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An INDCCA2 secure postquantum encryption scheme and a secure cloud storage use case
Humancentric Computing and Information Sciences volumeÂ 9, ArticleÂ number:Â 32 (2019)
Abstract
Codebased public key encryption (PKE) is a popular choice to achieve postquantum security, partly due to its capability to achieve fast encryption/decryption. However, codebased PKE has larger ciphertext and public key sizes in comparison to conventional PKE schemes (e.g., those based on RSA). In 2018, Lau and Tan proposed a new rank metric codebased PKE scheme, which has smaller public key and ciphertext sizes compared to other codebased PKE schemes. They also proved that their scheme achieves INDCPA security, assuming the intractability of the decisional rank syndrome decoding problem. It is known that INDCCA2 security is the strongest and most popular security assurance for PKE schemes. Therefore, in this paper, we obtain a new codebased PKE scheme from Lau and Tanâ€™s scheme, in order to inherit the underlying small public key and ciphertext sizes. However, our new scheme is shown to achieve INDCCA2 security, instead of the weaker INDCPA security. Specifically, the respective public key size and ciphertext size in our new scheme are 15.06 KB and 1.37 KB under 141bit security level, and 16.76 KB and 1.76 KB under 154bit security level. We then present a use case for the proposed scheme, that is for secure cloud storage.
Introduction
With rapid advances in Internet and information and communication technologies (ICT; e.g., computation devices, speed of communication and device processors), our society is becoming increasingly reliant on technologies. This is particularly true for technologically advanced countries. This has also given rise to a significant increase in the amount of data generated, and needed to be processed and stored (e.g., using distributed storage servers such as cloud server). Cloud storage service providers such as Amazonâ€™s S3 offer users public cloud storage services, where users can backup, access and/or process their data anywhere from some computing devices (e.g., Android and iOS devices) connected to the Internet. Popularity of such services is also driven by financial costs, as it may be cheaper to pay for what you use than building, maintaining and securing oneâ€™s local storage services. One downside, however, is the potential data privacy and integrity risks since data is now outsourced to some public cloud service providers that may be semitrusted. Thus, massive techniques are proposed for ensuring the data security and decreasing the useless data such as semisupervised learning, spammer detection framework and cryptographic protocol [1,2,3,4].
Generally, a cloud storage system should ensure availability (e.g., for any legitimate customer, (s)he can reach his/her uploaded data from some Internetconnected devices), reliability (of user data outsourced to the cloud), efficient retrieval (of user data outsourced to the cloud), data sharing (between authorized users), security (i.e., both confidentiality and integrity [5]), and other features required/stated in the service level agreements (SLAs). Cryptographic schemes, such as proxy reencryption, attributebased encryption, searchable encryption [6,7,8,9], have been designed to achieve several of these features.
Nowadays, most existing cryptographic schemes are numbertheoreticbased [10,11,12], which are not resilient against Shorâ€™s quantum attack algorithm [13, 14] using quantum computers. Codebased public key encryption (PKE) proposed by McEliece [15] in 1978 is widely considered to be postquantum secure [16]. There are two key advantages in McEliecelike codebased PKE schemes, namely: encryption/decryption speed is very fast, and its security relies on hard problems in coding theory which is believed to resist quantum computing attacks. However, such schemes are not widely used in practice, partly due to large public key size (in comparison to classic numbertheoreticbased schemes). Hence, how to design a codebased PKE scheme with small public key size is an area of active research. In 2017, for example, Loidreau proposed a codebased PKE scheme from rank metric [17], which can significantly reduce the size of public keys compared to Hamming metric encryption schemes. More recently in 2018, Lau and Tan introduced a public key generation approach and a new rank metric codebased PKE scheme (hereafter refer to LT scheme) to further reduce the public key size [18]. The LT scheme was also shown to be indistinguishability against chosen plaintext attacks (INDCPA) secure, based on the decisional rank syndrome decoding (DRSD) assumption.
Our contribution
For a PKE scheme, however, the strongest and the most popular security notion is indistinguishability against adaptive chosen ciphertext attacks (INDCCA2). This is the gap we seek to address in this paper. Our contribution in this paper include three aspects. Firstly, we propose a new INDCCA2 secure PKE scheme based on the LT scheme. While the semantic security in our PKE scheme has been improved to be INDCCA2 secure, the public key and ciphertext sizes remains small. Secondly, we give the formal proof of our PKE scheme under the DRSD assumption. Thirdly, we present a use case of our PKE scheme in secure cloud storage.
