$};
|
||||
|
||||
% op_3 : opération de P3
|
||||
% commence à x = 1.4 et finit à x = 2.5
|
||||
\draw[opbox,densely dotted] (1.4,-1.75) rectangle (2.5,-2.25);
|
||||
% ligne de vie de P_3 en pointillé dans l'intervalle de op_3
|
||||
\draw[thick,densely dotted] (1.4,-2) -- (2.5,-2);
|
||||
\node[font=\scriptsize, anchor=west] at (2.55,-2.2) {$op_3$};
|
||||
|
||||
\end{tikzpicture}
|
||||
}
|
||||
\end{frame}
|
||||
|
||||
\subsection{Les systèmes fédérés}
|
||||
|
||||
\begin{frame}{Système fédéré}
|
||||
\begin{itemize}
|
||||
\item Plusieurs serveurs, chacun gère un sous-ensemble de processus
|
||||
\item Les serveurs se synchronisent entre eux (fédération)
|
||||
\end{itemize}
|
||||
|
||||
\vspace{0.3cm}
|
||||
\centering
|
||||
\resizebox{0.4\linewidth}{!}{%
|
||||
\begin{tikzpicture}[
|
||||
node distance=1.8cm and 2.5cm,
|
||||
server/.style={draw, thick, rounded corners, minimum width=2.8cm,
|
||||
minimum height=1.2cm, font=\small, fill=gray!10},
|
||||
client/.style={draw, circle, thick, minimum size=0.9cm, font=\scriptsize},
|
||||
srvlink/.style={<->, thick}
|
||||
]
|
||||
|
||||
% --- Serveurs fédérés ---
|
||||
\node[server] (s1) at (0,1.5) {Serveur $S_1$};
|
||||
\node[server] (s2) at (-3,-0.5){Serveur $S_2$};
|
||||
\node[server] (s3) at ( 3,-0.5){Serveur $S_3$};
|
||||
|
||||
% --- Processus locaux à chaque serveur ---
|
||||
\node[client] (p1) at (0,3) {$P_1$};
|
||||
\node[client] (p2) at (-3,-2.5) {$P_2$};
|
||||
\node[client] (p3) at ( 3,-2.5) {$P_3$};
|
||||
|
||||
% --- Liens client -- serveur local ---
|
||||
\draw[-Latex, thick] (p1.south) -- (s1.north);
|
||||
\draw[-Latex, thick] (p2.north) -- (s2.south);
|
||||
\draw[-Latex, thick] (p3.north) -- (s3.south);
|
||||
|
||||
% --- Fédération entre serveurs ---
|
||||
\draw[srvlink] (s1) -- (s2);
|
||||
\draw[srvlink] (s1) -- (s3);
|
||||
\draw[srvlink] (s2) -- (s3);
|
||||
|
||||
\end{tikzpicture}
|
||||
}
|
||||
\end{frame}
|
||||
|
||||
\begin{frame}{Système fédéré : exécution parallèle des serveurs}
|
||||
\begin{itemize}
|
||||
\item Chaque serveur impose un ordre total local sur ses opérations
|
||||
\item Globalement, les opérations sur différents serveurs sont concurrentes
|
||||
\end{itemize}
|
||||
|
||||
\centering
|
||||
\resizebox{0.7\linewidth}{!}{%
|
||||
\begin{tikzpicture}[
|
||||
x=1.4cm, y=0.9cm,
|
||||
proc/.style={font=\small},
|
||||
opbox/.style={draw, rounded corners, thick}
|
||||
]
|
||||
|
||||
% --- Processus (serveurs fédérés) ---
|
||||
\node[proc, anchor=east] at (0,0) {$S_1$};
|
||||
\node[proc, anchor=east] at (0,-1) {$S_2$};
|
||||
\node[proc, anchor=east] at (0,-2) {$S_3$};
