521 lines
35 KiB
TeX
521 lines
35 KiB
TeX
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\subsection{Model extension}
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We consider a static set $\Pi$ of $n$ processes with known identities, communicating by reliable point-to-point channels, in a complete graph. Messages are uniquely identifiable. At most $f$ processes may be Byzantine, and we assume $n > 3f$.
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\paragraph{Synchrony.} The network is asynchronous.
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\paragraph{Communication.} Processes can exchange through a Reliable Broadcast ($\RB$) primitive (defined below) which is invoked with the functions $\RBcast(m)$ and $m = \rbreceived()$. The Reliable Broadcast is a Byzantine fault-tolerant primitive following Bracha's specification \cite{Bracha1987}: for any message $m$ broadcast by a process $p_j$, every correct process either receives the message $m$ or none at all; moreover, all correct processes that receive a broadcast from $p_j$ for a given sequence number or round receive exactly the same message. This ensures that Byzantine processes cannot send different messages to different correct processes for the same broadcast instance. There exists a shared object called DenyList ($\DL$) (defined below) that is interfaced with a set $O$ of operations. There exist three types of these operations: $\APPEND(x)$, $\PROVE(x)$ and $\READ()$.
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\paragraph{Byzantine behaviour}
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A process is said to exhibit Byzantine behaviour if it deviates arbitrarily from the prescribed algorithm. Such deviations may, for instance, consist in invoking primitives ($\RBcast$, $\APPEND$, $\PROVE$, etc.) with invalid inputs or inputs crafted maliciously, colluding with other Byzantine processes in an attempt to manipulate the system state or violate its guarantees, deliberately delaying or accelerating the delivery of messages to selected nodes so as to disrupt the expected timing of operations, or withholding messages and responses in order to induce inconsistencies in the system state.
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Byzantine processes are constrained by the following. They cannot forge valid cryptographic signatures or threshold shares without access to the corresponding private keys. They cannot violate the termination, validity, or anti-flickering properties of the $\DL$ object. They also cannot break the integrity, no-duplicates, or validity properties of the $\RB$ primitive.
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\paragraph{Notation.} For any indice $k$ we defined by $\DL[k]$ as the $k$-th DenyList object. For a given $\DL[k]$ and any indice $x$ we defined by $\Pi_x^k$ a subset of $\Pi$. Still for a given $k$ we consider $\Pi_M^k \subseteq \Pi$ and $\Pi_V^k \subseteq \Pi$ two authorization subsets for $\DL[k]$. Indice $i \in \Pi$ 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 $\prec_k$ for the $\DL[k]$ linearization: $x \prec_k y$ means that operation $x$ appears strictly before $y$ in the linearized history of $\DL[k]$. 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).
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For any operation $F \in O$,$F_i(...)$ denotes that the operation $F$ is invoked by process $p_i$, and by $F_i^k(...)$ the same operation invoked on the $\DL[k]$ object.
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% ------------------------------------------------------------------------------
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\subsection{Primitives}
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\subsubsection{t-BFT-DL}
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We consider a t-Byzantine Fault Tolerant DenyList (t-$\BFTDL$) with the following properties.
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There are 3 operations : $\BFTPROVE(x), \BFTAPPEND(x), \BFTREAD()$ such that :
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\paragraph{Termination.} Every operation $\BFTAPPEND(x)$, $\BFTPROVE(x)$, and $\BFTREAD()$ invoked by a correct process always returns.
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\paragraph{PROVE Validity.} The invocation of $op = \BFTPROVE(x)$ by a correct process is valid iff there exist a set of correct process $C$ such that $\forall c \in C$, $c$ invoke $op_2 = \BFTAPPEND(x)$ with $op_2 \prec op_1$ and $|C| \leq t$
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\paragraph{PROVE Anti-Flickering.} If the invocation of a operation $op = \BFTPROVE(x)$ by a correct process $p \in \Pi_V$ is invalid, then any $\BFTPROVE(x)$ operation that appears after $op$ in $\Seq$ is invalid.
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\paragraph{READ Liveness.} Let $op = \BFTREAD()$ invoke by a correct process such that $R$ is the result of $op$. For all $(i, \PROVEtrace(x)) \in R$ there exist a valid invocation of $\BFTPROVE(x)$ by $p_i$.
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\paragraph{READ Anti-Flickering.} Let $op_1, op_2$ two $\BFTREAD()$ operations that returns respectively $R_1, R_2$. Iff $op_1 \prec op_2$ then $R_2 \subseteq R_1$. Otherwise $R_1 \subseteq R_2$.
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\paragraph{READ Safety.} Let $op_1, op_2$ respectively a valid $\BFTPROVE(x)$ operation submited by the process $p_i$ and a $\BFTREAD()$ operation submited by any correct process such that $op_1 \prec op_2$. Let $R$ the result of $op_2$ then $R \ni (i, x)$
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\subsection{DL $\Rightarrow$ t-BFT-DL}
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Fix $3t < |M|$. Let
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\[
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\mathcal{U} = \{\, U \subseteq M \mid |U| = |M| - t \,\}.
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\]
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For each $U \in \mathcal{U}$, we instantiate one DenyList object $DL_U$ whose authorization sets are
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\[
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\Pi_M(DL_T) = S_T = U
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\qquad\text{and}\qquad
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\Pi_V(DL_T) = V.
