The present work deals with the modeling of multi-defected solids under the action of large deformation. A micromechanics constitutive model, formulated in terms of the compressible anisotropic NeoHookean strain energy density function, is presented to characterize the corresponding nonlinear effective elastic behavior. By employing a scalar energy parameter, a correspondence relation between the effective hyperelastic model and this energy parameter is established. The corresponding effective material coefficients are then evaluated through combined use of the “direct difference approach” and the extended “modified compliance contribution tensor” method. The proposed material constitutive model can be further used to estimate the effective mechanical properties for engineering structures with complicated geometry and mechanics and appears to be an efficient computational homogenization tool in practice.
Citation: Jui-Hung Chang, Weihan Wu. Evaluation of effective hyperelastic material coefficients for multi-defected solids under large deformation[J]. AIMS Materials Science, 2016, 3(4): 1773-1795. doi: 10.3934/matersci.2016.4.1773
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Abstract
The present work deals with the modeling of multi-defected solids under the action of large deformation. A micromechanics constitutive model, formulated in terms of the compressible anisotropic NeoHookean strain energy density function, is presented to characterize the corresponding nonlinear effective elastic behavior. By employing a scalar energy parameter, a correspondence relation between the effective hyperelastic model and this energy parameter is established. The corresponding effective material coefficients are then evaluated through combined use of the “direct difference approach” and the extended “modified compliance contribution tensor” method. The proposed material constitutive model can be further used to estimate the effective mechanical properties for engineering structures with complicated geometry and mechanics and appears to be an efficient computational homogenization tool in practice.
1.
Introduction
With the rapid growing of wireless networks, there is a corresponding increase in demand for multimedia-supported services, such as internet protocol television (IPTV) [1,2] video conference, and 3D holographic displays. Such services can be offered via voice over IP (VoIP) [2,3] services using session initial protocol (SIP) [3-7], a text-based signalling protocol. SIP has also been deployed for IP multimedia implementations [8,9], smart home and network management [10,11] and mobility management [12-14].
There are, however, open security and privacy challenges when transmitting voice packets over an open network. For example, designing a provably secure and efficient authentication protocol for SIP remains a challenging task. The original authentication mechanism in SIP was based on hypertext transport protocol (HTTP) [15] digested authentication, designed to provide only data authentication. Yang, Wang and Liu [16] introduced the first Diffie-Hellman authenticate key agreement protocol (AKAP) for SIP using the client-server model. However, it was later found to be vulnerable to server spoofing and off-line password guessing attacks [16]. Several elliptic curve based SIP AKAP has also been presented in the literature [17-24], due to the advantages elliptic curve based protocols offered over Diffie-Hellman based protocols (e.g., reduced computational and storage costs). For example, the key length for elliptic curve cryptography (ECC) is much smaller than RSA and ElGamal cryptosystem at the same level of security. Similar to the troubled history of Diffie-Hellman-based protocols, several published ECC-based protocols were found to be insecure after their publication. For example, the protocol of Wu, Zhang and Wang [25] was found to be vulnerable to a range of attacks in [26]. An improved protocol was then presented.
To minimize the impact of a security breach at the server, one of the many desired security features in SIP authentication is to ensure that users' passwords (which may be stored as plaintext) are securely stored. To withstand stolen verifiers attack at the SIP server, a number of smart card-based AKAPs were proposed in the literature, and some of them were subsequently found to be flawed [24,27-31].
More recently in 2016, Zhang, Tang and Zhu [32] introduced an energy-efficient AKAP for SIP. In this paper, we demonstrate that the proposed protocol is vulnerable to key-compromise impersonation attacks, in violation of their security claims. We then propose an ECC-based AKAP, and demonstrate its correctness and security respectively using BAN logic and the Automated Validation of Internet Security Protocols and Applications (AVISPA) simulation tool [33,34]. We also evaluate the performance of the proposed protocol.
2.
Preliminaries
The basic knowledge on ECC and some notations are introduced as follows [35].
Definition 1.Ep(a,b):y2=(x3+ax+b)modp be the form of an elliptic curve Ep(a,b) over Fp with a,b∈Fp satisfying (4a3+27b)modp≠0.