Paper organization
The rest of the paper is organized as follows. Prior to presenting our scheme, we introduce relevant background materials in "Preliminaries" section and the LT scheme [18] in "Revisiting the LT scheme" section. We then present our INDCCA2 secure codebased PKE scheme in "Our INDCCA2 secure PKE scheme" section, following by its security proof and performance evaluation in "Security proof" andÂ "Efficiency" sections. For example, we show that the public key size in our new scheme is 15.06 KB under 141bit security level which is acceptable in many todayâ€™s applications. We use cloud storage as a use case and explain its potential utility in "A secure cloud storage use case" section. Finally, we conclude this paper in "Conclusion" section.
Preliminaries
In this section, we revisit the definitions, notations and materials relevant to rank metric codes, which is required in the understanding of our proposed scheme. In the rest of this paper, we will use the notations and the functions as we defined in Table 1.
Definition 1
(Linear code and its generator matrix) An \([n, k]_{q}\) linear code \({\mathcal {C}}\) is a linear subspace of \({\mathbb {F}}_q^n\) with dimension k, and a matrix \(G\in {\mathbb {F}}_q^{k\times n}\) is called a generator matrix of \({\mathcal {C}}\) if its row vectors form a basic of \({\mathcal {C}}\).
Definition 2
(Rank metric) Let \({\mathbb {F}}_{q^m}\) be an extension field of \({\mathbb {F}}_q\) for some positive integer m, and \(\beta = (\beta _0, \beta _1, \ldots , \beta _{m1})\) a basis of \({\mathbb {F}}_{q^m}\) over \({\mathbb {F}}_q\). For each vector \({\mathbf {a}} = (a_0, a_1, \ldots , a_{n1}) \in {\mathbb {F}}_{q^m}^n\), we associate it with an \(m \times n\) matrix \(M({\mathbf {a}}) = (a_{ij})_{0\le i \le m1, 0 \le j \le n1}\in {\mathbb {F}}_q^{m\times n}\) s.t. \(a_j = \sum _{i=0}^{m1}m_{ij}\beta _i\), \(j=0, 1, \ldots , n1\). The rank of \({\mathbf {a}}\), denoted by \(\Vert {\mathbf {a}}\Vert\), is defined as the rank of the matrix \(M({\mathbf {a}})\), and the rank distance between any two vectors \({\mathbf {x}}\) and \({\mathbf {y}}\) in \({\mathbb {F}}_{q^m}^n\) is defined by \(\Vert {\mathbf {x}}{\mathbf {y}}\Vert\).
Definition 3
(Circulant matrix) Given a vector \({\mathbf {c}} = (c_0, c_1, \ldots , c_{n1}) \in {\mathbb {F}}_{q^m}^n\), we associate it with an \(n \times n\) circulant matrix \(Rot({\mathbf {c}})=(c_{ij})_{0\le i,j \le n1}\in {\mathbb {F}}_{q^m}^{n\times n}\) satisfying \(c_{ij} = c_{(ij) \mod n}\) for \(0 \le i \le n1\), \(0 \le j \le n1\). That is, we have
Further, we use the notation \(Rot_k({\mathbf {c}})\) to denote a \(k\times n\) matrix consisting of the first k rows of \(Rot({\mathbf {c}})\).
Definition 4
(Decisional rank syndrome decoding (DRSD) problem) For a given full rank matrix \(A\in {\mathbb {F}}_{q^m}^{k\times n}\) and an integer t, we define \({\mathcal {S}}_{A, t, 0}=\{(A, {\mathbf {y}})~~{\mathbf {y}}\xleftarrow {\$}{\mathbb {F}}_{q^m}^n\}\) and \({\mathcal {S}}_{A, t, 1}=\{(A, {\mathbf {xA}}+{\mathbf {e}})~~{\mathbf {x}}\in {\mathbb {F}}_{q^m}^k, {\mathbf {e}}\in {\mathbb {F}}_{q^m}^n \ \text{ and }\ \Vert {\mathbf {e}}\Vert =t \}\).
Input: a full rank matrix \(A\in {\mathbb {F}}_{q^m}^{k\times n}\), an integer t, and a vector \({\mathbf {s}} \in {\mathbb {F}}_{q^m}^n\).
Output: a bit \(b \in \{0,1\}\) s.t. \((A, {\mathbf {s}}) \in {\mathcal {S}}_{A, t, b}\).
Let \({\mathcal {A}}\) be a probabilistic polynomial time (PPT) algorithm for the DRSD problem. With respect to a security parameter \(\lambda\), we define the advantage of \({\mathcal {A}}\) as
In [19], the DRSD problem is proven to be NP hard in the worst case; thus, it is reasonable to assume that \(\text{ Adv }_{{\mathcal {A}}}^{{\text {DRSD}}}(\lambda )\) is negligible with regard to \(\lambda\).
Definition 5
(INDCPA attack) To describe the INDCPA attack we first introduce the power of the adversary, denoted by \({\mathcal {A}}\), in the attack.
Chosen Plaintext Adversary: \({\mathcal {A}}\) can choose the plaintext adaptively and he can get the corresponding ciphertext of the chosen plaintext through the encryption mechanism.
We define an INDCPA attack as an experiment. The experiment runs the key generation algorithm and gets a publicprivate key pair (pk,Â sk). An INDCPA adversary uses pk to get several messageciphertext pairs and stores them in a table. A distinguisher \({\mathcal {D}}\) gets pk and the messageciphertext table. \({\mathcal {D}}\) randomly chooses two messages \(m_0\) and \(m_1\) (not in the messageciphertext table) and sends them to the encryption mechanism. The encryption mechanism randomly chooses a message \(m_b\), \(b\in \{0,1\}\), and sends the corresponding ciphertext \(c_b\) to \({\mathcal {D}}\). \({\mathcal {D}}\) takes as input \((pk, m_0, m_1, c_b)\) and outputs a guess \(b'\) of b. If \(b' = b\), the experiment successes. Otherwise, the experiment failures. The probability that \({\mathcal {D}}\) succeeds in the INDCPA experiment is denoted by \(\Pr _{CPA}\).
A PKE scheme is called INDCPA secure if \(\Pr _{CPA}  \frac{1}{2}\) is negligible.
Revisiting the LT scheme [18]
The LT scheme [18] consists of four PPT algorithms, namely: Setup, KGen, Enc, and Dec which are described as follows.