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\]
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\[
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|\mathcal{U}| = \binom{|M|}{|M| - t}.
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\]
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\begin{algorithm}
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\caption{t-BFT-DL implementation using multiple DL objects}
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\Fn{$\BFTAPPEND(x)$}{
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\For{\textbf{each } $U \in \mathcal{U}$ st $i \in U$}{
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$DL_U.\APPEND(x)$\;
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}
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}
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\vspace{1em}
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\Fn{$\BFTPROVE(x)$}{
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$\state \gets false$\;
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\For{\textbf{each } $U \in \mathcal{U}$}{
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$\state \gets \state \textbf{ OR } DL_U.\PROVE(x)$\;\nllabel{code:prove-or}
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}
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\Return{$\state$}\;
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}
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\vspace{1em}
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\Fn{$\BFTREAD()$}{
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$\results \gets \emptyset$\;
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\For{\textbf{each } $U \in \mathcal{U}$}{
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$\results \gets \results \cup DL_U.\READ()$\;
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}
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\Return{$\results$}\;
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}
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\end{algorithm}
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\begin{lemma}[BFT-PROVE Validity]\label{lem:bft-prove-validity}
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The invocation of $op = \BFTPROVE(x)$ by a correct process is invalid iff there exist at least $t+1$ distinct processes in $M$ that invoked a valid $\BFTAPPEND(x)$ before $op$ in $\Seq$.
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\end{lemma}
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\begin{proof}
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Let $op=\BFTPROVE(x)$ be an invocation by a correct process $p_i$. Let $A\subseteq M$ be the set of distinct issuers that invoked $\BFTAPPEND(x)$ before $op$ in $\Seq$.
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\begin{itemize}
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\item \textbf{Case (i): $|A|\ge t+1$.}
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Fix any $U\in\mathcal{U}$. $A\cap U\neq\emptyset$. Pick $j\in A\cap U$. Since $j\in U$, the call $\BFTAPPEND_j(x)$ triggers $DL_U.\APPEND(x)$, and because $\BFTAPPEND_j(x)\prec op$ in $\Seq$, this induces a valid $DL_U.\APPEND(x)$ that appears before the induced $DL_U.\PROVE(x)$ by $p_i$. By \textbf{PROVE Validity} of $\DL$, the induced $DL_U.\PROVE(x)$ is invalid. As this holds for every $U\in\mathcal{U}$, there is \emph{no} component $DL_U$ where $\PROVE(x)$ is valid, so the field $\state$ at line~\ref{code:prove-or} is never becoming true, and $op$ return false.
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\item \textbf{Case (ii): $|A|\le t$.}
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There exists $U^\star\in\mathcal{U}$ such that $A\cap U^\star=\emptyset$. For any $j\in A$, we have $j\notin U^\star$, so $\BFTAPPEND_j(x)$ does \emph{not} call $DL_{U^\star}.\APPEND(x)$. Hence no valid $DL_{U^\star}.\APPEND(x)$ appears before the induced $DL_{U^\star}.\PROVE(x)$. Since also $i\in \Pi_V(DL_{U^\star})$, by \textbf{PROVE Validity} of $\DL$ the induced $DL_{U^\star}.\PROVE(x)$ is valid. Therefore, there exists a component with a valid $\PROVE(x)$, so $op$ is valid.
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\end{itemize}
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\smallskip
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Combining the cases yields the claimed characterization of invalidity.
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\end{proof}
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\begin{lemma}[BFT-PROVE Anti-Flickering]\label{lem:bft-prove-anti-flickering}
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If the invocation of a operation $op = \BFTPROVE(x)$ by a correct process $p \in \Pi_V$ is invalid, then any $\BFTPROVE(x)$ operation that appears after $op$ in $\Seq$ is invalid.
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\end{lemma}
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\begin{proof}
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Let $op=\BFTPROVE(x)$ be an invocation by a correct process $p_i$ that is \emph{invalid} in $\Seq$.
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By BFT-PROVE Validity, this implies that there exist at least $t+1$ \emph{distinct} processes in $M$ that invoked a \emph{valid} $\BFTAPPEND(x)$ before $op$ in $\Seq$. Let $A\subseteq M$ denote that set, with $|A|\ge t+1$.
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Fix any $U\in\mathcal{U}$. We have $A\cap U\neq\emptyset$. Pick $j\in A\cap U$. Since $j\in U$, the call $\BFTAPPEND_j(x)$ triggers a call $DL_U.\APPEND(x)$. Moreover, because $\BFTAPPEND_j(x)\prec op$ in $\Seq$, the induced $DL_U.\APPEND(x)$ appears before the induced $DL_U.\PROVE(x)$ of $op$ in the projection $\Seq_U$.
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Hence, in $\Seq_U$, there exists a \emph{valid} $DL_U.\APPEND(x)$ that appears before the $DL_U.\PROVE(x)$ induced by $op$. By \textbf{PROVE Validity} the base $\DL$ object, the induced $DL_U.\PROVE(x)$ is therefore \emph{invalid} in $\Seq_U$.
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Let $op'=\BFTPROVE(x)$ be any invocation such that $op\prec op'$ in $\Seq$. Fix again any $U\in\mathcal{U}$. Hence, the $DL_U.\PROVE(x)$ induced by $op'$ appears after the $DL_U.\PROVE(x)$ induced by $op$ in $\Seq_U$. Since the induced $DL_U.\PROVE(x)$ of $op$ is invalid, by \textbf{PROVE Anti-Flickering} of $\DL$, \emph{every} subsequent $DL_U.\PROVE(x)$ in $\Seq_U$ is invalid.
