Beginning of group encryption part
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chap-GE-LWE.tex
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\section{Syntax and Definitions of Group Encryption} \label{GE-model}
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We use the syntax and the security model of Kiayias, Tsiounis and Yung \cite{KTY07}.
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The group encryption (\textsf{GE}) primitive involves a sender, a verifier, a group manager~(\textsf{GM}) that manages the group of receivers and an opening
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authority~(\textsf{OA}) which is capable of identifying ciphertexts' recipients.
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In the syntax of \cite{KTY07}, a $\GE$ scheme is specified by the description of a
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relation $R$ as well as a tuple
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$\GE=\bigl(\mathsf{SETUP},\mathsf{JOIN},\langle
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\mathcal{G}_r,R,\mathsf{sample}_{R}
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\rangle,\mathsf{ENC},\mathsf{DEC},\mathsf{OPEN},\langle
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\mathcal{P},\mathcal{V} \rangle \bigr)$ of algorithms or protocols.
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In details, $\mathsf{SETUP}$ is a set of initialization procedures that all take (implicitly or explicitly) a security parameter $1^\lambda$ as input. We call them
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$\mathsf{SETUP}_{\mathsf{init}}(1^\lambda)$,
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$\mathsf{SETUP}_{\mathsf{GM}}(\param)$ and
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$\mathsf{SETUP}_{\mathsf{OA}}(\param)$. The first one of these procedures
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generates a set of public parameters $\param$ (like the KTY construction \cite{KTY07}, we rely on a common reference string even when using interaction between
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provers and verifiers). The latter two procedures are used to produce key pairs
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$(\pk_{\GM},\sk_{\GM})$, $(\pk_{\OA},\sk_{\OA})$ for the $\GM$ and the
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$\OA$. In the following, $\param$ is incorporated in the inputs of all algorithms although we sometimes omit to explicitly write it.
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$\mathsf{JOIN}=(\mathsf{J}_{\mathsf{user}},\mathsf{J}_{\GM})$ is an interactive protocol between the $\GM$ and the prospective user.
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After the execution of $\mathsf{JOIN}$, the $\GM$ stores the public key $\pk$ and its certificate $\crt_{\pk}$ in a public directory
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$\mathsf{database}$.
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As in \cite{KY05}, we will restrict this
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protocol to have minimal interaction and consist of only two messages: the first one is the user's public key $\pk$ sent by $\mathsf{J}_{\mathsf{user}}$ to $\mathsf{J}_{\GM}$
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and the latter's response is a certificate $\crt_{\pk}$ for $\pk$ that makes the user's group membership effective. We do not require the user to prove
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knowledge of his private key $\sk$ or anything else about it. In our construction, valid keys will be publicly recognizable and users will not have to prove
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their validity. By avoiding proofs of knowledge of private keys, the security proof never has to
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rewind the adversary to extract those private keys, which allows supporting concurrent joins as
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advocated by Kiayias and Yung \cite{KY05}. If applications demand it, it is possible to add
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proofs of knowledge of private keys in a modular way but our security proofs do not require
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rewinding the adversary in executions of $\mathsf{JOIN}$. \\
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\indent
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Algorithm $\mathsf{sample}_{R}$ allows sampling pairs $(x,w)\in R$ (made of a public value $x$ and a witness $w$) using keys $(\pk_{R},\sk_{R})$ produced by
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$\mathcal{G}_r(1^\lambda)$ which samples public/secret parameters for the relation $R$. Depending on the relation, $\sk_{R}$ may be the empty string (as in the scheme \cite{KTY07} and ours which both involve publicly samplable relations). The testing procedure $R(x,w)$ uses $\pk_{R}$ to
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return $1$ whenever $(x,w)\in R$. To encrypt a witness $w$ such that $(x,w) \in R$ for some public $x$, the sender fetches the pair $(\pk,\crt_{\pk})$
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from $\mathsf{database}$ and runs the randomized encryption algorithm. The latter takes as input $w$, a label $L$, the receiver's pair $(\pk,\crt_{\pk})$ as
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well as public keys $\pk_{\GM}$ and $\pk_{\OA}$. Its output is a ciphertext
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$\Psi \leftarrow \mathsf{ENC}(\pk_{\GM},\pk_{\OA},\pk,\crt_{\pk},w,L)$.
