Security Model. We assume that the reader is familier with the model of Xxxxxxx et al. [14], which is the model in which we prove security of our dynamic key aggreement protocol. For completeness, we review their definitions and refer the reader to [14] for more details. Let P = {U1, . . . , Un} be a set of n (fixed) users or participants. A user can execute the protocol for group key agreement several times with different partners, can join or leave the group at it’s desire by executing the protocols for Insert or Delete. We assume that users do not deviate from the protocol and adversary never participates as a user in the protocol. This adversarial model allows concurrent execution of the protocol. The interaction between the adversary A and the protocol participants occur only via oracle queries, which model the adversary’s capabilities in a real attack. These queries are as follows, where Π Πi . denotes the i-th instance of user U and ski denotes the session key after execution of the protocol by – Send(U, i, m) : This query models an active attack, in which the adversary may intercept a message and then either modify it, create a new one or simply forward it to the intended participant. The output of the query is the reply (if any) generated by the instance Πi upon receipt of message m. The adversary is allowed to prompt the unused instance Πi to initiate the protocol with partners U2, . . . , Ul, l ≤ n, by invoking Send(U, i, ⟨U2, . . . , Ul⟩). – Execute({(V1, i1), . . . , (Vl, il)}) : Here {V1, . . . , Vl} is a non empty subset of P. This query models passive attacks in which the attacker evesdrops on honest execution of group key agreement protocol among unused instances Πi1 , . . . , Πil and outputs the transcript of the execution. A transcript consists of V1 Vl the messages that were exchanged during the honest execution of the protocol. – Join({(V1, i1), . . . , (Vl, il)}, (U, i)) : This query models the insertion of a user instance Πi in the group (V1, . . . , Vl) ∈ P for which Execute have already been queried. The output of this query is the transcript generated by the invocation of algorithm Insert. If Execute({(V1, i1), . . . (Vl, il)}) has not taken place, then the adversary is given no output. – Leave({(V1, i1), . . . , (Vl, il)}, (U, i)) : This query models the removal of a user instance Πi from the group (V1, . . . Vl) ∈ P. If Execute({(V1, i1), . . . (Vl, il)}) has not taken place, then the adversary is given no output. Otherwise, algorithm Delete is...
Security Model. This section defines the components of the system, the adversary and its capabilities and the meaning of system breakdown.
4.1.1. System The system comprises nodes belonging to one administrative unit under the same TA. It is assumed that TA has access to a cryptographically secure random number generator. The master keys are assumed secure and cannot be stolen. If need be, they can be deleted after generating all of the possible public and private key sets. The nodes have access to secure cryptographic algorithms, such as AESencryption and hash algorithms.
Security Model. We prove our protocols secure in the Universal Composability framework intro- duced in [Can01]. This model is explained in Appendix A.
Security Model. The model is defined by the following game which is run between a challenger C H and an adversary A . A controls all communications from and to the protocol participants via accessing to a set of oracles as described below. Every participant involved in a session is treated as an oracle. We denote an instance i of the participant U as k = sr (R + PK − X ) = sr (r + s − x )P = ∏i , where U ∈ {C , · · · ,C } S S. Each client C has an 3 S C C C S C C C U 1 n (rC + sC − xC)rSsP = (rC + sC − xC)RS = k4. Thus the client C and the server S establish a common session key sk = H4(IDC, RS, RC,WC, Ppub, k3) = H4(IDC, RS, RC,WC, Ppub, k4).
Security Model.
3.1. We settle the basic notation of distinguishers in Sect.
3.2. For reference, the black-box duplex security model of Daemen et al. [15] is treated in Sect.
3.3. We lift the model to leakage resilience in Sect. 3.4.
3.1 Sampling of Keys D ←−− { } The duplex construction of Sect. 2 is based on an array of u k-bit keys. These keys may be generated uniformly at random, as K DK ( 0, 1 k)u. In our analysis of leakage resilience, however, we will require the scheme to be still secure if the keys are not uniformly random but as long as they have sufficient min-entropy. Xxxxxxxxxx, we will adopt the approach of Daemen et al. [15] to consider keys sampled using a distribution K , that distributes the key independently1 and with sufficient min-entropy, i.e., for which D∞ δ H ( K ) = min ∈[1,u] H∞(K[δ]) is sufficiently high. Note that if DK is the random distribution, H∞(DK ) = k.
Security Model. C A Π Ω The security model of the privacy-preserving authentication protocol for mobile Internet environment is defined by a game between a challenger and an adversary modeled by a probabilistic polynomial-time (P.P.T) Turing machine [37]. Let l denote an instance l of a participant Ω, where Ω is a server preserving mutual authentication is provably secure based on the security model. Second, we make a further security analysis or a user. In this game, can make a set of queries and must answer them as described below.
Security Model. We assume that there exists an adversary A. All messages available in the network are also available to A. This includes all the messages sent by any set S∗ of users within the system. The main goal of A is to attack the scheme by decrypting any messages sent in the network intended to any set of users in S∗ but not him. A is considered to be successful if he wins the following interactive experiment. • Init: A picks a set of users S ∗ = {ID∗ ,⋯, ID∗} that he wants to attack (with n≤N) and sents S* to challenger C . • Setup: Challenger C runs the setup algorithm and sends adversary A the public parameters PK.
Security Model. Our security model is a standard model for Group Key Agreement protocols executed over authenticated links. Since the players in our GKA protocols do not use long-term secrets, This define GKA security.
Security Model. The completeness of fund transferring denotes that a sender(s) can send a valid transaction to the cash system and pre-transactions to the receiver(s). The receiver(s) can secure the received coins by publishing another transaction with new secret keys. We define the completeness of the whole process of multi-party fund transferring as follows, Definition 12 (Completeness of Multi-Party Fund Transferring). MCT is complete if, [⊥̸← NFT(pp), ⊥̸← TNFT(pp)] is always true when pp ← MCT.Set(λ) for any λ. NFT(pp): n ←−$ N , v ∈ V∗ , k ∈ Zn×|v| ρ ←−$ V∗ s.t. Σv ≥ Σ ρ, m ←−$ N |ρ|, C ← MCT.Coin(v, [Σn
Security Model. We now introduce our security models for the analysis of privacy-preserving key agreement (PPKA) protocols. Our first security experiment is based on standard key-exchange models in the tradition of Xxxxxxx-Xxxxxxx [4] key indistin- guishability games. This allows our model to easily capture known key secrecy, as well as generically capture key ran- domness notions, since our adversary is tasked merely with the goal of distinguishing the targeted session key from a random session key from the same distribution. Our second security experiment allows us to capture privacy notions of sessions, by challenging an adversary to determine which of two previously selected nodes ran a given protocol ex- ecution. Our cleanness predicates (see Section 5.4) allows us to model KCI attacks by allowing the adversary to re- veal the long-term key of the node running the PPKA pro- tocol, as well as the notions of partial forward secrecy. We HN, running a number of instances of the PPKA protocol Π , and a set of (up to) nN nodes N1, . . . , NnN (representing nodes communicating with the hub node HN), each poten- tially running one stage of (up to) nS consecutive stages of
