Efficient and Secure Many-to-One Signature Delegation

Abstract

We propose an IBPMS scheme from pairings, which is more efficient in the sense of computation and operation time than the existing schemes. We also prove on random oracle that the propose d scheme is secure against existential forgery on adaptive chosen-message and adaptive-chosen ID attack under the k-CAA assumption. Additionally, our scheme fulfills all the security requirements of a proxy signature scheme. Moreover we do an efficiency analysis and show that our scheme is significantly more efficient than the existing IBPMS schemes in the sense of computation and operation time.

1 Introduction

In a proxy signature scheme, an original signer $o$ can transfer its signing rights to a proxy signer $\rho$ without transferring its private key; and the proxy signer can sign any document on behalf of the original signer. Proxy signature schemes are applicable in distributed systems, grid computing, mobile agent environment etc. where delegation of rights is quite common.

The notion of proxy signature has been around since 1989 due to Gasser et al. [3] but the first formal construction of a proxy signature scheme [6] was proposed in 1996. The notion of proxy multi-signature was introduced by Yi et al. [10] in 2000 and then in 2005, Li and Chen [5] proposed the first proxy multi-signature scheme in ID-based setting using bilinear pairings. Since then, many identity (ID)-based proxy multi-signature (IBPMS) schemes have been proposed using bilinear pairings, but most of the schemes are either too much inefficient or insecure, hence cannot be considered for the practical implementation.

We propose here an IBPMS scheme from bilinear pairings. Our scheme is significantly more efficient than the existing IBPMS schemes [1, 5, 8, 9] in the sense of computation and operation time. Moreover, we prove the security of our scheme against existential forgery on adaptive chosen-message and adaptive chosen-ID attacks in random oracle model. Additionally, we also show that the proposed scheme fulfills all the security requirements of a proxy signature scheme listed in [4].

2 Preliminaries

Definition 1

Let $G_{1}$ be an additive cyclic group with generator P and $G_{2}$ be a multiplicative cyclic group with generator g. Let both the groups are of the same prime order q. A map $e: G_1\times G_1\rightarrow G_2$ is called a cryptographic bilinear map or a pairing if it satisfies the following properties:

• 1. Bilinearity: For all $a, b \in \mathbb {Z}_q^*$, $e(aP, bP) = e(P,P)^{ab}$, or equivalently, for all $Q,R,S \in G_{1}$, $e(Q+R,S)=e(Q,S)e(R,S)$ and $e(Q,R+S)=e(Q,R)e(Q,S)$.

• 2. Non-degeneracy: There exists $Q,R \in G_{1}$ such that $e(Q,R)\ne 1$. Note that since $G_{1}$ and $G_{2}$ are groups of prime order, this condition is equivalent to the condition $e(P,P) \ne 1$, which again is equivalent to the condition that e(P, P) is a generator of $G_{2}$.

• 3. Computability: There exists an efficient algorithm to compute $e(Q,R)\in G_2$, for any $Q,R \in G_{1}$.

Definition 2

The k -CAA Problem [7] is to compute $\frac{1}{s+e_0}P$, for some $e_0\in Z_{q}^*$ when given $P, sP\in G_1, e_1,e_2,\dots ,e_k\in Z_{q}^*$ and $\frac{1}{s+e_1}P, \frac{1}{s+e_2}P,\dots ,\frac{1}{s+e_k}P\in G_1$.

Definition 3

The $(t, \epsilon )~k$ -CAA assumption holds in $G_1$ if there is no algorithm which takes at most t running time and can solve the k-CAA problem with at least a non-negligible advantage $\epsilon$.

3 Proposed IBPMS Scheme

3.1 Setup

Given a security parameter $1^K$, the private key generator (PKG) generates the system’s master secret $s\in \mathbb {Z}_q^*$ and the system’s public parameters

$$ params = (K,q,G_{1},G_{2},e,H_1,H_2,P,Q), $$

where $G_{1}$ is an additive cyclic group of prime order q; $G_{2}$ is a multiplicative cyclic group of prime order q; $e:G_1\times G_1\rightarrow G_2$ is a bilinear map defined as above; $H_1:\{0,1\}^*\rightarrow \mathbb {Z}_q^*$ and $H_{2}:\{0,1\}^*\times G_{1} \rightarrow \mathbb {Z}_q^*$ are two cryptographic hash functions; P is a generator of $G_{1}$; and $Q:=sP\in G_1$ is system’s public key.