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On input of the same elements, the certificate $\crt_{\pk}$, the ciphertext $\Psi$ and the random coins $coins_{\Psi}$ that were used to produce $\Psi$, the
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non-interactive algorithm $\mathsf{PP}$ generates a proof $\pi_{\Psi}$ that there exists a certified receiver whose public key was registered in $\mathsf{database}$ and
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who is able to decrypt $\Psi$ and obtain a witness $w$ such that $(x,w) \in R$. The verification algorithm $\mathcal{V}$ takes as input $\Psi$, $\pk_{\GM}$,
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$\pk_{\OA}$, $\pi_{\Psi}$ and the description of $R$ and outputs $0$ or $1$. Given $\Psi$, $L$ and the receiver's private key $\sk$, the output of
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$\mathsf{DEC}$ is either a witness $w$ such that $(x,w) \in R$ or a rejection symbol $\bot$. Finally,
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$\mathsf{OPEN}$ takes as input a ciphertext/label pair $(\Psi,L)$ and the OA's secret key $\sk_{\OA}$ and returns a receiver's public key $\pk$.\\
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\indent
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The model of \cite{KTY07} considers four properties termed correctness, message security, anonymity and soundness.
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In the security definitions, stateful oracles capture the adversary's
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interaction with the system. In the soundness game, the KTY model requires
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that pk belongs to the language of valid public keys. Here, we are implicitly assuming that the space
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of valid public keys is dense (all matrices are valid keys, as is the case in our scheme).
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In the upcoming definitions, we sometimes use the notation
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\[ \langle \mathsf{output}_A |\mathsf{output}_B \rangle \allowbreak \leftarrow \langle A(\mathsf{input}_A),B(\mathsf{input}_B) \rangle (\mathsf{common\textrm{-}input}) \]
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to denote the execution of a protocol between $A$ and $B$ obtaining their own outputs from their respective inputs.
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\medskip
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\paragraph{Correctness.}
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The correctness property
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requires that the following experiment returns $1$ with overwhelming
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probability.
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\begin{center}
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\procedure{Experiment $\Expt^{\mathrm{correctness}}(\lambda)$}{
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\mathsf{param} \leftarrow
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\mathsf{SETUP}_{\mathsf{init}}(1^\lambda); (\pk_{R},\sk_{R})
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\gets \mathcal{G}_r (\lambda); (x,w) \leftarrow \mathsf{sample}_{R}
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(\pk_{R},\sk_{R}); \\
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(\pk_{\GM},\sk_{\GM}) \leftarrow \mathsf{SETUP}_{\GM}(\mathsf{param}); (\pk_{\OA},\sk_{\OA}) \leftarrow \mathsf{SETUP}_{\OA}(\mathsf{param}); \\
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\langle \pk,\sk,\crt_{\pk} | \pk,\crt_{\pk} \rangle \leftarrow \langle \mathsf{J}_{\mathsf{user}},\mathsf{J}_{\GM}(\sk_{\GM}) \rangle (\pk_{\GM}); \\
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\Psi \leftarrow
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\mathsf{ENC}(\pk_{\GM},\pk_{\OA},\pk,\crt_{\pk},w,L);\\
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\pi_{\Psi} \leftarrow \mathcal{P}(\pk_{\GM},\pk_{\OA},\pk,\crt,w,L,\Psi,coins_{\Psi}); \\
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\pcif \bigl( (w \neq \mathsf{DEC}(\sk,\Psi,L)) \vee ( \pk \neq
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\mathsf{OPEN}(\sk_{\OA},\Psi,L )) \\ \quad \qquad \vee (\mathcal{V}(\Psi,L,\pi_{\Psi},\pk_{\GM},\pk_{\OA})=0)
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\bigr) \pcthen\\
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\pcind \pcreturn 0\\
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\pcelse\\
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\pcind \pcreturn 1;
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}
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\end{center}
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\paragraph{Message Secrecy.}
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The message secrecy property is defined by an experiment where the adversary has access to oracles
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that may be stateful (and maintain a state across queries) or
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stateless:
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%These oracles are the following:
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\begin{itemize}
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\item[-] $\mathsf{DEC}(\sk)$: is a stateless oracle for the user decryption function
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$\mathsf{DEC}$. When this oracle is restricted not to decrypt a
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ciphertext-label pair $(\Psi,L)$, we denote it by
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$\mathsf{DEC}^{\neg \langle \Psi, L \rangle}$.