Assumption 1. Elliptic curve discrete logarithm problem (ECDLP): Known two points aP and P over Ep(a,b), to determine the random number a∈Z∗q. It is a hard problem in polynomial time such that the security is achieved.
Assumption 2. Elliptic curve Diffie-Hellman problem (ECDHP): Known two points aP and bP over Ep(a,b), to calculate the point abP, where the random numbers a,b∈Z∗q. abP can not be solved with non-negligible probability in polynomial time.
3.
Review of Zhang, Tang and Zhu's protocol
Zhang, Tang and Zhu's protocol consists of four phases: initiation, registration, authentication, and password change.
3.1. Initiation
S selects its private key as s, gets its public key as sP. Also, S releases the {Ppub,P,h(),E(Fp)}, where h() be the hash function.
3.2. Registration
Step 1. Ui computes C1=h(PWi⊕r) and sends {IDi,C1} to S via a secure channel. In the formula, r is a random number.
Step 2. S calculates C3=h(IDi⊕s)⊕C1 and sends back a smart card including {C3} to Ui.
Step 3. Ui writes r into the smart card, which has stored the values {C3,r,h()}.
3.3. Authentication
Step 1. The smart card asks Ui to enter: identity IDi and password PWi. Based on the two values, the smart card derives h(IDi⊕s) by computing C3⊕h(IDi⊕PWi). Subsequently, the card picks two random numbers r1, r2, and computes C4=r1P, C5=r1C2Ppub, and C6=h(C5)⊕(h(IDi⊕s)⊕r2,(C5)x,(C5)y). In the formula, (C5)x and (C5)y are x- and y-coordinate values of point C5. Then, Ui delivers a request message REQUEST {IDi,C4,C6} to S.
Step 2. S calculates C2=h(IDi⊕s) and retrieves (h(IDi⊕s)⊕r2,(C5)x,(C5)y) by computing h(sC2C4)⊕C6. Then S checks whether (C5)x,(C6)y?=(sC2C4)x,(sC2C4)x. If this holds, S derives r2 by means of computing C2⊕h(IDi⊕s)⊕r2. Then, S picks two random numbers r3, r4 and computes C7=r3P, SK=h(C4,r3C4,C7), and Auths=h(h(IDi⊕s),r2,(SK)x,(C5)x,(SK)y,(C5)y). Next, S replays a challenge message CHALLENGE {realm,C7,Auths,r4}.
Step 3. U computes SK=h(C4,r1C7,C7) and verifies whether h(C2,r2,(SK)x,(C5)x,(SK)y,(C5)y)?=Auths. If this holds, U computes Authu=h((SK)x,(r4+1),(SK)y) and transmits a response message REPONSE {realm,Authu} to S.
Step 4. Once S verifies whether h((SK)x,(r4+1),(SK)y)?=Authu. If the equation is correct, S and U successfully shares a common session key SK=r1r3P.
3.4. Password change
When Ui plans to change a password into a new one, he needs a helping of S. The process is as follows.
Step 1. Ui keys his old password PWi and retrieves Z=h(IDi⊕s) by computing h(PWi⊕r)⊕C3 and V=Enc(SK)x(h(PW∗i⊕r∗),IDi,R,Z). In the formula, PW∗i and r∗ are the new password and the random number. Next, Ui submits {V} to S.
Step 2. S computes h(IDi⊕s), and and examines whether it is equivalent to the decrypted value Z from V, using x-coordinate of point SK. Subsequently, S computes C∗3=h(PW∗i⊕r∗)⊕h(IDi⊕s) and W=Enc(SK)x(C∗3,h(C∗3,(R+1))). S then returns {W} to Ui.
Step 3. Ui checks whether h(C∗3,R+1) is equivalent to the derived one which comes from received message W. If it is true, Ui uses the values {C∗3,r∗} instead of the old one.
4.
Cryptanalysis of Zhang, Tang and Zhu's protocol
Before performing security analysis on the protocol of Zhang, Tang and Zhu, we assume that A could intercept any packets which are delivered in the authentication phase. Also, assume A can perform the corresponding computation.