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As this holds for every $U\in\mathcal{U}$, there is no component $DL_U$ in which the induced $\PROVE(x)$ of $op'$ is valid.
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\end{proof}
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\begin{lemma}[BFT-READ Liveness]
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Let $op = \BFTREAD()$ invoke by a correct process such that $R$ is the result of $op$. For all $(i, x) \in R$ there exist a valid invocation of $\BFTPROVE(x)$ by $p_i$.
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\end{lemma}
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\begin{proof}
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Let $R$ the result of a $READ()$ operation submit by any correct process. $(i, \PROVEtrace(x)) \in R$ implie that $\exists U^\star \in \mathcal{U}$ such that $(i, x) \in R^{U^\star}$ with $R^{U^\star}$ the result of $DL_{U^\star}.\READ()$. By \textbf{READ Validity} $(i, x) \in R^{U^\star}$ implie that there exist a valid $DL_{U^\star}.\PROVE_i(x)$. The for loop in the $\BFTPROVE(x)$ implementation return true iff there at least one valid $DL_{U}.\PROVE_i(x)$ for any $U \in \mathcal{U}$.
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Hence because there exist a $U^\star$ such that $DL_{U^\star}.\PROVE_i(x)$, there exist a valid $\BFTPROVE_i(x)$.
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$(i, x) \in R \implies \exists \BFTPROVE_i(x)$
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\end{proof}
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\begin{lemma}[BFT-READ Anti-Flickering]\label{lem:bft-read-anti-flickering}
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Let $op_1, op_2$ two $\BFTREAD()$ operations that returns respectively $R_1, R_2$. Iff $op_1 \prec op_2$ then $R_2 \subseteq R_1$. Otherwise $R_1 \subseteq R_2$.
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\end{lemma}
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\begin{proof}
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Let $R_1, R_2$ respectively the output of two $\BFTREAD()$ operations $op_1, op_2$ such that $op_1 \prec op_2$.
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By the implementation of $\BFTREAD$, $R_k = \bigcup_{U \in \mathcal{U}} R_k^U$ where $R_k^U$ is the result of $DL_U.\READ()$ during $op_k$.
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Because $op_1 \prec op_2$ for any $U \in \mathcal{U}$, the $DL_U.\READ()$ induced by $op_1$ happen before the $DL_U.\READ()$ induced by $op_2$. Hence we have for all $U, R_2^U \subseteq R_1^U$.
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Therefore
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\[
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\bigcup_U R_2^U \subseteq \bigcup_U R_1^U \implies
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R_2 \subseteq R_1
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\]
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\end{proof}
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\begin{lemma}[BFT-READ Safety]\label{lem:bft-read-safety}
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Let $op_1, op_2$ respectively a valid $\BFTPROVE(x)$ operation submited by the process $p_i$ and a $\BFTREAD()$ operation submited by any correct process such that $op_1 \prec op_2$. Let $R$ the result of $op_2$ then $R \ni (i, x)$
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\end{lemma}
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\begin{proof}
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Let $op_1 = \BFTPROVE_i(x)$ be a valid operation by a correct process $p_i$ and $op_2 = \BFTREAD()$ be any $\BFTREAD()$ operation such that $op_1 \prec op_2$ in $\Seq$.
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By BFT-PROVE Validity, there exist at most $t$ distinct processes in $M$ that invoked a valid $\BFTAPPEND(x)$ before $op_1$ in $\Seq$. Let $A\subseteq M$ denote that set, with $|A|\le t$.
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There exists $U^\star\in\mathcal{U}$ such that $A\cap U^\star=\emptyset$. For any $j\in A$, we have $j\notin U^\star$, so $\BFTAPPEND^{(j)}(x)$ does \emph{not} call $DL_{U^\star}.\APPEND(x)$. Hence no valid $DL_{U^\star}.\APPEND(x)$ appears before the induced $DL_{U^\star}.\PROVE(x)$ of $op_1$. Since also $i\in \Pi_V(DL_{U^\star})$, by \textbf{PROVE Validity} of $\DL$ the induced $DL_{U^\star}.\PROVE_i(x)$ is valid.
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Now, because $op_1 \prec op_2$ in $\Seq$, the induced $DL_{U^\star}.\PROVE_i(x)$ appears before the induced $DL_{U^\star}.\READ()$ of $op_2$ in $\Seq_{U^\star}$. By \textbf{READ Safety} of $\DL$, the result $R^{U^\star}$ of the induced $DL_{U^\star}.\READ()$ contains $(i, x)$.
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Finally, by the implementation of $\BFTREAD()$, we have $R = \bigcup_{U \in \mathcal{U}} R^U$, so $(i, x) \in R$.
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\end{proof}
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\begin{theorem}
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For any fixed value $t$ such that $3t < |M|$, multiple DenyList Object can be used to implement a t-Byzantine Fault Tolerant DenyList Object.
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\end{theorem}
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\begin{proof}
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Follows directly from the previous lemmas.