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\item[-] $\mathsf{CH}_{\mathsf{ror}}^b(\lambda,\pk,w,L)$: is a
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real-or-random challenge oracle which is called \textit{once}. It
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returns $(\Psi,coins_{\Psi})$ such that $\Psi \leftarrow
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\mathsf{ENC}(\pk_{\GM},\pk_{\OA},\pk,\crt_{\pk},w,L)$ if $b=1$
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whereas, if $b=0$, $\Psi \leftarrow
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\mathsf{ENC}(\pk_{\GM},\pk_{\OA},\pk,\crt_{\pk},w',L)$ encrypts a
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random plaintext of
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length $O(\lambda)$ uniformly sampled in the plaintext space. In both cases, $coins_{\Psi}$ denote the random
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coins used to generate $\Psi$.
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\item[-]
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$\mathsf{PROVE}_{\mathsf{PP},\mathsf{PP}'}^b(\pk_{\GM},\pk_{\OA},\pk,\crt_{\pk},\pk_{R},x,w,\Psi,L,coins_{\Psi})$:
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is a stateful oracle that can be invoked a polynomial number times. If $b=1$, it replies by running the real prover $\mathsf{PP}$ on the inputs to
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create an actual proof $\pi_{\Psi}$. If $b=0$, the oracle runs a
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simulator $\mathsf{PP}'$ that uses the same inputs as $\mathsf{PP}$ except witness
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$w,coins_{\Psi}$ and generates a simulated proof.
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\end{itemize}
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These oracles are used in an experiment where the adversary controls
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the $\GM$, the $\OA$ and all members except the honest receiver. The
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adversary $\adv$ embodies the dishonest $\GM$ that certifies the honest
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receiver in an execution of $\mathsf{JOIN}$. It is granted access to an oracle $\mathsf{DEC}$ which decrypts on behalf of that receiver. In the
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challenge phase, it transmits a state information $\mathsf{aux}$ to itself and invokes the challenge oracle for a label and a
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pair $(x,w) \in R$ of its choice. After the challenge phase, it
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can also query the $\mathsf{PROVE}$ oracle many times
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and finally attempts to guess the challenger's bit $b$.\\
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\indent As pointed out in \cite{KTY07,CLY09}, designing an efficient
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simulator $\mathsf{PP}'$ (for executing $\mathsf{PROVE}_{\mathsf{PP},\mathsf{PP}'}^b(.)$
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when $b=0$) is part of the security proof.
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\begin{definition} \label{security-def}
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A $\GE$ scheme satisfies \textit{message security}
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if, for any PPT adversary $\adv$, the experiment below returns $1$
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with probability at most $1/2 + \mathsf{negl}(\lambda)$.
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\begin{center}
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\procedure{Experiment $\Expt_{\adv}^{\mathrm{sec}}(\lambda)$}{
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\param \leftarrow \mathsf{SETUP}_{\mathsf{init}}(1^\lambda); (\mathsf{aux},\pk_{\GM},\pk_{\OA}) \leftarrow \adv(\param); \\
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\langle \pk,\sk,\crt_{\pk} | \mathsf{aux} \rangle
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\leftarrow \langle \mathsf{J}_{\mathsf{user}},\adv(\mathsf{aux}) \rangle
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(\pk_{\GM}); \\
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(\mathsf{aux},x,w,L,\pk_{R}) \leftarrow
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\adv^{\mathsf{DEC}(\sk,.)}(\mathsf{aux});\\
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\pcif (x,w) \not\in R \pcthen\\
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\pcind\pcreturn 0;\\
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b\sample \{0,1\}; ( \Psi,coins_{\Psi} ) \leftarrow \mathsf{CH}_{\mathsf{ror}}^b(\lambda,\pk,w,L) ; \\
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b' \leftarrow
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\adv^{\mathsf{PROVE}_{\mathsf{PP},\mathsf{PP}'}^b(\pk_{\GM},\pk_{\OA},\pk,\crt_{\pk},\pk_{R},x,w,\Psi,L,coins_{\Psi}),\mathsf{DEC}^{\neg
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\langle \Psi, L \rangle}(\sk,.)}(\mathsf{aux},\Psi) ; \\
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\pcif b=b' \pcthen\\ \pcind\pcreturn 1 \\\pcelse\\ \pcind \pcreturn 0;
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}
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\end{center}
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\end{definition}
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\paragraph{Anonymity.}
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In the experiment modeling the anonymity property, the adversary
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controls the entire system except the opening authority and two well-behaved users.