Resilience to key compromise impersonation (KCI) attacked is considered as an important, and a desired security attributed to a key exchange protocol. It means an adversary A once has obtained one party's long-term private key. The saboteur may disguise another entity to the corrupted one. However, the three-round protocol of Zhang, Tang, and Zhu does not provide KCI-resilient. We show more detail on how A launches such an attack and achieves its goal. It breaks the confidentiality of the session key established.
Step 1. A compromises the perpetual secret key s of the server S. And A obtains the user's identity IDi utilizing intercepting the communication messages;
Step 2. Based on the two parameters, A calculates C2=h(IDi⊕s) and then generates two random numbers r′1, r′2;
Step 3. A computes C4=r1P, C5=r′1C2Ppub and C6=h(C5)⊕(h(IDi⊕s)⊕r′2,(C5)x,(C5)y) and transmits the REQUEST message {IDi,C4,C6} to S;
Step 4. S calculates C2 and verifies (Cs5)x,(C5)y?=(sC2C4)x,(sC2C4)y. If this holds, S computes C7=r3P, SK=h(C4,r3C4,C7), and Auths=h(h(IDi⊕s),r′2,(SK)x,(C5)x,(SK)y,(C5)y). In the formula, r3 and r4 are the random numbers. Next, S transmits the CHALLENGE message (realm,C7,Auths,r4) to Ui;
Step 5. A calculates SK=h(C4,r1C7,C7) and verifies whether h(C2,r2,(SK)x,(C5)x,(SK)y,(C5)y)?=Auths. Also, A computes Authu=h((SK)x,(r4+1),(SK)y) and sends the RESPONSE message (realm,Authu) to S.
Step 6. S needs to check whether h((SK)x,(r4+1),(SK)y) is equivalent to the received Authu. If it is false, S discards the packets; Otherwise, S accepts A, and agrees on the session key SK=r′1r3P as their subsequent shared key.
More seriously, A could even guess the correct password PWi once the secret key of S is compromised. A is by checking whether h(PW∗i⊕r)⊕h(IDi⊕s) until the equation holds. In this formula, PW∗i is an arbitrary element that is selected from the password candidate set.
5.
The proposed protocol
To solve the security problems found in Zhang, Tang and Zhu's protocol, we develop a secure SIP authentication protocol with a key agreement facility. Unlike Zhang, Tang and Zhu's design, our proposal comprises of five phases: initialization, registration (Table 1), login (Table 2), key agreement (Table 3) and password update, where the password update process is concise and convenient, that lying in it does not require interaction with the server. The proposed protocol is explained below along with Figure 1.
S selects an additive group G, with a generator P of large prime order q, including points of an elliptic curve E over a finite field Fp. S publishes the public parameters {h(),sP,P,E,q}, where s∈Fp,P∈Ep(a,b).
5.2. Registration
Once a new user Ui attempts to access services, he first selects his identity IDi and password PWi of his/her choice.
Step 1. The user Ui picks out a random number r, and calculates h(PWi,r). Ui then transmits the registration request message {IDi} to the proxy server, through a private channel.
Step 2. The proxy sever S calculates X1=h(IDi,s)P and X2=h(IDi)nsP. In the formula, n is a random number. Next, S personalizes Ui's smart card which remains the values {X1,X2} to Ui via a secure channel.
Step 3. Ui computes X3=h(PWi,r)h(IDi,s)P and X4=h(h(IDi)⊕h(PWi,r)). Ui finally stores r into his smart card. Note that the smart card includes the information {r,X2,h(),X3,X4}.
5.3. Login
Ui enters his identity IDi and password PWi after putting the smart card into the card slot.
Step 1. The smart card verifies whether the condition h(h(IDi)⊕h(PWi,r))?=X4 holds.
Step 2. If the equality holds, the smart card then gets h(IDi,s)P by calculating h(PWi,r)−1X3 and nsP=h(IDi)−1X2. Also, a random number r1 is generated, the value X5=EncnsP(IDi,r1P,sP) to verify.
Step 3. Ui delivers the login request message REQUEST {X5} to S through a public path.
5.4. Key agreement
Step 1. S uses its private key s and secrets parameter n to retrieve (IDi,r1P,sP).
Step 2. S proceeds to generate two random numbers r2, r3, and calculates the temporary key SK=h(r2r1P,IDi) to be shared with the user Ui. Also, the smart card calculates X6=Ench(IDi,s)P(r2P,r1P) and Auths=h(SK,r1P,r3). And then the smart card responds with the challenge message CHALLENGE {realm,X6,Auths,r3} to Ui.