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\end{proof}
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\subsection{Algorithm}
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\begin{algorithm}[H]
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\caption{t-BFT ARB at process $p_i$}\label{alg:bft-arb}
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\SetKwBlock{LocalVars}{Local Variables:}{}
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\LocalVars{
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$\unordered \gets \emptyset$,
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$\ordered \gets \epsilon$,
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$\delivered \gets \epsilon$\;
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$\prop[r][j] \gets \bot, \forall j, r \in \Pi \times \mathbb{N}$\;
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$\done[r] \gets \emptyset, \forall r \in \mathbb{N}$\;
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}
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\vspace{0.3em}
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\For{$r = 1, 2, \ldots$}{\nllabel{alg:main-loop}
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\textbf{wait until} $\unordered \setminus \ordered \neq \emptyset$\;
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$S \gets \unordered \setminus \ordered$;
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$\RBcast(i, \texttt{PROP}, S, r)$\;
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\textbf{wait until} $|\validated(r)| \geq n - t$\;\nllabel{alg:check-validated}
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\BlankLine
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\lForEach{$j \in \Pi$}{
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$\BFTAPPEND(\langle j, r\rangle)$\;\nllabel{alg:append}
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}
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\lForEach{$j \in \Pi$}{
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$\send(j, \texttt{DONE}, r)$\;
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}
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\BlankLine
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\textbf{wait until} $|\done[r]| \geq n - t$\;\nllabel{alg:check-done}
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\textbf{wait until} $\forall j \in \winners[r],\ \prop[r][j] \neq \bot$ \textbf{ with } $\winners[r] \gets \validated(r)$\;
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\BlankLine
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$M \gets \bigcup_{j \in \winners[r]} \prop[r][j]$\;\nllabel{code:Mcompute}
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$\ordered \gets \ordered \cdot \ordered(M)$\;
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}
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\vspace{0.3em}
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\Fn{\validated($r$)}{
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\Return{$\{j: |\{k: (k, r) \in \BFTREAD()\}| \geq t+1\}$}\;
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}
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\vspace{0.3em}
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\Upon{$\ABbroadcast(m)$}{
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$\unordered \gets \unordered \cup \{m\}$\;
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}
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\vspace{0.3em}
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\Upon{$\rdeliver(\texttt{PROP}, S, \langle j, r \rangle)$ from process $p_j$}{
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$\unordered \gets \unordered \cup S$;
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$\prop[r][j] \gets S$\;
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$\BFTPROVE(\langle j, r\rangle)$\;
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}
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\vspace{0.3em}
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\Upon{$\receive(\texttt{DONE}, r)$ from process $p_j$}{
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$\done[r] \gets \done[r] \cup \{j\}$\;
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}
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\vspace{0.3em}
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\Upon{$\ABdeliver()$}{
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\lIf{$\ordered \setminus \delivered = \emptyset$}{
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\Return{$\bot$}
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}
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let $m$ be the first message in $(\ordered \setminus \delivered)$\;
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$\delivered \gets \delivered \cdot \{m\}$\;
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\Return{$m$}
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}
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\end{algorithm}
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% \textbf{Everything below has to be updated}
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% \begin{definition}[BFT Closed round for $k$]
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% Given $Seq^{k}$ the linearization of the $\BFTDL$ $Y[k]$, a round $r \in \mathcal{R}$ is \emph{closed} in $\Seq$ iff there exist at least $n - f$ distinct processes $j \in \Pi$ such that $\BFTAPPEND_j(r)$ appears in $\Seq^k$. Let call $\BFTAPPEND(r)^\star$ the $(n-f)^{th}$ $\BFTAPPEND(r)$.
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% \end{definition}
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% \begin{definition}[BFT Closed round]\label{def:bft-closed-round}
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% A round $r \in \mathcal{R}$ is \emph{closed} iff for all $\DL[k]$, $r$ is closed in $\Seq^k$.
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% \end{definition}
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% \subsection{Proof of correctness}
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% \begin{remark}[BFT Stable round closure]\label{rem:bft-stable-round-closure}
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% If a round $r$ is closed, no more $\BFTPROVE(r)$ can be valid and thus linearized. In other words, once $\BFTAPPEND(r)^\star$ is linearized, no more process can make a proof on round $r$, and the set of valid proofs for round $r$ is fixed. Therefore $\Winners_r$ is fixed.
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% \end{remark}
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% \begin{proof}
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% By definition $r$ closed means that for all process $p_i$, there exist at least $n - f$ distinct processes $j \in \Pi$ such that $\BFTAPPEND_j(r)$ appears in $\Seq^k$. By BFT-PROVE Validity, any subsequent $\BFTPROVE(r)$ is invalid because at least $n - f$ processes already invoked a valid $\BFTAPPEND(r)$ before it. Thus no new valid $\BFTPROVE(r)$ can be linearized after $\BFTAPPEND(r)^\star$. Hence the set of valid proofs for round $r$ is fixed, and so is $\Winners_r$.
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% \end{proof}
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% \begin{lemma}[BFT Across rounds]\label{lem:bft-across-rounds}
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% For any $r, r'$ such that $r < r'$, if $r'$ is closed, $r$ is also closed.
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% \end{lemma}
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% \begin{proof}
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% Let $r \in \mathcal{R}$. By \cref{def:bft-closed-round}, if $r + 1$ is closed, then for all $\DL[k]$, $r + 1$ is closed in $\Seq^k$. By the implementation, a process can only invoke $\BFTAPPEND(r + 1)$ after observing at least $n - f$ valid $\BFTPROVE(r)$, which means that for all $\DL[k]$, $r$ is closed in $\Seq^k$. Hence by \cref{def:bft-closed-round}, $r$ is closed.