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The challenger thus introduces two honest users' public keys $\pk_0,\pk_1$ in $\mathsf{database}$ and thus obtains certificate for both $\pk_0,\pk_1$ from the adversarially-controlled $\GM$.
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For a pair $(x,w) \in R$ of its choice, the adversary obtains an encryption of $w$ under $\pk_b$ for some $b\in \bit$ chosen by the challenger.
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The adversary is provided with decryption oracles w.r.t. both keys $\pk_0,\pk_1$. In addition, it has the following oracles at disposal:
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\begin{itemize}
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\item[-] $\mathsf{CH}_{\mathsf{anon}}^b(\pk_{\GM},\pk_{\OA},\pk_0,\pk_1,w,L)$: is a
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challenge oracle that is only queried once by the adversary. It
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returns a pair $(\Psi,coins_{\Psi})$ consisting of a ciphertext
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$\Psi \leftarrow
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\mathsf{ENC}(\pk_{\GM},\pk_{\OA},\pk_b,\crt_{\pk_b},w,L)$ and the
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coin tosses $coins_{\Psi}$ that were used to generate $\Psi$.
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\item[-]
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$\mathsf{USER}(\pk_{\GM})$: is a stateful oracle that obtains certificates from the adversary by simulating two
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executions of $\mathsf{J}_{\mathsf{user}}$ to introduce two honest users
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in the group. It uses a string $\mathsf{keys}$ where the outputs $(\pk_0,\sk_0,\crt_{\pk_0})$, $(\pk_1,\sk_1,\crt_{\pk_1})$ of honest users
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are written as long as the adversarially-supplied certificates $\{\crt_{\pk_d}\}_{d=0}^1$ are valid w.r.t. $\pk_{\GM}$ (i.e., invalid certificates are ignored
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by the oracle and no entry is introduced in $\mathsf{keys}$ for them).
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\item[-]
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$\mathsf{OPEN}(\sk_{\OA},.)$: is a stateless oracle that simulates
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the opening algorithm and, on input of a $\GE$
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ciphertext, returns the receiver's public key.
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\end{itemize}
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The reason why
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the $\mathsf{USER}$ oracle is needed is that both honest users' public keys $\pk_0, \pk_1$ must have been properly
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certified by the adversarially-controlled $\mathsf{GM}$ before the challenge phase because the adversary subsequently obtains
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proofs generated using $(\pk_b,\crt_{\pk_b})$.
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\begin{definition} \label{anonymity-def}
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A $\GE$ scheme satisfies \textit{anonymity} if, for any PPT adversary $\adv$, the experiment below returns $1$
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with a probability not exceeding $1/2 + \mathsf{negl}(\lambda)$.