|
||||
|
||||
% --- Lignes de temps pour chaque serveur ---
|
||||
\foreach \y in {0,-1,-2}{
|
||||
\draw[thick] (0.2,\y) -- (3.2,\y);
|
||||
}
|
||||
|
||||
% op_1 : opération sur S_1 (longue)
|
||||
\draw[opbox,densely dotted] (0.35,0.25) rectangle (2.3,-0.25);
|
||||
\draw[thick,densely dotted] (0.35,0) -- (2.3,0);
|
||||
\node[font=\scriptsize, anchor=west] at (2.35,-0.2) {$op_1$};
|
||||
|
||||
% op_2 : opération sur S_2 (décalée, plus courte)
|
||||
\draw[opbox,densely dotted] (0.8,-0.75) rectangle (2.0,-1.25);
|
||||
\draw[thick,densely dotted] (0.8,-1) -- (2.0,-1);
|
||||
\node[font=\scriptsize, anchor=west] at (2.05,-1.2) {$op_2$};
|
||||
|
||||
% op_3 : opération sur S_3 (commence plus tard, finit après op_2)
|
||||
\draw[opbox,densely dotted] (1.3,-1.75) rectangle (2.8,-2.25);
|
||||
\draw[thick,densely dotted] (1.3,-2) -- (2.8,-2);
|
||||
\node[font=\scriptsize, anchor=west] at (2.85,-2.2) {$op_3$};
|
||||
|
||||
\end{tikzpicture}
|
||||
}
|
||||
\end{frame}
|
||||
|
||||
|
||||
|
||||
\begin{frame}{Jeu : placer les applications dans les architectures}
|
||||
\begin{itemize}
|
||||
\item À quel type d'architecture appartient chaque application ?
|
||||
\end{itemize}
|
||||
|
||||
\vspace{0.3cm}
|
||||
\centering
|
||||
\resizebox{0.95\linewidth}{!}{%
|
||||
\begin{tikzpicture}[
|
||||
archibox/.style={draw, rounded corners, thick,
|
||||
minimum width=3.5cm, minimum height=1.6cm,
|
||||
font=\small, align=center, fill=gray!10},
|
||||
app/.style={draw, rounded corners, thick,
|
||||
minimum width=2.0cm, minimum height=0.8cm,
|
||||
font=\scriptsize, align=center}
|
||||
]
|
||||
|
||||
% --- Zones d'architecture ---
|
||||
\node[archibox] (central) at (-4, 1.5) {Centralisé};
|
||||
\node[archibox] (p2p) at ( 0, 1.5) {Pair à pair};
|
||||
\node[archibox] (federated) at ( 4, 1.5) {Fédéré};
|
||||
|
||||
% --- Jetons d'applications à placer ---
|
||||
\node[app] at (-5, -0.2) {cryptpad};
|
||||
\node[app] at (-2.5,-0.2) {grist};
|
||||
\node[app] at ( 0, -0.2) {colabora};
|
||||
\node[app] at ( 2.5,-0.2) {matrix};
|
||||
\node[app] at ( 5, -0.2) {signal};
|
||||
|
||||
\node[app] at (-5, -1.8) {XMPP};
|
||||
\node[app] at (-2.5,-1.8) {IRC};
|
||||
\node[app] at ( 0, -1.8) {IMAP};
|
||||
\node[app] at ( 2.5,-1.8) {PeerTube};
|
||||
\node[app] at ( 5, -1.8) {DNS};
|
||||
|
||||
\end{tikzpicture}
|
||||
}
|
||||
\end{frame}
|
||||
|
||||
|
||||
\begin{frame}{Un petit Quizz}
|
||||
% \qrcode{partici.fi/38956702}
|
||||
\end{frame}
|
||||
|
||||
\begin{frame}
|
||||
\frametitle{MERCI !}
|
||||
|
||||
|
||||
|
||||
\end{frame}
|
||||
|
||||
|
||||