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% Because $r$ is monotonically increasing, we can reccursively apply the same argument to conclude that for any $r, r'$ such that $r < r'$, if $r'$ is closed, $r$ is also closed.
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% \end{proof}
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% \begin{lemma}[BFT Progress]\label{lem:bft_progress}
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% For any correct process $p_i$ such that
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% \[
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% \received \setminus (\delivered \cup (\cup_{r' < r} \cup_{j \in W[r'] \prop[r'][j]})) \neq \emptyset
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% \]
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% with $r$ the highest closed round in the $\DL$ linearization. Eventually $r+1$ will be closed.
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% \end{lemma}
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% \begin{lemma}[BFT Winners invariant]\label{lem:bft-winners-invariant}
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% For any closed round $r$, define
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% \[
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% \Winners_r = \{j: \BFTPROVE_j(r) \prec \BFTAPPEND^\star(r)\}
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% \]
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% called the unique set of winners of round $r$.
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% \end{lemma}
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% \begin{lemma}[BFT n-f lower-bounded Winners]
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% Let $r$ a closed round, $|W[r]| \geq n-f$.
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% \end{lemma}
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% \begin{remark}\label{rem:correct-in-winners}
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% Because we assume $n \geq 2f+ 1$, if $|W[r]| \geq n-f$ at least 1 correct have to be in $W[r]$ to progress.
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% \end{remark}
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% \begin{lemma}[BFT Winners must purpose]\label{lem:bft-winners-purpose}
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% Let $r$ a closed round, for all process $p_j$ such that $j \in W[r]$, $p_j$ must have executed $\RBcast(j, PROP, \_, r)$ and hence any correct will eventually set $\prop[r][j]$ to a non-$\bot$ value.
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% \end{lemma}
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% \begin{lemma}[BFT Messages Incariant]\label{lem:bft-messages-invariant}
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% For any closed round $r$ and any correct process $p_i$ such that $\forall j \in \Winners_r$: $\prop^{(i)}[r][j] \neq \bot$ define
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% \[
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% \Messages_r = \cup_{j \in \Winners_r} prop^{(i)}[r][j]
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% \]
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% as the set of messages proposed by the winners of round $r$
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% \end{lemma}
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% \begin{lemma}[BFT EVentual proposal closure]\label{lem:bft-eventual-proposal-closure}
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% If a correct process $p_i$ define $M$ at line~\ref{code:Mcompute}, then for every $j \in \Winners_r$, $\prop^{(i)}[r][j] \neq \bot$.
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% \end{lemma}
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% \begin{lemma}[BFT Unique proposal per sender per round]\label{lem:bft-unique-proposal}
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% For any round $r$ and any process $p_i$, if $p_i$ invokes two $\RBcast$ call for the same round, such that $\RBcast(i, PROP, S, r) \prec \RBcast(i, PROP, S', r)$. Then for any correct process $p_j$, $\prop^{(j)}[r][i] \in \{\bot, S\}$
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% \end{lemma}
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% \begin{lemma}[BFT $W_r$ as grow only set]\label{lem:bft-wr-grow-only}
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% For any correct process $p_i$. If $p_i$ computes $W_r$ at two different times $t_1$ and $t_2$ with $t_1 < t_2$, then $W_r^{t_1} \subseteq W_r^{t_2}$.
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% \end{lemma}
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% \begin{proof}
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% By the implementation, $W_r$ is computed exclusively from the results of $\{j: (j, \PROVEtrace(r)) \in \bigcup_{k \in \Pi} Y[k].\BFTREAD()\}$.
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% We know by BFT-READ Anti-Flickering that for any two $\BFTREAD()$ operations $op_1, op_2$ such that $op_1 \prec op_2$, the result of $op_2$ is included in the result of $op_1$. Therefore, if $p_i$ computes $W_r$ at two different times $t_1$ and $t_2$ with $t_1 < t_2$, then $W_r^{t_1} \subseteq W_r^{t_2}$.
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% \end{proof}
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% \begin{lemma}[BFT well defined winners]\label{lem:bft-well-defined-winners}
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% For any closed round $r$, if a correct process $p_i$ compute $W_r$, then $W_r = \Winners_r$ with $|W_r| \geq n - f$.
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% \end{lemma}
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% \begin{proof}
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% By \Cref{lem:bft-read-safety}, any correct process $p_i$ computing $W_r$ after round $r$ is closed includes all valid $\BFTPROVE(r)$ in its computation of $W_r$. Therefore $W_r = \Winners_r$.
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% By \Cref{def:bft-closed-round}, at least $n - f$ distinct processes invoked a valid $\BFTAPPEND(r)$ before $\BFTAPPEND(r)^\star$. By the implementation in algorithm D, if a process correct $j$ invoked a valid $\BFTAPPEND(r)$, thats means that he observed at least $n - f$ valid $\BFTPROVE(r)$ submitted by distinct processes. By \Cref{lem:bft-wr-grow-only}, once $p_j$ observed $n - f$ valid $\BFTPROVE(r)$, any correct process computing $W_r$ will eventually observe at least these $n - f$ valid $\BFTPROVE(r)$. By \Cref{lem:bft-stable-round-closure}, no more valid $\BFTPROVE(r)$ can be linearized after round $r$ is closed, so any correct process computing the same fixed set $W_r$ of at least $n - f$ distinct processes.