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\begin{center}
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\procedure{Experiment $\Expt_{\adv}^{\mathrm{anon}}(\lambda)$}{
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\param \leftarrow
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\mathsf{SETUP}_{\mathsf{init}}(1^\lambda); (\pk_{\OA},\sk_{\OA})
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\leftarrow \mathsf{SETUP}_{\mathsf{OA}}( \param); \\
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(\mathsf{aux},\pk_{\GM}) \leftarrow \adv(\param,\pk_{\OA});
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\mathsf{aux} \leftarrow
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\adv^{\mathsf{USER}(\pk_{\GM}),\mathsf{OPEN}(\sk_{\OA},.)}
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(\mathsf{aux}) ; \\
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\pcif \mathsf{keys} \neq (\pk_0,\sk_0,\crt_{\pk_0},\pk_1,\sk_1,\crt_{\pk_1})(\mathsf{aux}) \pcthen\\
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\pcreturn 0;\\
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(\mathsf{aux},x,w,L,\pk_{R}) \leftarrow
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\adv^{\mathsf{OPEN}(\sk_{\OA},.),
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\mathsf{DEC}(\sk_0,.),\mathsf{DEC}(\sk_1,.)}(\mathsf{aux}); \\
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\pcif (x,w) \not\in R \pcthen\\
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\pcind \pcreturn 0; \\
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b\sample \{0,1\}; ( \Psi,coins_{\Psi} ) \leftarrow \mathsf{CH}_{\mathsf{anon}}^b(\pk_{\GM},\pk_{\OA},\pk_0,\pk_1,w,L) ; \\
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b' \leftarrow
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\adv^{\mathcal{P}(\pk_{\GM},\pk_{\OA},\pk_b,\crt_{\pk_b},x,w,\Psi,L,coins_{\Psi},}
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\\ ^{\mathsf{OPEN}^{\neg \langle \Psi,L \rangle
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}(\sk_{\OA},.),\mathsf{DEC}^{\neg \langle \Psi, L
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\rangle}(\sk_0,.),\mathsf{DEC}^{\neg
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\langle \Psi, L \rangle}(\sk_1,.))}(\mathsf{aux},\Psi) ; \\
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\pcif b=b' \pcthen\\
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\pcind \pcreturn 1\\
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\pcelse\\
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\pcind \pcreturn 0;
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}
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\end{center}
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\end{definition}
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\paragraph{Soundness.}
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Here, the adversary creates the group of receivers by interacting with the honest GM.
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Its goal is to produce a ciphertext $\Psi$ and a convincing proof
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that $\Psi$ is valid w.r.t. a relation $R$ of its choice but
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either: (1) The opening of $\Psi$ reveals a receiver's public key $\pk$ that
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does not belong to any group member; (2) The output $\pk$ of
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$\mathsf{OPEN}$ is not a valid public key (\textit{i.e.}, $\pk \not\in
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\mathcal{PK}$, where $\mathcal{PK}$ is the language of valid public keys); (3) The ciphertext $C$ is not in the space
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$\mathcal{C}^{x,L,\pk_{R},\pk_{\GM},\pk_{\OA},\pk}$ of valid
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ciphertexts. This notion is formalized by a game where the adversary
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is given access to a user registration oracle
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$\mathsf{REG}(\sk_{\GM},.)$ that simulates $\mathsf{J}_{\GM}$. This oracle
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maintains a list $\mathsf{database}$ where registered public keys and
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their certificates are stored.
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\begin{definition} \label{soundness-def}
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A $\GE$ scheme is \textit{sound} if, for any PPT adversary $\adv$, the experiment below returns $1$
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with negligible probability.
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\begin{center}
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\procedure{Experiment $\Expt_{\adv}^{\mathrm{soundness}}(\lambda)$}{
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\param \leftarrow
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\mathsf{SETUP}_{\mathsf{init}}(1^\lambda); (\pk_{\OA},\sk_{\OA})
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\leftarrow \mathsf{SETUP}_{\OA}( \param); \\
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(\pk_{\GM},\sk_{\GM}) \leftarrow \mathsf{SETUP}_{\GM}( \param); \\
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(\pk_{R},x,\Psi,\pi_{\Psi},L,\mathsf{aux}) \leftarrow
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\adv^{\mathsf{REG}(\sk_{\GM},.)}(\param,\pk_{\GM},\pk_{\OA},\sk_{\OA});
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\\
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\pcif \mathcal{V}( \Psi,L,\pi_{\Psi},\pk_{\GM},\pk_{\OA})=0 \pcthen\\
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\pcind \pcreturn 0; \\
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\pk \leftarrow \mathsf{OPEN}(\sk_{\OA},\Psi,L); \\
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\pcif \big( (\pk \not\in \mathsf{database} ) \vee (\pk \not \in \mathcal{PK})
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\vee (\Psi \not \in
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\mathcal{C}^{x,L,\pk_{R},\pk_{\GM},\pk_{\OA},\pk}) \big)\pcthen\\
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\pcind \pcreturn 1\\
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\pcelse \\
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\pcind \pcreturn 0;
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}
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\end{center}
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\end{definition}
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The model of Kiayias \textit{et al.} \cite{KTY07} requires
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that $\pk$ belongs to the language of valid public keys, so that the adversary is considered to defeat the soundness property when
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$(\Psi,L)$ opens to a key outside the language (i.e.,
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$\pk \not \in \mathcal{PK}$). In our scheme, we will assume that the space
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of valid public keys is dense in that all matrices of a given dimension are valid public keys, which have an underlying private key.