Step 3. The smart card gets the values (r2P,r1P) by decrypting X6. And then, the temporary key SK=h(r2r1P,IDi) is gotten by Ui and to be shared with the server S. Verifying whether h(SK,r1P,r3)?=Auths. If the condition holds, Ui deems S as the legitimate server.
Step 4. The smart card computes Authu=h(SK,r2P,r3+1), and transmits a response message RESPONSE {realm,Authu} to S through a public channel.
Step 5. S examines the verification condition h(SK,r2P,r3+1)?=Authu. If this equation holds, S ensures Ui as authentic and agrees on the session key SK as valid key.
5.5. Password updating
The following mechanism achieves altering the password of a legal user Ui without interacting S.
Step 1. Ui puts his smart card into the card slot and waits for commands of the terminal to provide the identity IDi and password PWi. The smart card verifies that if the condition h(h(IDi)⊕h(PWi,r))?=X4 holds. If the validation does not validate, the session is quitted promptly. Otherwise, the smart card derives h(IDi,s)P by computing X3h(PWi,r)−1, and requests a new password.
Step 2. Ui picks his new password PW∗i and the random number r∗. The smart card calculates X∗3=h(PW∗i,r∗)h(IDi,s)P and X∗4=h(h(IDi)⊕h(PW∗i,r∗)).
Step 3. The smart card discards X3 and X4 but keeps X∗3 and X∗4 in its memory for renewal.
6.
Security analysis of the proposed protocol
We confirm our proposal could achieve a mutual handshake using well-popular BAN logic [36]. Also, the robustness of our proposal is validated via the universally applicable simulation tool-AVISPA [33,34]. In addition, we provide informal cryptanalysis so as to demonstrate our proposal is well protecting against relevant security attacks.
6.1. Verification of the proposal under BAN Logic
BAN logic is a well-known formal method used to strictly prove the authentication protocols' security, or find security vulnerabilities. Subsequently, we will introduce more details about BAN logic. BAN logic includes basic logical notations and some logic postulates. According to these preliminaries, we show the desired goals, idealized form, assumptions for our protocol. And we finally demonstrate its correctness.
Notations
⋅P|≡X: P deems X is true;
⋅P◃X: P observes X;
⋅P|∼X: P ever have sent X;
⋅P⇒X: P judges X;
⋅#X: X is fresh;
⋅PK↔Q: share a key K between P and Q;
(X,Y): X or Y is one portion of the formula (X,Y);
⋅ From P2, A4 and fresh conjuncatenation rule, we derive
P3. S|≡#(IDi,r1P,sP)
⋅ Since P2, A3 and nonce-verification rule, we get
P4. S|≡Ui|≡(IDi,r1P,sP)
⋅ According to P4 and belief rule, we get
P5. S|≡Ui|≡IDi, S|≡Ui|≡r1P
⋅ From A8, A9 and jurisdiction rule, we derive
P6. S|≡r1P, Goal1. S|≡IDi
⋅ Since SK=h(r1r2P,IDi), Goal1, and P6, we have
Goal2. S|≡UiSK⟷S
⋅ According to message Authu, we have
P7. S◃<r2P,r3+1,SK>UiSK⟷S
⋅ By P7, Goal2 and message-meaning rule, we obtain
P8. S|≡Ui|∼(r2P,r3+1,UiSK⟷S)
⋅ Since P8, A10, A11 and nonce-verification rule, we have
Goal3. S|≡Ui|≡UiSK⟷S
⋅ From message X6, we attain
P9. Ui◃<r2P,r1P>Uih(IDi,s)⟷S
⋅ According to A7, P9 and message-meaning rule, we get
P10. Ui|≡S|∼(r2P,r1P)
⋅ By A2 and fresh conjuncatenation rule, we have
P11. Ui|≡#(r2P,r1P)
⋅ According to P10, P11 and nonce-verification rule, we attain
P12. Ui|≡S|≡r2P
⋅ Since P12, A11 and jurisdiction rule, we derive
P13. Ui|≡r2P
⋅ From SK=h(r1r2P,IDi), A1, A2, P13, we have
Goal4. Ui|≡UiSK⟷S
⋅ By message Auths, we attain
P14. Ui◃<r1P,r3,SK>UiSK⟷S
⋅ Since P14, Goal4 and message-meaning rule, we have
P15. Ui|≡S|∼(r1P,r3,UiSK⟷S)
⋅ According to P13, P15, Goal4 and nonce-verification rule, we attain
Goal5. Ui|≡S|≡UiSK⟷S
6.2. Formal security analysis
Theorem 1. The probability that an attacker A breaks the AKE security of our AKAP is
AdvAkeP(A)≤q2h2l−1+2qsend|D|
where qsend, qh and D denote the number of Send queries, Hash queries, and a uniformly distributed dictionary, respectively.