3.2 Extraction

Given a user’s identity $\text {ID}$, the PKG computes its

• public key as: $Q_{\text {ID}}=H_1(\text {ID})$ and

• private key as: $S_{\text {ID}}=\frac{1}{s+H_{1}(\text {ID})}P$.

Thus, the proxy signer $A_{0}$ and the original signers, $A_{i}$, $i=1,\dots ,n$, have their public keys and the corresponding private keys

• $Q_{\text {ID}_{A_i}}=H_{1}(\text {ID}_{A_i})$ and

• $S_{\text {ID}_{A_i}}=\frac{1}{s+H_{1}(\text {ID}_{A_i})}P$,

for $i=0,1,\dots ,n$.

3.3 Proxy Key Generation

In this phase, all the original signers interact with the proxy signer to delegate their signing rights through a signed warrant. The warrant w includes some specific information about the message like nature of the message, time of delegation, identity information of the original signers and the proxy signer, period of validity, some public keys etc. After the successful interaction, the proxy signer outputs its proxy signing key. The interaction and proxy key generation can be described in the following phases:

• Delegation generation:

  Each of the n original signers, $A_{i}$, $i=1,\dots ,n$, and the proxy signer, $A_{0}$, interact and do the following:

  • – set $I = \sum _{i=0}^{n} H_{1}(\text {ID}_{A_i}) \in \mathbb {Z}_q^*$ and $J=(n+1)Q+IP \in G_1$;

  • – select ;

  • – compute and publish $V_{i}=e(J,S_{\text {ID}_{A_{i}}})\in G_2$, and $W_{i}=V_{i}^{x_{i}} \in G_{2}$;

  • – create a warrant w which includes the identities $\text {ID}_{A_{i}}$ of the proxy and original signers, the values $V_{i}$ and $V_{o}=\prod _{i=0}^{n}V_{i}$, the values $W_{i}$ and $W_{o}=\prod _{i=0}^{n}W_{i}$, the scope of messages to be signed, time of delegation, period of validity etc.;

  • – computes $h_{o}=H_2(w,W_{o})\in \mathbb {Z}_q^*$;

  • – and $S_i=(x_i+h_{o})S_{\text {ID}_{A_i}}$.

  • – Finally each original signer sends $(w,S_i)$, $i=1,\dots ,n$, to the proxy signer as a partial delegation.

• Delegation verification:

  On receiving $(w,S_i)$ from each original signer $A_i$, the proxy signer $A_0$ obtains $W_{i}$ and $W_{o}$ from the warrant, computes $h_{o}=H_2(w,W_{o}) \in \mathbb {Z}_q^*$ and validates each partial delegation by checking

  $$\begin{aligned} e(J,S_i)= W_{i} V_{i}^{h_{o}}. \end{aligned}$$

  If the above equality does not hold for any $1\le i\le n$, the proxy signer terminates the protocol.

• Proxy key generation:

  In this phase, the proxy signer computes its proxy secret key to sign the message on behalf of the group of original signers to be

  $$\begin{aligned} S_{pk}=\sum \nolimits _{i=0}^{n} S_i. \end{aligned}$$

3.4 Proxy Multi-signature

To sign a message $m\in \{0,1\}^*$ under the warrant w on behalf of the group of original signers, the proxy signer does the following:

• selects ;

• computes $V_{\rho }=V_{o}^y$ and $W_{\rho }=W_{o}^y$;

• computes $h_{\rho }=H_2(m,W_{\rho })\in Z_q^*$

• and $\sigma =(y+h_{\rho })S_{pk}$.

• Finally, $(\sigma ,V_{\rho },W_{\rho })$ is the IBPMS by the proxy signer on behalf of the group of original signers on message m under the warrant w.

3.5 Verification:

On receiving the IBPMS $(\sigma ,V_{\rho },W_{\rho })$ on message m under the warrant w, a verifier validates it as follows:

• checks if the message m confirms to the warrant w. Stops if not. Continues otherwise.