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% \end{proof}
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% \begin{lemma}[BFT Non-empty winners proposal]\label{lem:bft-non-empty-winners-proposal}
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% For every process $p_i$ such as $i \in W_r$, eventually $\prop[r][i] \neq \bot$.
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% \end{lemma}
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% \begin{proof}
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% By the implementation, if $i \in W_r$, then $(i, \PROVEtrace(r))$ is included in the result of at least one $\BFTREAD()$ operation. Hence there exist a valid $\BFTPROVE(r)$ operation.
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% By \Cref{lem:bft-prove-validity}, this implies that there exist at least $f + 1$ valid $\PROVE(r)$ operation invoked by processes. At least one of these processes is correct, say $p_j$. By the implementation, $p_j$ invoked $\BFTPROVE(r)$ after receiving a $Rdeliver(j, \texttt{PROP}, S, r)$ message from $p_i$. Therefore, by the reliable broadcast properties, the message will eventually be delivered to every correct process, hence eventually for any correct process $\prop[r][i] \neq \bot$.
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% \end{proof}
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% \begin{definition}[BFT Message invariant]\label{def:bft-message-invariant}
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% For any closed round $r$, for any correct process $p_i$, such that $\nexists j \in W_r : \prop[r][j] = \bot$, twe define the set
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% \[
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% \Messages_r = \bigcup_{j \in \Winners_r} \prop[r][j]
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% \]
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% as the unique set of messages proposed during round $r$.
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% \end{definition}
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% \begin{lemma}[BFT Proposal convergence]\label{lem:bft-proposal-convergence}
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% For any closed round $r$, for any correct process $p_i$, that define $M_r$ at line B10, we have $M_r = \Messages_r$.
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% \end{lemma}
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% \begin{proof}
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% By \Cref{lem:bft-well-defined-winners}, any correct process $p_i$ computing $W_r$ after round $r$ is closed has $W_r = \Winners_r$.
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% By \Cref{lem:bft-non-empty-winners-proposal}, for any correct process $p_i$, such as $i \in W_r$, eventually $\prop[r][i] \neq \bot$.
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% Therefore, eventually for any correct process $p_i$, at line B10 we have
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% \[
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% M_r = \bigcup_{j \in W_r} \prop[r][j] = \bigcup_{j \in \Winners_r} \prop[r][j] = \Messages_r
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% \]
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% \end{proof}
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% \begin{lemma}[BFT Inclusion]\label{proof:bft-inclusion}
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% If a correct process $p_i$ ABroadcasts a message $m$, then eventually any correct process $p_j$ ADelivers $m$.
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% \end{lemma}
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% \begin{proof}
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% Let $m$ be a message ABroadcast by a correct process $p_i$ and eventually exit the \texttt{ABroadcast} function at line A10.
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% By the implementation, if $p_i$ exits the \texttt{ABroadcast} function at line A10, then there exists a round $r'$ such that $m \in \prop[r'][j]$ for some $j \in W_{r'}$.
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% Since $p_i$ is correct, seeing that $m \in \prop[r'][j]$ for some $j \in W_{r'}$ implies that $p_i$ received a $Rdeliver(j, \texttt{PROP}, S, r')$ message from $p_j$ such that $m \in S$. And because $p_j$ is in $W_{r'}$, at least $n - f$ correct processes invoked a valid $Y[j].\BFTPROVE(r')$ before the round $r'$ were closed. By the reliable broadcast properties, the $Rdeliver(j, \texttt{PROP}, S, r')$ message will eventually be delivered to every correct process, hence eventually for any correct process $m \in \prop[r'][j]$ with $j \in W_{r'}$. Hence $m$ will eventually be included in the set $\Messages_{r'}$ defined in \Cref{def:bft-message-invariant} and thus eventually be ADelivered by any correct process.
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% \end{proof}
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\subsection{Correctness Lemmas}
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\begin{definition}[Closed Round]
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A round $r \in \mathcal{R}$ is said to be \emph{closed} for a correct process $p_i$ if at the moment $p_i$ computes $\winners[r]$, it satisfies $|\{k : (k, r) \in \BFTREAD()\}| \geq n - t$.
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\end{definition}
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\begin{lemma}[Round Monotonicity]\label{lem:round-monotonicity}
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If a round $r$ is closed, then every round $r_1 < r$ is also closed.
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\end{lemma}
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\begin{proof}
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\textbf{Base:} $r = 0$. Suppose round $1$ is closed, the set $\{r' \in \mathcal{R}: r' < r\}$ is empty, so the claim holds.
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\textbf{Inductive step:} Let $r \geq 1$. Assume that the property holds for all rounds $r' < r$: whenever round $r'$ is closed, all rounds $r_2 < r'$ are also closed. We show that if round $r$ is closed, then all rounds $r_1 < r$ are closed.
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Suppose round $r$ is closed. This means at least one correct process $p_i$ has reached line \ref{alg:check-validated} and satisfied the condition $|\validated(r)| \geq n - t$. Consequently, $p_i$ has invoked $\BFTAPPEND(\langle j, r \rangle)$ for each $j \in \Pi$ at line \ref{alg:append}. Since these operations are invoked by a correct process, and they depend on valid proofs in the byzantine fault tolerant deny list, at least one of these operations must have succeeded in the DL object.