\end{document}
|
||||
0
misc/synthese2024/programme/main.pdf
Normal file
0
misc/synthese2024/programme/main.pdf
Normal file
0
misc/synthese2024/programme/main.synctex(busy)
Normal file
0
misc/synthese2024/programme/main.synctex(busy)
Normal file
138
misc/synthese2024/programme/main.tex
Normal file
138
misc/synthese2024/programme/main.tex
Normal file
@@ -0,0 +1,138 @@
|
||||
\documentclass[a4paper,12pt]{article}
|
||||
|
||||
\usepackage[utf8]{inputenc}
|
||||
\usepackage[T1]{fontenc}
|
||||
\usepackage[french]{babel}
|
||||
\usepackage{geometry}
|
||||
\usepackage{xcolor}
|
||||
\usepackage{graphicx}
|
||||
\usepackage{tabularx}
|
||||
\usepackage{array}
|
||||
\usepackage{setspace}
|
||||
|
||||
\geometry{margin=1.5cm}
|
||||
\pagestyle{empty}
|
||||
|
||||
% Quelques couleurs douces pour l'affiche
|
||||
\definecolor{maincolor}{RGB}{40,70,140}
|
||||
\definecolor{lightbg}{RGB}{230,236,250}
|
||||
|
||||
\renewcommand{\familydefault}{\sfdefault} % Police sans serif
|
||||
\setlength{\parindent}{0pt}
|
||||
\renewcommand{\arraystretch}{1.5}
|
||||
|
||||
\begin{document}
|
||||
\thispagestyle{empty}
|
||||
|
||||
% --- Bandeau haut : logo + titre ---
|
||||
\begin{minipage}[t]{0.25\textwidth}
|
||||
% Remplacer "logo.pdf" par le fichier de votre logo, ou commenter la ligne
|
||||
%\includegraphics[width=\textwidth]{logo.pdf}
|
||||
\end{minipage}
|
||||
\hfill
|
||||
\begin{minipage}[t]{0.7\textwidth}
|
||||
\raggedleft
|
||||
{\Huge\bfseries Après-midi Synthèse}\\[0.4em]
|
||||
{\Large Rencontre des Doctorants et Post-Doctorants du LIS}\\[0.4em]
|
||||
{\large Date 03/12/2025\quad--\quad Lieu Saint-Charles FRUMAN}
|
||||
\end{minipage}
|
||||
|
||||
\vspace{1cm}
|
||||
|
||||
% --- Bloc "Programme" ---
|
||||
\begingroup
|
||||
\color{maincolor}
|
||||
{\LARGE\bfseries Programme de l'après-midi (14h--18h)}
|
||||
\endgroup
|
||||
|
||||
\vspace{0.5cm}
|
||||
|
||||
\setlength{\fboxsep}{8pt}
|
||||
\colorbox{lightbg}{%
|
||||
\begin{minipage}{0.98\textwidth}
|
||||
\vspace{0.3cm}
|
||||
\small % on réduit un peu la taille dans le tableau pour que tout tienne
|
||||
\begin{tabularx}{\textwidth}{>{\bfseries}p{2.8cm} X}
|
||||
14h00 -- 15h00 & \textit{Brise Glace} \\
|
||||
& \textit{Intervenant·e : Marius} \\[0.3em]
|
||||
15h00 -- 15h15 & \textit{Présentation du LIS} \\
|
||||
& \textit{Intervenant·e : Amaury} \\[0.3em]
|
||||
15h15 -- 15h30 & \textit{Présentation du Comité JCC} \\[0.3em]
|
||||
& \textit{Intervenant·e : Amaury} \\[0.3em]
|
||||
15h30 -- 15h45 & \textit{Présentation du Comité Parité} \\
|
||||
& \textit{Intervenant·e·s : Elie, Aliénor} \\[0.3em]
|
||||