Setup(\(1^{\lambda }\)): Given a security parameter \(\lambda\), the algorithm chooses a prime power q, a group of integers \((n, m, k, k_0, k_1, t)\) s.t. \(n>k, ~k_0 = \lfloor \frac{k}{2}\rfloor , ~k_1 = k k_0, t \le \lfloor \frac{nk}{2}\rfloor\) according to \(\lambda\). The algorithm then outputs the public parameter \(params = (q, n, m, k, k_0, k_1, t)\).

KGen(params): Give the public parameter params, the algorithm

Chooses a generator matrix \(G \in {\mathbb {F}}_{q^m}^{k \times n}\) of a \([n, k]_{q^m}\) linear code \({\mathcal {C}}\) with error correcting ability t and an efficient decoding algorithm \(\mathcal {DEC}(\cdot )\);

Chooses a vector \({\mathbf {x}} \xleftarrow {\$}{\mathbb {F}}_{q^m}^n\) s.t. \(\Vert {\mathbf {x}}\Vert = n\);

Chooses two invertible matrices \(Q \in {\mathbb {F}}_{q^m}^{k \times k}\) and \(P \in {\mathbb {F}}_q^{n \times n}\) uniformly at random;

Computes \(\hat{G} = QG + Rot_k({\mathbf {x}})P\).
The public/private key pair is (\(pk = (\hat{G}, {\mathbf {x}}),sk = (Q, G, P, \mathcal {DEC}(\cdot ))\)).

The algorithms Enc and Dec are listed in Figs.Â 1 andÂ 2, respectively.
As mentioned previously, the LT scheme was proven to be INDCPA secure under the DRSD assumption in [18].
Our INDCCA2 secure PKE scheme
In this section, we convert the LT scheme into an INDCCA2 secure PKE scheme in which two error vectors come from a random value. This allows the new scheme to avoid the expansion of ciphertext.
Our new INDCCA2 secure PKE scheme consists of the following four algorithms:

Setup(\(1^{\lambda }\)): Given a security parameter \(\lambda\), the algorithm

Chooses a prime number power q, a group of integers \((n, m, k, k_0, k_1, t, t_0, t_1)\) s.t. \(n>k, ~k_0 = \lfloor \frac{k}{2}\rfloor , ~k_1 = k k_0,\ t \le \lfloor \frac{nk}{2}\rfloor\), \(~t_0 = \lfloor \frac{t}{2}\rfloor , ~t_1 = t t_0\);

Chooses a random vector \(\mathbf {con}\) over \({\mathbb {F}}_{q^m}\) with \(Len(\mathbf {con}) \le \frac{1}{2}\ell\), where \(\ell _0=\lfloor \log _{q^m}\left( {\begin{array}{c}n\\ t_0\end{array}}\right) \rfloor\), \(\ell _1=\lfloor \log _{q^m}\left( {\begin{array}{c}n\\ t_1\end{array}}\right) \rfloor\), and \(\ell =\ell _0+\ell _1\);

Outputs the public parameter \(params = (q, n, m, k, k_0, k_1, t, t_0, t_1, \mathbf {con})\).


KGen(params): Give the public parameter \(params= (q, n, m, k, k_0, k_1, t, t_0, t_1, \mathbf {con})\), the algorithm

Chooses a generator matrix \(G \in {\mathbb {F}}_{q^m}^{k \times n}\) of an \([n, k]_{q^m}\) linear code \({\mathcal {C}}\) with error correcting ability t and an efficient decoding algorithm \(\mathcal {DEC}(\cdot )\);

Chooses a vector \({\mathbf {x}} \xleftarrow {\$}{\mathbb {F}}_{q^m}^n\) s.t. \(\Vert {\mathbf {x}}\Vert = n\);

Chooses two invertible matrices \(Q \in {\mathbb {F}}_{q^m}^{k \times k}\) and \(P \in {\mathbb {F}}_q^{n \times n}\) uniformly at random;

computes \(\hat{G} = QG + Rot_k({\mathbf {x}})P\).
The public key is \(pk = (\hat{G}, {\mathbf {x}})\) and the private key is \(sk = (Q, G, P, \mathcal {DEC}(\cdot ))\).