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Moreover, by the algorithm structure, a correct process can only reach the beginning of round $r$ after completing round $r-1$. In particular, process $p_i$ reaches line \ref{alg:main-loop} for round $r$ only after having completed round $r-1$, which requires the process to have observed at least $|\done[r-1]| \geq n - t$ (line \ref{alg:check-done} for round $r-1$). This condition can only be satisfied if round $r-1$ is closed: at least $n - t$ processes must have issued DONE messages, and since $n > 3f$, at least one of these is correct.
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Since $n - t > 2f + 1$, among the $n - t$ processes satisfying the condition for round $r-1$, at least one is correct. A correct process that issued a DONE message for round $r-1$ must have previously completed its execution of lines \ref{alg:append} for round $r-1$, which in turn required observing $|\validated(r-1)| \geq n - t$ from line \ref{alg:check-validated}. Thus, round $r-1$ is closed.
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Now, considering any round $r_1 < r-1$: by the inductive hypothesis applied to round $r-1$, since round $r-1$ is closed and $r_1 < r-1$, we have that $r_1$ is also closed.
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By strong induction, if round $r$ is closed, all rounds $r_1 < r$ are closed.
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\end{proof}
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\begin{lemma}[Uniqueness of Winners' Proposals]\label{lem:unique-proposal}
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For any closed round $r$ and any process $p_j \in \winners[r]$, there exists a unique set $S_j \subseteq \mathcal{M}$ such that every correct process $p_i$ that has received a reliable delivery of a $\texttt{PROP}$ message from $p_j$ for round $r$ receives exactly this set $S_j$.
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\end{lemma}
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\begin{proof}
|
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Let $r$ be a closed round and $p_j \in \winners[r]$. Since $j \in \winners[r]$, it means that $(j, r) \in \BFTREAD()$, i.e., at least $t+1$ distinct processes have invoked a valid $\BFTPROVE(\langle j, r \rangle)$ before the round $r$ closed.
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By the algorithm, a correct process $p_i$ invokes $\BFTPROVE(\langle j, r \rangle)$ only upon receiving a reliable delivery of a $\texttt{PROP}$ message from $p_j$ for round $r$ at line \ref{alg:append}. Let $S_j^{(i)}$ denote the set received by $p_i$.
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Since at least $t+1$ distinct processes have performed a valid $\BFTPROVE(\langle j, r \rangle)$, and since $t < n/3$, at least one of these processes is correct. Thus, at least one correct process must have received a reliable broadcast from $p_j$ for round $r$.
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Now, by Bracha's Byzantine Reliable Broadcast specification \cite{Bracha1987}, for any message broadcast instance by process $p_j$ in round $r$, all correct processes that deliver this broadcast either receive the identical message or none at all. Formally, for all correct processes $p_i, p_i'$ that received a delivery from $p_j$ for round $r$, we have $S_j^{(i)} = S_j^{(i')}$.
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Therefore, there exists a unique set $S_j$ such that every correct process $p_i$ that receives a reliable delivery from $p_j$ for round $r$ receives exactly $S_j$.
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\end{proof}
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\begin{lemma}[Winners Stability]\label{lem:winners-stability}
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For any closed round $r$, the set $\winners[r]$ is stable.
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\end{lemma}
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\begin{proof}
|
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Let $r$ be a closed round. By definition, a closed round $r$ means that at least one correct process $p_i$ has observed $|\{k : (k, r) \in \BFTREAD()\}| \geq n - t$ and computed $\winners[r]$ at line \ref{alg:check-done}.
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When round $r$ becomes closed, this implies that at least $n - t$ distinct processes have invoked $\BFTAPPEND(\langle j, r \rangle)$ for $j \in \Pi$ (by the algorithm structure, since $\done[r]$ contains DONE messages from processes that completed line \ref{alg:append}). Since $n > 3f$, we have $n - t > 2f + 1 > t + 1$.
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Consider a fixed $j \in \Pi$. Since at least $n - t > t + 1$ processes have invoked a valid $\BFTAPPEND(\langle j, r \rangle)$, by \cref{lem:bft-prove-validity}, any subsequent invocation of $\BFTPROVE(\langle j, r \rangle)$ will be invalid.
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By \cref{lem:bft-prove-anti-flickering}, once a $\BFTPROVE(\langle j, r \rangle)$ operation becomes invalid, all subsequent $\BFTPROVE(\langle j, r \rangle)$ operations are also invalid.
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This holds for all $j \in \Pi$. Therefore, after the round $r$ is closed, the set of valid $\BFTPROVE(\langle j, r \rangle)$ operations for each $j$ cannot grow. By \cref{lem:bft-read-anti-flickering}, any subsequent $\BFTREAD()$ invocation will not return any new $(k, r)$ pairs beyond those already returned. Thus, $\winners[r] = \{j : (j, r) \in \BFTREAD()\}$ becomes stable and cannot change.
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Since this reasoning applies to all correct processes that compute $\winners[r]$ after round $r$ is closed, all correct processes will compute the same stable set $\winners[r]$.
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\end{proof}
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\begin{lemma}[Message Content Invariance]\label{lem:message-content-invariance}
|
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For any closed round $r$ and any correct process $p_i$, the set $M$ computed at line \ref{code:Mcompute} by $p_i$ is identical to the set computed by any other correct process $p_j$ for the same round $r$.
