Security Model Clause Samples
The Security Model clause defines the standards and protocols that must be followed to protect data, systems, or assets from unauthorized access, breaches, or other security threats. It typically outlines the technical and organizational measures required, such as encryption, access controls, regular security assessments, and incident response procedures. By establishing clear security expectations and responsibilities, this clause helps prevent data breaches and ensures both parties understand their obligations to maintain a secure environment.
Security Model. We assume that the reader is familier with the model of ▇▇▇▇▇▇▇ 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. 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. ▇▇▇▇▇▇▇▇▇▇, 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. 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. Before going to prove that the session key security is preserved by the proposed scheme, we discuss describe the ROR model [46]. • Participants. Let V , Dj , and CC denote the αth Lemma 1 (Difference Lemma): Let A, B, F denote the events defined in some probability distribution, and assume instance of vehicle Vi, the βth instance of drone Dj and that A ∧ ¬F ⇐⇒ B ∧ ¬F . Then | Pr[A] − Pr[B] ≤ Pr[F ]. the γth instance of CC, respectively. These instances are named the oracles. • Accepted state. If an instance V α jumps to the accepted state after the last expected protocol message is received, it will be in the accepted state. The session identification (sid) of V α for the current session that is constructed
Security Model. Players. We assume that two users A and B participate in the key agreement protocol P. Each of them may have several instances called oracles involved in distinct executions of P. We denote instance s of i ∈ {A, B} by Πs for an integer s ∈ N. We also use the notation Πs to define the s-th instantiation of A executing with B. Adversarial Model. We allow a probabilistic polynomial time (PPT) adver- sary F to access to all message flows in the system. All oracles only communicate with each other via F . F can replay, modify, delay, interleave or delete messages. At any time, the adversary F can make the following queries: – Execute(A, B ): This query models passive attacks, where F gets access to an honest execution of P between A and B by eavesdropping. – Send(Πs, m): This query models F sending a message m to instance Πs.
Security Model. We consider a hybrid security model by combining a computational assumption, that there exist a short-term- secure computational encryption, and conversely assum- ing that any optical quantum memory is technologically bound to decohere within a timescale shorter than the time for which the computational encryption is secure. This new, Quantum Computational Hybrid (QCH) secu- rity model, is formally defined as:
Security Model. We formulated a series of games between challenger ▇ and adversary 𝐴 to define our security model. Assume that participant ∏𝑖 ∈ {𝑈, 𝐸𝐷, 𝐶𝑆} represents the i-th instance and 𝛬 represents the entire protocol. The 𝐴 can ask the 𝐶 oracle queries, and the C can respond. ⚫ 𝑆𝑒𝑛𝑑(∏𝑖, 𝑚): If A asks the query for the message 𝑚, the 𝐶 executes the specific steps of the protocol and returns the result. ⚫ 𝐸𝑥𝑒𝑐𝑢𝑡𝑒(∏𝑈 , ∏𝐸𝐷 , ∏𝐶𝑆): This oracle query models a
Security Model. The security of ring signature schemes is defined via the following notions.