Similar to in "Revisiting the LT scheme" section, we list the algorithms Enc and Dec of our proposed PKE scheme in Figs.Â 3 andÂ 4, respectively.
Security proof
In this section, we prove the correctness and INDCCA2 security of our codebased PKE scheme presented in the preceding section.
Correctness
According to the Enc and Dec algorithms of our scheme, we can easily get:
Then, we can compute the error vectors \({\mathbf {e}}_0\) and \({\mathbf {e}}_1\) correctly with the vector \({\mathbf {u}}_4\) and get \({\mathbf {u}}_3' = Conv^{1}_{t_0}({\mathbf {e}}_0)Conv^{1}_{t_1}({\mathbf {e}}_1) = {\mathbf {u}}_3\). From \({\mathbf {u}}_3\) and \({\mathbf {u}}_4\), we can recover \({\mathbf {u}}_1\) and \({\mathbf {u}}_2\) correctly and obtain the message \({\mathbf {m}}\) by computing
This concludes the correctness proof of our scheme.
INDCCA2 security
The INDCCA2 security for any PKE scheme can be defined as follows.
Definition 6
(INDCCA2 security) For any twostage (find stage and guess stage) adversary \({\mathcal {A}} =({\mathcal {A}}_1, {\mathcal {A}}_2)\) against a PKE scheme \({\mathcal {E}}=(\textsf {Setup}, \textsf {KGen}, \textsf {Enc}, \textsf {Dec})\), the INDCCA2 security is modeled as an experiment \(\mathbf{Exp }_{{\mathcal {A}}, \textsf {PKE}}^{\textsf {INDCCA2}}(\lambda )\) with respect to a security parameter \(\lambda\).

\(\mathbf{Exp }_{{\mathcal {A}}, \textsf {PKE}}^{\textsf {INDCCA2}}(\lambda )\):

\(params \leftarrow \textsf {Setup}(\lambda )\)

\((pk, sk) \leftarrow \textsf {KGen}(params)\)

\(({\mathbf {m}}_0, {\mathbf {m}}_1, \text{ state })\leftarrow {\mathcal {A}}_1^{\textsf {Dec}(sk,\cdot )}(pk)\), where \({\mathbf {m}}_0\) and \({\mathbf {m}}_1\) have the same length

\(b \xleftarrow {\$}\{0,1\}\)

\({\mathbf {c}}' \leftarrow \textsf {Enc}(pk,{\mathbf {m}}_b)\)

\(b' \leftarrow {\mathcal {A}}_2^{\textsf {Dec}(sk,\cdot )}({\mathbf {c}}',\text{ state })\). Note that \({\mathcal {A}}_2\) is not allowed to make decryption query on \({\mathbf {c}}'\).

Outputs 1 if \(b' = b\), otherwise outputs 0.
The advantage, \(\text{ Adv }_{{\mathcal {A}}, \textsf {PKE}}^{\textsf {INDCCA2}}(\lambda )\), of \({\mathcal {A}}\) in \(\mathbf{Exp }_{{\mathcal {A}}, \textsf {PKE}}^{\textsf {INDCCA2}}(\lambda )\) can be defined by
A PKE scheme \({\mathcal {E}}=(\textsf {Setup}, \textsf {KGen}, \textsf {Enc}, \textsf {Dec})\) is INDCCA2 secure if for any PPT \({\mathcal {A}}\), \(\text{ Adv }_{{\mathcal {A}}, \textsf {PKE}}^{\textsf {INDCCA2}}(\lambda )\) is negligible with respect to \(\lambda\). In addition, the above experiment can become an INDCPA experiment if \({\mathcal {A}}\) is not allowed to make decryption query. We denote by \(\mathbf{Exp }_{{\mathcal {A}}, \textsf {PKE}}^{\textsf {INDCPA}}(\lambda )\) the INDCPA experiment and the advantage of \({\mathcal {A}}\) is defined by
We divide the security proof of our scheme into two phases following the approach in [20]. First, we give the INDCPA security proof based on the INDCPA security of the LT scheme. Then, we prove the INDCCA2 security of our scheme in the random oracle model.
Theorem 1
Assume that \({\mathcal {A}}\) is an attacker in the experiment \(\mathbf{Exp }_{{\mathcal {A}}, {\textsf {PKE}}}^{{\textsf {INDCPA}}}(\lambda )\) with advantage \(\epsilon\) against the LT scheme. Then, we have
where \(q_{Rd}\) is the query times to oracle Rd and r is the length of the output vectors of the function Rd.
Proof
We assume that \({\mathcal {C}}\) is a challenger, who is responsible for the generation of the public parameters params and a publicprivate key pair (pk,Â sk) by running the algorithms \(\textsf {Setup}\) and \(\textsf {KGen}\), respectively. It is obvious that \({\mathcal {A}}\) should have the advantage \(\text{ Adv }_{{\mathcal {A}}, \textsf {PKE}}^{\textsf {INDCPA}}(\lambda )\) against our scheme if \({\mathcal {C}}\) simulates two oracles \({\mathcal {H}}\) and Rd correctly. After receiving \({\mathbf {m}}_0\) and \({\mathbf {m}}_1\) from \({\mathcal {A}}\), \({\mathcal {C}}\) randomly chooses a vector \(\mathbf {v}\), a bit b and computes \({\mathcal {H}}\) and Rd as
If \({\mathcal {A}}\) makes a query to Rd on a vector \({\mathbf {v}}'\), \({\mathcal {C}}\) first checks whether \({\mathbf {v}}' = {\mathbf {v}}\). If yes, \({\mathcal {C}}\) outputs \(Rd({\mathbf {v}}) = {\mathbf {u}}_1  ({\mathbf {m}}_b\mathbf {con})\). Otherwise, \({\mathcal {C}}\) outputs a random vector. If \({\mathcal {A}}\) makes queries to \({\mathcal {H}}\) on a vector \({\mathbf {u}}'_1\), \({\mathcal {C}}\) first checks whether \({\mathbf {u}}'_1 = {\mathbf {u}}_1\). If yes, \({\mathcal {C}}\) outputs \({\mathcal {H}}({\mathbf {u}}_1) = {\mathbf {u}}_2  {\mathbf {v}}.\) Otherwise, \({\mathcal {C}}\) outputs a random vector. It is obvious that \({\mathcal {C}}\) can not simulate the Rd correctly if \({\mathcal {C}}\) doesnâ€™t know \({\mathbf {u}}_1\). Hence, \({\mathcal {C}}\) could simulate Rd and \({\mathcal {H}}\) correctly if and only if \({\mathcal {A}}\) made queries to \({\mathcal {H}}\) on \({\mathbf {u}}_1\) first and then to Rd on \({\mathbf {v}}\). We define two events \(\textsf {Evt}_1\) and \(\textsf {Evt}_2\) as follows:

\(\textsf {Evt}_1\): the event that \({\mathbf {v}}\) is queried to Rd in \(q_{Rd}\) queries before \({\mathbf {u}}_1\) is queried to \({\mathcal {H}}\).

\(\textsf {Evt}_2\): the event that \({\mathbf {u}}_1\) is queried to \({\mathcal {H}}\) in \(q_{{\mathcal {H}}}\) times before \({\mathbf {v}}\) is queried to Rd.
It is obvious that the probability \(\Pr [\textsf {Evt}_1 \wedge \textsf {Evt}_2] = 0\) from the definition of \(\textsf {Evt}_1\) and \(\textsf {Evt}_2\) and we have
While the event \(\lnot \textsf {Evt}_1 \wedge \lnot \textsf {Evt}_2\) happens, \({\mathcal {A}}\) can not get any information about the connection between \({\mathbf {m}}_b\mathbf {con}\) and \({\mathbf {u}}_1{\mathbf {u}}_2\). Hence, the advantage of \({\mathcal {A}}\) is negligible in this case. For another case that \(\textsf {Evt}_1 \vee \textsf {Evt}_2\) happens, \({\mathcal {A}}\) can guess b correctly. Then we have the inequation \(2\Pr [\mathbf{Exp }_{{\mathcal {A}}, \textsf {PKE}}^{\textsf {INDCPA}}(\lambda )=1] \le 2\Pr [\textsf {Evt}_1\vee \textsf {Evt}_2] + 1  \Pr [\textsf {Evt}_1\vee \textsf {Evt}_2]\). That is, we have
According to Eq. (2), we have
and
i.e.,
As we know, b and \(\mathbf {v}\) are generated randomly by \({\mathcal {C}}\), so \({\mathcal {A}}\) could not obtain any information about them if the event \(\textsf {Evt}_1\vee \textsf {Evt}_2\) did not happen. For instead, \({\mathcal {C}}\) can recover \({\mathbf {u}}_4\) from \({\mathbf {c}}\) if the event \(\textsf {Evt}_2\) happens. In the other words, while event \(\textsf {Evt}_2 \wedge \lnot \textsf {Evt}_1\) happens, \({\mathcal {C}}\) could use \({\mathcal {A}}\) to break the LT scheme. Thus, we have
On the other hand, the probability that \(\mathbf {v}\) is exactly queried to Rd is \(\frac{1}{q^{mr}}\). Hence, the probability that \(\textsf {Evt}_1\) happens during the \(q_{Rd}\) queries on Rd is
According to Eqs. (3), (5), (6), (7) we can get:
That is, we have
\(\square\)
Theorem 2
Assume that \({\mathcal {A}}\) is an attacker in the experiment \(\mathbf{Exp }_{{\mathcal {A}}, {\textsf {PKE}}}^{{\textsf {INDCCA2}}}(\lambda )\) who has the advantage \(\epsilon\) against the LT scheme. Then, we have
where \(q_{Rd}\) and \(q_{{\mathcal {D}}}\) are the query times to the oracles Rd and \({\mathcal {D}}\), respectively, and r is the length of the output vectors of the function Rd.
Proof
We assume that \({\mathcal {C}}\) is a challenger, who is responsible for the generation of the public parameters params and a publicprivate key pair (pk,Â sk) by running the algorithms \(\textsf {Setup}\) and \(\textsf {KGen}\), respectively. It is clear that \({\mathcal {A}}\) should have the advantage \(\text{ Adv }_{{\mathcal {A}}, {\textsf {PKE}}}^{{\textsf {INDCCA2}}}(\lambda )\) against our scheme if \({\mathcal {C}}\) simulates oracles \({\mathcal {H}}\), Rd and the decryption oracle \({\mathcal {D}}\) correctly. \({\mathcal {C}}\) simulates \({\mathcal {H}}\) and Rd in the same way of the proof of TheoremÂ 1. Then, \({\mathcal {C}}\) uses the plaintextextractor described in [21] to construct the decryption oracle via the following steps.