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\end{lemma}
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|
|
|
\begin{proof}
|
|
Let $r$ be a closed round and let $p_i, p_{i'}$ be two correct processes. By the algorithm, both processes compute $M$ as:
|
|
\[
|
|
M^{(i)} = \bigcup_{j \in \winners[r]} \prop^{(i)}[r][j]
|
|
\]
|
|
and
|
|
\[
|
|
M^{(i')} = \bigcup_{j \in \winners[r]} \prop^{(i')}[r][j]
|
|
\]
|
|
at line \ref{code:Mcompute}.
|
|
|
|
By \cref{lem:winners-stability}, the set $\winners[r]$ is stable once round $r$ is closed. Therefore, both $p_i$ and $p_{i'}$ compute the same set of winners $\winners[r]$.
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|
|
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Now, consider any winner $j \in \winners[r]$. By \cref{lem:unique-proposal}, there exists a unique set $S_j \subseteq \mathcal{M}$ such that every correct process that receives a reliable delivery of a $\texttt{PROP}$ message from $p_j$ for round $r$ receives exactly $S_j$.
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|
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By the algorithm, each correct process stores this received set in its local variable: $\prop^{(i)}[r][j] = S_j$ and $\prop^{(i')}[r][j] = S_j$.
|
|
|
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Since both processes compute the union over the same set of winners, and each winner's proposal is identical for all correct processes:
|
|
\[
|
|
M^{(i)} = \bigcup_{j \in \winners[r]} \prop^{(i)}[r][j] = \bigcup_{j \in \winners[r]} S_j = \bigcup_{j \in \winners[r]} \prop^{(i')}[r][j] = M^{(i')}
|
|
\]
|
|
|
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Therefore, all correct processes compute the same set $M$ for any closed round $r$.
|
|
\end{proof}
|
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|
|
\begin{lemma}[Inclusion]\label{lem:inclusion}
|
|
If a correct process $p_i$ invokes $\ABbroadcast(m)$, then there exist a closed round $r$ and a winner $j \in \winners[r]$ such that $p_j$ invoked $\RBcast(j, \texttt{PROP}, S, r)$ with $m \in S$.
|
|
\end{lemma}
|
|
|
|
\begin{proof}
|
|
Let $p_i$ be a correct process that invokes $\ABbroadcast(m)$. By the algorithm, $p_i$ adds $m$ to $\unordered$. Consequently, $\unordered \setminus \ordered \neq \emptyset$, and the process enters the main loop and eventually invokes $\RBcast(i, \texttt{PROP}, S_i^{(1)}, r_1)$ for some round $r_1$ where $m \in S_i^{(1)}$.
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|
|
|
We distinguish two cases:
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|
|
|
\textbf{Case 1: $p_i$ becomes a winner in round $r_1$.}
|
|
If $p_i$ is elected as a winner for round $r_1$ (i.e., $i \in \winners[r_1]$), then the claim holds with $r = r_1$ and $j = i$.
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|
|
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\textbf{Case 2: $p_i$ does not become a winner in round $r_1$.}
|
|
If $p_i$ is not elected as a winner, by the properties of Reliable Broadcast (Bracha's specification), at least one correct process $p_{i_1}$ will eventually receive the reliable delivery of the $\texttt{PROP}$ message from $p_i$ for round $r_1$. Process $p_{i_1}$ adds all messages from $S_i^{(1)}$ to its $\unordered$ set at line \ref{alg:append}. In particular, $m$ is now in $p_{i_1}$'s $\unordered$ set.
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|
|
In round $r_2 > r_1$ (and any subsequent round), process $p_{i_1}$ computes $S_{i_1}^{(2)} = \unordered \setminus \ordered$ and invokes $\RBcast(i_1, \texttt{PROP}, S_{i_1}^{(2)}, r_2)$ with $m \in S_{i_1}^{(2)}$.
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|
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This process repeats: either $p_{i_1}$ becomes a winner and the claim holds, or another correct process receives $p_{i_1}$'s proposal and includes $m$ in its own proposal.
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|
|
|
Since $n > 3f$, we have $n - t > t + 1 > f$. By the pigeonhole principle, there eventually exists a round $r$ where at least $n - t$ distinct processes have proposed sets containing $m$. Since more than $2f$ processes have proposed $m$, at least one of them must be correct and be elected as a winner. Therefore, there exist a round $r$ and a winner $j \in \winners[r]$ such that $m \in S_j$ was broadcast by $p_j$.
|
|
\end{proof}
|
|
|
|
\begin{lemma}[Eventual Closure]\label{lem:eventual-closure}
|
|
For any correct process $p_i$, if $\unordered \setminus \ordered \neq \emptyset$ and if $r$ is the highest closed round observed by $p_i$, then eventually round $r+1$ will be closed with $|\winners[r+1]| \geq n - t$.
|
|
\end{lemma}
|
|
|
|
\begin{lemma}[Total Order]\label{lem:total-order}
|
|
For any two correct processes $p_i$ and $p_j$, the sequence $\ordered$ maintained locally by $p_i$ and the sequence maintained by $p_j$ contain the same messages in the same order, provided that both have reached the same set of closed rounds.
|
|
\end{lemma}
|
|
|
|
\begin{theorem}
|
|
The algorithm implements a BFT Atomic Reliable Broadcast.
|
|
\end{theorem}
|