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We nevertheless use the same definition as \cite{KTY07} in order to emphasize that we are not relaxing the model in any way.
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\section{Building Blocks}
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\subsection{The Agrawal-Boneh-Boyen IBE Scheme} \label{ap:ABB-IBE}
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\subsubsection{Identity-Based Encryption.} \label{ap:IBE}
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An IBE scheme is a tuple of efficient algorithms $(\mathsf{Setup}, \mathsf{Extract}_\mathsf{PP}, \mathsf{Encrypt}_\mathsf{PP},$ $\mathsf{Decrypt}_\mathsf{PP})$ such that
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\begin{description}
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\item[\textsf{Setup}$(1^\lambda)$:] On security parameter $\lambda$, this algorithm outputs public parameters $\mathsf{PP}$ and a master secret key $\textsf{msk}$.
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\item[\textsf{Extract}$_\mathsf{PP}(\textsf{msk}, \ID)$:] Takes as input a master secret key $\textsf{msk}$ and an identity $\ID$ and outputs a secret key $\sk_\ID$.
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\item[\textsf{Encrypt}$_\mathsf{PP}(\ID, M)$:] Given an identity $\ID$ and a message $M$, it outputs a ciphertext $C$.
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\item[\textsf{Decrypt}$_\mathsf{PP}(\sk_\ID, C)$:] Given a secret key $\sk_\ID$ and a ciphertext $C$, outputs either a decryption error symbol $\bot$, or a message $M$.
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\end{description}
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\noindent Correctness requires that, for any pair $(\mathsf{PP}, \textsf{msk}) \gets \Setup(1^\lambda)$, any $\ID$ and any message $M$, we have
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$\mathsf{Decrypt}_\mathsf{PP}\bigl(\textsf{Extract}_\mathsf{PP}(\textsf{msk}, \ID), \mathsf{Encrypt}_\mathsf{PP}(\ID, M)\bigr) = M.$
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Our proofs rely on the semantic security of the scheme against selective adversaries (\textsf{IND-sID-CPA})
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but also on the stronger property of ciphertext pseudo-randomness. %in Lemma~\ref{ABB-deux}.
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Informally, this notions demands that the adversary be unable to distinguish an
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encryption of a message of its choice from a random element of the ciphertext space $\mathcal{C}$. Notice that this property implies \textsf{IND-sID-CPA} security.
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\begin{definition}
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\label{de:pseudorand-cipher}
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An IBE scheme has pseudo-random-ciphertexts if no PPT adversary $\adv$ with access to private key extraction oracle \textsf{Extract$_\mathsf{PP}(\textsf{msk}, \cdot)$} has non-negligible advantage
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$ \advantage{\mathrm{ROR}}{\adv}{\lambda} = | \Pr\bigl[ \mathbf{Expt}_{\adv}^\mathrm{ROR} = 1 \bigr] - \frac 1 2 | $ in the game described in Figure~\ref{fig:expt-ror}
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\begin{figure}
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\centering
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\procedure{Experiment $\Expt^{\mathrm{ROR}}_{\adv}(\lambda)$}{
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\ID^\star \gets \adv(\textsf{id}, \lambda); (\mathsf{PP}, \textsf{msk}) \gets \mathsf{Setup}(1^\lambda);~\\
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M \gets \adv^{\mathsf{Extract}_\mathsf{PP}(\textsf{msk}, \cdot)}_\textsf{Ch}(\mathsf{PP});\\
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b \sample U(\bit);\\
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\pcif b = 1 \pcthen\\
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\pcind C^\star \gets \mathsf{Encrypt}_\mathsf{PP}(M, \ID^\star) \\
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\pcelse\\
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\pcind C^\star \gets U(\mathcal{C});\\
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b' \gets \adv^{\mathsf{Extract}_\mathsf{PP}(\textsf{msk}, \cdot)}(\textsf{guess},C^\star);\\
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\pcif b = b' \pcthen\\
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\pcind \pcreturn 1\\
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\pcelse \\
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\pcind \pcreturn 0
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|
}
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\caption{Security experiment for the pseudo-random-ciphertext property for an IBE}
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\label{fig:expt-ror}
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|
\end{figure}
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|
\end{definition}
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\subsubsection{The ABB System.} \label{ap:ABB-desc}
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Agrawal, Boneh and Boyen described~\cite{ABB10} a compact IBE scheme in the standard model which allows encrypting multi-bit messages.