The plaintextextractor takes as input a ciphertext \({\mathbf {c}}\).

For \(q_{Rd}\) times queries to Rd, the input and the output pairs of Rd are denoted by \(({\mathbf {v}}_i, V_i)\), \(1\le i \le q_{Rd}\). Similarly, the input and the output pairs of \({\mathcal {H}}\) are denoted by \(({\mathbf {u}}_{1j}, U_{1j})\), \(1\le j \le q_{{\mathcal {H}}}\).

The plaintextextractor finds the pair \(({\mathbf {v}}_i, V_i) = ({\mathbf {u}}_2  U_{1j}, {\mathbf {u}}_{1j}  ({\mathbf {m}}\mathbf {con}))\) s.t.
$$\begin{aligned} {\mathbf {e}}_0=\, & {} Conv_{t_0}(Lt_{\ell _0}(Lt_{\ell _0+\ell _1}({\mathbf {u}}_{1j}{\mathbf {u}}_2))), \end{aligned}$$(8)$$\begin{aligned} {\mathbf {e}}_1=\, & {} Conv_{t_1}(Rt_{\ell _1}(Lt_{\ell _0+\ell _1}({\mathbf {u}}_{1j}{\mathbf {u}}_2)) \end{aligned}$$(9)are two error vectors of \({\mathbf {c}}\) and \(Rt_{k}({\mathbf {u}}_{1j}{\mathbf {u}}_2)\) is the plaintext of \({\mathbf {c}}\) in the LT scheme.

If the plaintextextractor cannot find the pair \(({\mathbf {v}}_i, V_i)\) or \(({\mathbf {u}}_{1j}, U_{1j})\), it rejects the ciphertext \({\mathbf {c}}\).
Based on the above construction, we know that the plaintextextractor cannot simulate \({\mathcal {D}}\) unless \({\mathcal {A}}\) made queries \({\mathbf {u}}_{1j}\) to \({\mathcal {H}}\) and \({\mathbf {v}}_i\) to Rd before the ciphertext \({\mathbf {c}}\) is queried to \({\mathcal {D}}\). It should be noted that in this case we have Eqs. (8) and (9). For further analysis, we define a new event, denoted by \(\textsf {Evt}_3\), as