• checks whether the proxy signer $A_0$ is authorized or not in the warrant w, by the group of n original signers. Stops if not. Continues otherwise.

• obtains $V_{o}$ and $W_{o}$ from the warrant w, computes $h_{o}=H_2(w,W_{o}) \in Z_q^*$ and $h_{\rho }=H_2(m,W_{\rho }) \in Z_q^*$ and accepts $(\sigma ,V_{\rho },W_{\rho })$ as a valid IBPMS on message m, if and only if the following equality holds:

  $$ e(J,\sigma ) = W_{\rho } V_{\rho }^{h_{o}} W_{o}^{h_{\rho }} V_{o}^{h_{o}h_{\rho }}. $$

4 Analysis of the Proposed Scheme

In this section, we first give the correctness of our scheme then analyze the security of our scheme and show that the proposed scheme satisfies all the security requirements of a proxy signature scheme [4].

4.1 Correctness

Correctness of the delegation verification holds since for all $0\le i\le n$,

$$\begin{aligned} e(J,S_i) = e(J,(x_i+h_{o})S_{\text {ID}_{A_i}}) = e(J,S_{\text {ID}_{A_i}})^{x_i+h_{o}} = V_{i}^{x_i+h_{o}} = W_{i} V_{i}^{h_{o}}. \end{aligned}$$

(1)

Correctness of the IBPMS verification holds since

$$\begin{aligned} e(J,\sigma )&= e(J,(y+h_{\rho })S_{pk}) = e(J,(y+h_{\rho })\sum _{i=0}^{n} S_i) \\&= \prod _{i=0}^{n} e(J,(y+h_{\rho }) S_i) = \prod _{i=0}^{n} e(J, S_i)^{y+h_{\rho }} = \prod _{i=0}^{n} \left( W_{i}V_{i}^{h_{o}} \right) ^{y+h_{\rho }} \qquad \text {from (1)}\\&= \left( \prod _{i=0}^{n} W_{i}V_{i}^{h_{o}} \right) ^{y+h_{\rho }} = \left( \prod _{i=0}^{n} W_{i} \prod _{i=0}^{n} V_{i}^{h_{o}} \right) ^{y+h_{\rho }} \\&= \left( W_{o} V_{o}^{h_{o}} \right) ^{y+h_{\rho }} = W_{o}^{y} (V_{o}^{y})^{h_{o}} W_{o}^{h_{\rho }} (V_{o}^{h_{o}})^{h_{\rho }} = W_{\rho } V_{\rho }^{h_{o}} W_{o}^{h_{\rho }} V_{o}^{h_{o}h_{\rho }}. \end{aligned}$$

4.2 Security Analysis

Theorem 1

The proposed IBPMS scheme is strongly unforgeable if the k-CAA is intractable in $G_1$.

Proof. For security parameter $1^k$, the challenger $\mathcal {C}$ runs the setup algorithm and provides $\langle q, G_1, P, sP, (e_1,f_1),\dots ,(e_k,f_k) \rangle$ to $\mathcal {B}$ where $G_{1}$ is an additive cyclic group of prime order q; P is a generator of $G_{1}$ and $s,e_1,\dots ,e_k\in \mathbb {Z}_q^*$ are randomly chosen elements and $f_i:=\frac{1}{s+e_i}P\in G_1$, $i=1,\dots ,k$. The goal of $\mathcal {B}$ is to solve the k-CAA problem by producing a pair $\left( e_0,\frac{1}{s+e_0}P\right)$ for some $e_0 \in \mathbb {Z}_q^*$, $e_0 \ne e_i$ for all $i=1,\dots ,n$.

Let $\mathcal {A}$ be a forger algorithm who claims is to break the proposed identity based proxy multi-signature scheme. The adversary $\mathcal {B}$ simulates the challenger and interacts with $\mathcal {A}$. We facilitate the adversary $\mathcal {A}$ to adaptively select the identity $\text {ID}^{*}$ on which it wants to forge the signature. Further the adversary can obtain the private keys associated to the identities. The adversary also can access the proxy multi-generation oracles on warrants $w'$ of its choice, and proxy multi-signature oracles on the warrant, messages pair $(w', m')$ of its choice upto polynomial many times.