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|
%In
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%the description hereunder, algorithms \textsf{Extract}, \textsf{Encrypt} and \textsf{Decrypt} have implicit access to the public param
|
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|
%eters $\mathsf{PP}$ defined in the \textsf{Setup} algorithm.
|
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|
|
||||||
|
\begin{description}
|
||||||
|
\item[\textsf{Setup}$(1^\lambda)$:] Given a security parameter $\lambda$, choose parameters
|
||||||
|
$q, n, \sigma, \alpha$ and define $k =\lfloor \log q \rfloor$, $\bar{m}= nk$, $m = 2 \bar{m}$ and choose a noise distribution $\chi$ for $\LWE$.
|
||||||
|
\begin{enumerate}
|
||||||
|
\item Compute $(\bar{\mathbf A}, \mathbf{T}_{\bar{\mathbf{A}}}) \gets \TrapGen(1^n, 1^m, q)$.
|
||||||
|
\item Define $\mathbf{G} = \mathbf{I}_n \otimes [1|2|\ldots |2^{k-1}] \in \ZZ_q^{n \times \bar{m}}$. Sample matrices $\mathbf B \sample U(\ZZ_q^{ n \times \bar{m}}) $,
|
||||||
|
$ \mathbf U \sample U(\Zq^{n \times m})$.
|
||||||
|
\item Let $\mathsf{FRD}: \Zq^n \to \Zq^{n \times n}$ be the full-rank difference mapping from~\cite{ABB10}.
|
||||||
|
\end{enumerate} Output
|
||||||
|
$
|
||||||
|
\mathsf{PP}= \bigl(\bar{\mathbf A}, \mathbf B, \mathbf U \bigr)$ and $\textsf{msk} = \mathbf{T}_{\bar{\mathbf A}}$.
|
||||||
|
|
||||||
|
\item[\textsf{Extract}$_\mathsf{PP}(\textsf{msk}, \ID)$:] Given $\textsf{msk} = \mathbf{T}_{\bar{\mathbf A}}$ and an identity $\ID \in \Zq^n$, do as follows: \smallskip
|
||||||
|
\begin{enumerate}
|
||||||
|
\item Define the matrix $\mathbf B_\ID = \mathbf B + \mathsf{FRD}(\ID) \cdot \mathbf G \in \Zq^{n \times \bar{m}}$.
|
||||||
|
%\item Use $\mathbf T_A$ to compute a delegated basis $\mathbf T_\ID$ for the dual lattice of the matrix $\mathbf B_{\mathbf A, \ID} = \left[ \mathbf A \mid \mathbf B_\ID \right]$.
|
||||||
|
\item Let $\mathbf B_{\mathbf A, \ID} = \left[ \mathbf A \mid \mathbf B_\ID \right] \in \ZZ_q^{n \times (m + \bar{m})}$, use $\mathbf T_A$ to compute a delegated basis $\mathbf T_\ID$ for the lattice $\Lambda^\perp(\mathbf B_{\mathbf A, \ID})$.
|
||||||
|
\item Use $\mathbf T_\ID$ to sample a small-norm matrix $\mathbf E_\ID \in \ZZ^{(m+ \bar{m}) \times m}$ satisfying the equality $\mathbf B_{\mathbf A, \ID} \cdot \mathbf E_\ID = \mathbf U \bmod q$.
|
||||||
|
\item Output $\sk_\ID = \mathbf E_\ID \in \ZZ^{(m+ \bar{m}) \times m}$. \smallskip \smallskip
|
||||||
|
\end{enumerate}
|
||||||
|
\item[\textsf{Encrypt}$_\mathsf{PP}(\ID,\mathbf m)$:] Given an identity $\ID$ and a message $\mathbf m \in \bit^m$, \smallskip
|
||||||
|
\begin{enumerate}
|
||||||
|
\item Compute the matrix $\mathbf B_\ID = \mathbf B + \mathsf{FRD}(\ID) \cdot \mathbf G \in \Zq^{n \times \bar{m}}$.