\(\textsf {Evt}_3\): the event that \({\mathcal {A}}\) queries the appropriate ciphertext \({\mathbf {c}}\) to \({\mathcal {D}}\) before \({\mathcal {A}}\) queries \({\mathbf {u}}_{1j}\) to \({\mathcal {H}}\) and \({\mathbf {v}}_i\) to Rd.
It is obvious that the probability that the appropriate ciphertext \({\mathbf {c}}\) is queried to \({\mathcal {D}}\) is \(\frac{1}{q^{2mn}}\). Hence, we get the probability that the event \(\textsf {Evt}_3\) occurs in at most \(q_{{\mathcal {D}}}\) queries satisfies
\(\square\)
As in the proof of TheoremÂ 1, \({\mathcal {C}}\) could simulate these oracles correctly while the event \(\textsf {Evt}_2\) (equivalent to event \(\textsf {Evt}_2 \wedge \lnot \textsf {Evt}_3 \wedge \lnot \textsf {Evt}_1\)) happened. Then, we have
That is, we have \(\text{ Adv }_{{\mathcal {A}}, \textsf {PKE}}^{\textsf {INDCCA2}}(\lambda ) \le \varepsilon + \frac{q_{Rd}}{q^{mr}}+ \frac{q_{{\mathcal {D}}}}{q^{2mn}}.\)
Efficiency
In this section, we analyze the efficiency of our proposed PKE scheme by comparing with two recently proposed schemes (i.e., Wangâ€™s scheme [22] and Loidreauâ€™s scheme [17]), in terms of public key and ciphertext sizes. It should be noted that both Loidreauâ€™s scheme and our proposed scheme are based on rank metric codes, while Wangâ€™s scheme is based on Hamming metric codes. For a better understanding of our strengths, we introduce a family of the most representative rank metric codes called Gabidulin codes first.
Definition 7
(Gabidulin codes) For an element \(a\in {\mathbb {F}}_{q^m}\) and an integer \(i\in {\mathbb {Z}}\), we denote \(a^{[i]}=a^{q^{i}}\). Further for a vector \(\mathbf {a} = (a_0, a_1, \ldots , a_{n1})\in {\mathbb {F}}_{q^m}^n\), we denote \(\mathbf {a}^{[i]}= (a_0^{[i]}, a_1^{[i]}, \ldots , a_{n1}^{[i]})\). Then, an \([n,k]_{{\mathbb {F}}_{q^m}}\) Gabidulin code \({\mathcal {C}}({\mathbf {a}})\) for a vector \(\mathbf {a} = (a_0, a_1, \ldots , a_{n1}) \in {\mathbb {F}}_{q^m}^n\) is defined with the following k by n generator matrix:
According to the analysis in [18, 23,24,25], a secure codebased PKE scheme should be able to resist several structural attacks, such as key recovery attack, reduction attack and moore decomposition attack [23]. In particular, we can obtain the lower bound of the computation costs that an adversary needs to break a PKE scheme based on Gabidulin codes. The lower bound, known as the PQsecurity level [17, 26], can be computed by the following formula:
where t is the correction ability of the underlying Gabidulin codes.
As the same in the LT scheme [18], we choose two sets of parameters \(params_1=(q = 2, m = 75, n = 73, k = 21, t_0 = t_1 = 13)\) and \(params_2=(q = 2, m = 85, n = 83, k = 18, t_0 = t_1 = 16)\), which can achieve the security levels \(2^{141} \approx 75^3 \cdot 2^{\cdot (131) \lfloor \frac{21\cdot \min \{75,73\}}{146} \rfloor }\) and \(2^{154} \approx 85^3 \cdot 2^{\cdot (161) \lfloor \frac{18\cdot \min \{85,83\}}{166} \rfloor }\), respectively, according to Eq. (11). Recall that the ciphertext in our scheme is of the form \({\mathbf {c}} = ({\mathbf {c}}_0, {\mathbf {c}}_1)\), where \({\mathbf {c}}_0\) and \({\mathbf {c}}_1\) are two vectors of the same length n over \({\mathbb {F}}_{q^m}\). For the first parameter set \(params_1\), we have the ciphertext size of \(2\cdot 73\cdot 75\ \text{ bits }=10950 \ \text{ bits }\approx 1.37\ \text{ KB }\). Similarly, we can get the ciphertext size of \(2\cdot 83\cdot 85\ \text{ bits }=14110\ \text{ bits } \approx 1.76\ \text{ KB }\) for the second parameter set \(params_2\).
As to the public key, it has the form of \(pk = (\hat{G}, {\mathbf {x}})\) in our scheme, where \(\hat{G}\) is a \(k\times n\) matrix over \({\mathbb {F}}_{q^m}\) and \({\mathbf {x}}\) is a vector of the length n over \({\mathbb {F}}_{q^m}\). For the first parameter set \(params_1\), we have the public key size of \(75\cdot (21+1)\cdot 73\ \text{ bits }=120450 \ \text{ bits }\approx 15.06\ \text{ KB }\). Similarly, we can get the public key size of \(85\cdot (18+1)\cdot 83\ \text{ bits }=134045\ \text{ bits } \approx 16.76\ \text{ KB }\) for the second parameter set \(param_2\).
We note that the ciphertext and public key sizes of our INDCCA2 secure scheme are exactly the same to the ones of the INDCPA secure LT scheme. Compared to Loidreauâ€™s scheme [17], which is a recently proposed PKE scheme based on rank metric codes, our PKE scheme has a significant advantage in terms of public key size and ciphertext size (see TableÂ 2).
For a fair comparison between our rank metric codesbased PKE scheme and Wangâ€™s scheme from Hamming metric codes [22], we choose two new sets of parameters \(params_3 = (q = 2, m = 71, n = 69, k = 19, t_0 = t_1 = 13)\) and \(params_4 = (q = 2, m = 101, n = 99, k = 22, t_0 = t_1 = 17)\), which can achieve the security levels \(71^3 \cdot 2^{\cdot (131) \lfloor \frac{19\cdot \min \{71,69\}}{138} \rfloor }\approx 2^{133}\) and \(101^3 \cdot 2^{\cdot (171) \lfloor \frac{22\cdot \min \{101,99\}}{198} \rfloor } \approx 2^{196}\), respectively, according to Eq. (11). For the parameter set \(params_3\), we have the ciphertext size of \(2\cdot 69\cdot 71\ \text{ bits }=9798 \ \text{ bits }\approx 1.22\ \text{ KB }\) and public key size of \(71\cdot (19+1)\cdot 69\ \text{ bits }=97890\ \text{ bits } \approx 12.25\ \text{ KB }\). Similarly, we can get the ciphertext size of \(2\cdot 99\cdot 101\ \text{ bits }=19998\ \text{ bits } \approx 2.50\ \text{ KB }\) and public key size of \(101\cdot (22+1)\cdot 99\ \text{ bits }=229977\ \text{ bits } \approx 28.75\ \text{ KB }\) for the parameter set \(params_4\). From TableÂ 2, we observe that our codebased PKE scheme also has a significant advantage to Wangâ€™s scheme in terms of public key size.
A secure cloud storage use case