Setup: For security parameter $1^k$, $\mathcal {B}$ generates the system’s public parameter $params = \langle q, G_1, G_2, e, H_1, H_2, P, Q=sP \rangle$ where $G_{2}$ is a multiplicative cyclic group of prime order q; $e:G_1\times G_1\rightarrow G_2$ is a bilinear map defined as in Sect. 2; and $H_1:\{0,1\}^*\rightarrow \mathbb {Z}_{q}^*$ and $H_{2}:\{0,1\}^*\times G_{1} \rightarrow \mathbb {Z}_{q}^*$ are two cryptographic hash functions and provides $params$ to $\mathcal {A}$. $\mathcal {B}$ picks a random index $i^*\in [1,k+1]$ and a random $e_0\in \mathbb {Z}_q^*$. It then resets the values $(e_i,f_i)=(e_i,f_i)$ for $1\le i<i^*$, $(e_{i^*},f_{i^*})=(e_0,\perp ))$ and $(e_i,f_i)=(e_{i-1},f_{i-1})$ for $i^*<i\le k+1$.

$H_1$ -queries: To respond to the $H_1$ hash function queries, $\mathcal {B}$ maintains a list $L_{H_1}=\{\langle \text {ID}, e, f \rangle \}$. When $\mathcal {A}$ requests the $H_1$ query on some identity $\text {ID}_i\in \{0,1\}^*$, $i\le k+1$, $\mathcal {B}$ responds as follows:

• 1. If the query $\text {ID}_i$ already appears in the list $L_{H_1}$ in some tuple $\langle \text {ID}_j, e_j, f_j \rangle$, $j<i$, then algorithm $\mathcal {B}$ responds to $\mathcal {A}$ with $H_1(\text {ID}_i)=e_j$. So WLOG we assume $\text {ID}_i \ne \text {ID}_j$ for $i \ne j$.

• 2. Otherwise $\mathcal {B}$ responds to $\mathcal {A}$ with $H_1(\text {ID}_i)=e_i$ and adds the tuple $\langle \text {ID}_i, e_i, f_i \rangle$ to the list $L_{H_1}$.

$H_2$ -queries: To respond to the $H_2$ hash function queries, $\mathcal {B}$ maintains a list $L_{H_2}=\{\langle w, U, g \rangle \}$. When $\mathcal {A}$ requests the $H_2$ query on $(w', U')$ for some $w'\in \{0,1\}^{*}$ and $U'\in G_1$, $\mathcal {B}$ responds as follows:

• 1. If the query $(w', U')$ already appears on the list $L_{H_2}$ in some tuple $\langle w', U', g \rangle$ then algorithm $\mathcal {B}$ responds to $\mathcal {A}$ with $H_2(w'\Vert U')=g$.

• 2. Otherwise $\mathcal {B}$ picks a random integer $g \in \mathbb {Z}_q^*$ and adds the tuple $\langle w', U', g \rangle$ to the list $L_{H_2}$ and responds to $\mathcal {A}$ with $H_2(w'\Vert U')=g$.

Extraction Queries: When $\mathcal {A}$ makes a private key query on some identity $\text {ID}_i$, $i\le k+q$, $\mathcal {B}$ responds as follows:

• 1. If $i=i^*$, then $\mathcal {B}$ reports failure and terminates. The probability of such failure is $1/(k+1)$.

• 2. Otherwise $\mathcal {B}$ responds to $\mathcal {A}$ with $S_{\text {ID}_i}=f_i$ and adds the tuple $\langle \text {ID}_i, e_i, f_i \rangle$ to the list $L_{H_1}$.

Recall that, for $i \ne i^*$, $H(\text {ID}_{i})=e_i$ and $f_i=\frac{1}{s+e_i}P$. So, $S_{\text {ID}_i}=\frac{1}{s+H(\text {ID}_i)}P$ is a valid private key of the user with identity $\text {ID}_i$.