15h45 -- 16h30& & \textit{Wooclap des Doctorants} \\
|
||||
& \textit{Intervenant·e : Marius} \\[0.3em]
|
||||
16h00 -- 16h15 & \textit{Pause café} \\[0.3em]
|
||||
16h30 -- 18h00 & \textit{Speed-Dating} \\
|
||||
& \textit{Intervenant·e : Ioana} \\[0.3em]
|
||||
18h00 -- 19h00 & \textit{Nom Nom Nom} \\[0.3em]
|
||||
19h00 -- \dots & \textit{Soirée Jeu à la Barakajeu} \\
|
||||
\end{tabularx}
|
||||
\vspace{0.4cm}
|
||||
\end{minipage}
|
||||
}
|
||||
|
||||
\vspace{1cm}
|
||||
|
||||
|
||||
\vfill
|
||||
|
||||
\newpage
|
||||
\thispagestyle{empty}
|
||||
|
||||
% ===================== PAGE 2 : VERSION ANGLAISE =====================
|
||||
\selectlanguage{english}
|
||||
|
||||
% --- Top banner: logo + title ---
|
||||
\begin{minipage}[t]{0.25\textwidth}
|
||||
% Replace "logo.pdf" with your logo file, or comment the line
|
||||
%\includegraphics[width=\textwidth]{logo.pdf}
|
||||
\end{minipage}
|
||||
\hfill
|
||||
\begin{minipage}[t]{0.7\textwidth}
|
||||
\raggedleft
|
||||
{\Huge\bfseries SynThèse Afternoon}\\[0.4em]
|
||||
{\Large Meeting of LIS PhD Students and Postdocs}\\[0.4em]
|
||||
{\large Date 03/12/2025\quad--\quad Venue Saint-Charles FRUMAN}
|
||||
\end{minipage}
|
||||
|
||||
\vspace{1cm}
|
||||
|
||||
% --- "Programme" block ---
|
||||
\begingroup
|
||||
\color{maincolor}
|
||||
{\LARGE\bfseries Afternoon schedule (14:00--19:00)}
|
||||
\endgroup
|
||||
|
||||
\vspace{0.5cm}
|
||||
|
||||
\setlength{\fboxsep}{8pt}
|
||||
\colorbox{lightbg}{%
|
||||
\begin{minipage}{0.98\textwidth}
|
||||
\vspace{0.3cm}
|
||||
\small
|
||||
\begin{tabularx}{\textwidth}{>{\bfseries}p{2.8cm} X}
|
||||
14h00 -- 15h00 & \textit{Icebreaker} \\
|
||||
& \textit{Facilitator: Marius} \\[0.3em]
|
||||
15h00 -- 15h15 & \textit{Presentation of LIS} \\
|
||||
& \textit{Speaker: Amaury} \\[0.3em]
|
||||
15h15 -- 15h45 & \textit{PhD students Wooclap quiz} \\
|
||||
& \textit{Facilitator: Marius} \\[0.3em]
|
||||
15h45 -- 16h00 & \textit{Presentation of the JCC Committee} \\
|
||||
& \textit{Speaker: Amaury} \\[0.3em]
|
||||
16h15 -- 16h30 & \textit{Presentation of the Parity Committee} \\
|
||||
& \textit{Speakers: Elie, Aliénor} \\[0.3em]
|
||||
16h00 -- 16h15 & \textit{Coffee break} \\[0.3em]
|
||||
16h30 -- 18h00 & \textit{Speed-dating} \\
|
||||
& \textit{Facilitator: Ioana} \\[0.3em]
|
||||
18h00 -- 19h00 & \textit{Nom Nom Nom (food \& drinks)} \\[0.3em]
|
||||
19h00 -- \dots & \textit{Board games night at Barakajeu} \\
|
||||
\end{tabularx}
|
||||
\vspace{0.4cm}
|
||||
\end{minipage}
|
||||
}
|
||||
|
||||
\vfill
|
||||
|
||||
\end{document}
|
||||
Reference in New Issue
Block a user