In this section, we explain how our proposed codebased PKE scheme can be deployed in a cloud storage system. We use the notations \(\textsf {CBEnc}(\cdot )/\textsf {CBDec}(\cdot )\) and \(\textsf {SEnc}(\cdot )/\textsf {SDec}(\cdot )\) to denote the encryption/decryption algorithms in our proposed codebased PKE scheme and a secure symmetric encryption scheme, respectively. In our cloud storage system, we assume that there are two users (Alice and Bob) and multiple storage servers. Alice splits her data into some data blocks and sends each block to storage servers (one data block can be sent to several storage severs). Bob can get the data blocks from the cloud and recover them into the original data (see Fig.Â 5).
If Alice wants to store data M in the cloud and shares it with Bob, she can select a random session key K for the symmetric encryption algorithm and encrypts it under Bobâ€™s public key \(pk_{B}\) to get a ciphertext \(EK=\textsf {CBEnc}(pk_{B}, K)\). Then, Alice can split the data M into m blocks \(M=M_1\Vert M_2\Vert \cdots \Vert M_m\) and uses K to encrypt them to get \(C_i = \textsf {SEnc}(K, M_i)\), \(1\le i \le m\). Next, Alice computes the hash values \(h_i=Hash(C_i)\), \(1\le i\le m\), and saves them in her local storage, where \(Hash(\cdot )\) is a secure collisionresistant hash function. Then, Alice sets the information \(ID_{A}\), EK, and \(EH_i = \textsf {SEnc}(K, h_i)\) to each ciphertext block \(C_i\), \(1\le i\le m\), as the header data (see Fig.Â 6). Finally, she sends the header data and ciphertext \(C_i\), \(1\le i \le m\), to m storage servers. If Alice wants to check the data integrity of some ciphertext block \(C_i\), she can ask the corresponding storage server to compute \(Hash(C_i)\) and send it back to her. If the received \(Hash(C_i)\) matches her local hash value for \(C_i\), she is assured that the data are intact and have not modified by the cloud storage server.
On the other hand, if Bob wants to get the data M from Alice via the cloud storage system, he can find the corresponding m ciphertext data blocks \(C_i\), \(1\le i\le m\), and their header data \((ID_A, EK, EH_i)\) by searching for \(ID_{A}\) in the cloud. Then, Bob can download them and runs \(\textsf {CBDec}(sk_B, EK)\) to recover the symmetric key K with his private key \(sk_B\). Next, for each \(1\le i\le m\), Bob uses the obtained symmetric key K to decrypt \(EH_i\) and \(C_i\) to get \(Hash(C_i)\) and \(M_i\), respectively. To verify the data integrity of \(M_i\), Bob computes \(Hash(C_i)\) and checks whether it is equal to \(h_i\) obtained from \(EH_i\). If yes, it means that \(C_i\) has not been modified by the storage server and the corresponding \(M_i\) is the original data from Alice. Otherwise, Bob asks the corresponding \(SS_i\) of the nonmatching \(C_i\) to send the data again.
From the above description, we find that the security of the proposed cloud storage system relies mainly on the security of our codebased PKE scheme and the symmetric encryption scheme. As mentioned in "Introduction" section, our scheme is able to resist quantum computerfacilitated attacks and achieves INDCCA2 security as demonstrated in "Security proof" section. Clearly, Alice can access her data stored in the cloud using any devices by searching for the data with her ID. Moreover, both Alice and Bob can check the integrity of the downloaded data from the cloud. In the event that the data were modified by the cloud storage server or during the communication, Bob can detect this and then proceed to download the corresponding blocks again. Thus, the proposed cloud storage system provides availability, reliability, efficient retrieval and data sharing, even in the postquantum era.
Conclusion
Codebased public key cryptography (PKC) can potentially be deployed in realworld applications to ensure quantum computing attack resilience. Existing codebased PKC schemes are broadly categorized into those based on Hamming metric codes and those based on rank metric codes, and it is believed that the latter generally has smaller public key and ciphertext sizes than the former under the same security level.
In this paper, we presented a new rank metric codesbased public key encryption scheme from Lau and Tanâ€™s scheme [18], and hence inherits the latterâ€™s small public key and ciphertext size properties. However, our new scheme achieves INDCCA2 secure (the highest security assurance possible), as shown in this paper. We also analyzed its efficiency in term of the public key size and the ciphertext size. Finally, we presented a use case to demonstrate its utility in practice.
Future work includes constructing more rank metric codesbased cryptographic primitives, such as proxy reencryption and attributebased encryption, to achieve other desirable properties in practical applications. We also intend to implement and evaluate a prototype of the proposed/extended scheme in collaboration with a (small) cloud storage service provider.
Availability of data and materials
The datasets used and analysed in this study are available from the corresponding author on reasonable request.
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We would like to thank the reviewers for their valuable comments.
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The work is supported in part by the National Key R&D Program of China under Grant No. 2017YFB0802300, the NSFCZhejiang Joint Fund for the Integration of Industrialization and Informatization under Grant No. U1509219, the Shanghai Natural Science Foundation under Grant No. 17ZR1408400, the National Natural Science Foundation of China under Grant Nos. 61601129, 11701179, the Shanghai Science and Technology Commission Program under Grant No. 18511105700, and the Shanghai Sailing Program under Grant No. 17YF1404300.
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Zeng, P., Chen, S. & Choo, KK.R. An INDCCA2 secure postquantum encryption scheme and a secure cloud storage use case. Hum. Cent. Comput. Inf. Sci. 9, 32 (2019). https://doi.org/10.1186/s1367301901936
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DOI: https://doi.org/10.1186/s1367301901936