Delegation Queries: To respond to the delegation queries, $\mathcal {B}$ maintains a list $L_{del}=\{\langle w,(x_0,S_0),(x_1,S_1),\dots ,(x_n,S_n) \rangle \}$ and responds to identical queries in a consistent fashion. It uses $L_{H_1}$ and $L_{H_2}$ to generate the needed hash values and the secret keys and computes the delegations $\langle w, S_{1}, \dots , S_{n} \rangle$ as in the actual scheme. $\mathcal {B}$ may have to terminate if the identity of one of the original signers is $\text {ID}_{i^*}$ and the probability for that event is bounded by $(n+1)/(k+1)$.

Proxy Key Generation Queries: To respond to the proxy key generation queries, $\mathcal {B}$ maintains a list $L_{pkg}=\{\langle w, S \rangle \}$ and responds to identical queries in a consistent fashion. It uses $L_{H_1}$ and $L_{H_2}$ to generate the needed hash values and the secret keys and computes the proxy key $\langle w, S \rangle$ using $L_{del}$ as in the actual scheme. $\mathcal {B}$ may have to terminate if the identity of one of the original signers or the proxy signer is $\text {ID}_{i^*}$ and the probability for that event is $(n+1)/(k+1)$.

Proxy Multi-Signature Queries: To respond to the proxy multi-signature queries, $\mathcal {B}$ maintains a list $L_{pms}=\{\langle w, m, y, V, W, \sigma \rangle \}$ and responds to identical queries in a consistent fashion. It uses $L_{H_1}$ and $L_{H_2}$ to generate the needed hash values and the secret keys and computes the delegations $\langle w, m, V, W, \sigma \rangle$ using $L_{del}$ and $L_{pkg}$ as in the actual scheme.

$\mathcal {B}$ may have to terminate if the identity of one of the original signers or the proxy signer is $\text {ID}_{i^*}$ and the probability for that event is $(n+1)/(k+1)$.

Output: $\mathcal {A}$ outputs a valid ID-based proxy multi-signature $(\sigma , V_{\rho }, W_{\rho })$ on a message m under the warrant w by the proxy signer $A_0$ on behalf of the group of original signers $A_1,\dots ,A_n$ such that

$$\begin{aligned} e(J,\sigma ) = W_{\rho } V_{\rho }^{h_{o}} W_{o}^{h_{\rho }} V_{o}^{h_{o}h_{\rho }} \end{aligned}$$

(2)

where $J,V_{o},W_{o},h_{o},h_{\rho }$ are defined as in Sect. 3.

If $\mathcal {A}$ does not query any hash function, that is, if responses to any of the hash function query is picked randomly then the probability that verification equality holds is less than 1/q. Thus, with probability greater than $1-1/q$, all the public keys and were computed using $H_1$-oracle.

For the forgery to be valid, $\mathcal {A}$ must not have queried the private key of at least one of the signers, say $A_{i}$, and must not have received $(\sigma , V_{\rho }, W_{\rho })$ as a response to a proxy key generation query. The probability that the identity of $A_i$ is $\text {ID}_{i^*}$ is $1/(k+1)$ and in that case, $H_1(ID_{A_i})=e$.

Then, using the Eq. (2) and the values returned by the adversary we can reverse compute the secret key $S_{\text {ID}_{A_i}}$ of $A_i$ as in [2]. But by definition, $S_{\text {ID}_{A_i}} = \frac{1}{s+H_1(ID_{A_i})}P = \frac{1}{s+e}P$. Thus $\mathcal {B}$ can then return the pair $(e,S_{\text {ID}_{A_i}})$ to the challenger $\mathcal {C}$ and win the k-CAA game.

Hence the proposed IBPMS scheme is secure.

5 Efficiency Analysis

We compare the total number of bilinear pairings (P), map-to-point hash functions (H), modular exponentiations (E) and pairing-based scalar multiplications (PSM) in Proxy key generation phase, Proxy multi-signature phase and the Verification phase with those of other IBPMS schemes [1, 5, 8, 9] and show that our scheme is computationally more efficient and takes less operation time than the known best IBPMS schemes given in [1, 5, 8, 9] (Table 1).

Table 1. Efficiency comparison