|
||||||
|
Sample vectors $\mathbf s \sample U(\Zq^n), \mathbf x, \mathbf y \sample \chi^m$, $\mathbf R \sample D_{\ZZ,\sigma}^{m \times \bar{m}}$ and compute
|
||||||
|
$\mathbf z = \mathbf R^\top \cdot \mathbf y \in \ZZ^m$.
|
||||||
|
\item Compute
|
||||||
|
\begin{equation} \label{eq:ABB-c}
|
||||||
|
\begin{cases}
|
||||||
|
\mathbf c^{(1)} = \bar{\mathbf A}^\top \cdot \mathbf s + \mathbf y \bmod q,\\
|
||||||
|
\mathbf c^{(2)} = \mathbf B_\ID^\top \cdot \mathbf s + \mathbf z \bmod q,\\
|
||||||
|
\mathbf c^{(3)} = \mathbf U^\top \cdot \mathbf s + \mathbf x + \mathbf m \cdot \left\lfloor \dfrac{q}{2} \right\rfloor.
|
||||||
|
\end{cases}
|
||||||
|
\end{equation}
|
||||||
|
\item Output $\mathbf c = \bigl(\mathbf c^{(1)},\mathbf c^{(2)},\mathbf c^{(3)}\bigr) \in \ZZ_q^m \times \ZZ_q^{\bar{m}} \times \ZZ_q^m$. \smallskip \smallskip
|
||||||
|
\end{enumerate}
|
||||||
|
\item[\textsf{Decrypt}$_\mathsf{PP}(\sk_\ID, \mathbf c)$:] Given $\sk_\ID = \mathbf E_\ID$ and $\mathbf c=\bigl(\mathbf c^{(1)},\mathbf c^{(2)},\mathbf c^{(3)}\bigr) \in \ZZ_q^m \times \ZZ_q^{\bar{m}} \times \ZZ_q^m$, compute and output
|
||||||
|
% \begin{equation} \label{eq:ABB-dec}
|
||||||
|
$ \mathbf m' = \left\lfloor \left( \mathbf c^{(3)} - \mathbf E_\ID \cdot \begin{bmatrix} \mathbf c^{(1)} \\ \mathbf c^{(2)} \end{bmatrix} \right) \cdot \left\lfloor \dfrac{q}{2} \right\rfloor^{-1} \right\rceil \in \bit^m .$
|
||||||
|
% \end{equation}
|
||||||
|
\end{description}
|
||||||
|
\smallskip
|
||||||
|
|
||||||
|
|
||||||
|
|
||||||
|
\begin{theorem}[{\cite[Th. 23]{ABB10}}] \label{ABB-pseudorand-prop}
|
||||||
|
The ABB IBE scheme has pseudo-random ciphertexts if the $\LWE_{n,q,\chi}$ assumption holds.
|
||||||
|
\end{theorem}
|
||||||
|
|
||||||
|
|
||||||
|
%**************************************************
|
||||||
|
|
||||||
@ -19,6 +19,7 @@
|
|||||||
\newcommand{\QANIZK}{\textsf{QA-NIZK}\xspace}
|
\newcommand{\QANIZK}{\textsf{QA-NIZK}\xspace}
|
||||||
\newcommand{\PKE}{\textsf{PKE}\xspace}
|
\newcommand{\PKE}{\textsf{PKE}\xspace}
|
||||||
\newcommand{\OT}{\textsf{OT}\xspace}
|
\newcommand{\OT}{\textsf{OT}\xspace}
|
||||||
|
\newcommand{\GE}{\textsf{GE}\xspace}
|
||||||
%% Common
|
%% Common
|
||||||
\newcommand{\Setup}{\ensuremath{\mathsf{Setup}}\xspace}
|
\newcommand{\Setup}{\ensuremath{\mathsf{Setup}}\xspace}
|
||||||
\newcommand{\Keygen}{\ensuremath{\mathsf{Keygen}}\xspace}
|
\newcommand{\Keygen}{\ensuremath{\mathsf{Keygen}}\xspace}
|
||||||
|
|||||||
Loading…
Reference in New Issue
Block a user