Proof: Game Gi(i=0,1,2) defines three games. Game G0 is the factual attack, and game G3 concludes a breach of the AKE security of our AKAP is asymptotically optima:
Game G0: This game corresponds to the actual attack.
AdvAkeP(A)=|2Pr[Succ0]−1|.
Game G1: This game simulates the eavesdropping attack by querying Execute(Ui,Sj) oracle, and then by querying Test(Pi) oracle. It decides whether the result of Test is the real session key SK or a random value. We know that r1P is derived by the server's secret key s, and secrets parameter n. That is, A has no way to compute r1P through eavesdrop on the communication channel unless S is compromised. Also, r2P is not impossible to obtain, unless it possesses both the smart card and password. Hence, intercepting is not probable for helping A to win in this game. Thus,
Pr[Succ1]=Pr[Succ0].
Game G2: This game models Send(M,Pi) query, in which A can eavesdrop or alter the information from the transcripts. Then, games G2 and G1 are undistinguishable unless the collision occurring in G2. Thus,
|Pr[Succ2]−Pr[Succ1]|≤q2h/2l.
Game G3: This game models Corrupt(SC) query, in which A has obtained the smart card to simulate the smart card breach attack. Since the password PWi is protected by a cryptographic one-way function, where X3=h(PWi,r)h(IDi,s)P and X4=h(h(IDi)⊕h(PWi,r)). This implies that A has no way to check the password excepts possession of user's identity, or corrupts the server to get s. Hence, |Pr[Succ3]−Pr[Succ1]|≤qsend/|D|.
6.3. Verifying protocol using AVISPA tool
AVISPA is a simulation engine for the automated validation of Internet security protocols and applications. Upon Dolev and Yao model, four model back-ends, called OFMC (On-the-fly Model-Checker), CL-AtSe (Constraint-Logic-based Attack Searcher), SATMC (SAT-based Model-Checker), and TA4SP (Tree Automata-based Protocol Analyzer) (Figure 1) are utilized for the validation using HLPSL (High-Level Protocol Specification Language). The HLPSL presentation of the protocol is compiled to IF by the translator-HLPSL2IF. IF is an entrance of the four different back-ends. The output OF is exported by using one of the four back-ends, which shows the conclusion if the AKAP is secure or insecure.
During the protocol execution, each entity act a role, which is a feature of AVISPA. We show the role specifications in HLPSL of Appendix the initiator, responder, session, and environment and goal in Appendix Figures 1–4. In our implementation, we assume S's private key is a public parameter. The privacy of three parameters and client-server authentication is verified:
⋅ The secrecy_of subs1: the user confidential parameters IDi is gotten only Ui and S.
⋅ The secrecy_of subs2: the user confidential parameters PWi is gotten only Ui.
⋅ The secrecy_of subs3: the session key is gotten only Ui and S.
⋅ Authenticaion_si_ui_auths: Ui validates S by receiving r2 securely, r2 is a ephemeral number of S.
⋅ Authenticaion_si_ui_authu: Ui validates S by receiving r1 securely, r1 is a ephemeral number of S.
After running the program under two back-ends CL-AtSe and OFMC, Figures 2 and 3 show that our proposal realizes the session security without imperfection.
Figure 2.
Simulation result in CL-AtSe model checker.
We demonstrate that our proposal holds many security attributes, such as mutual authentication, anonymity, privileged insider attack, perfect forward secrecy, KCI-resistance, etc., under a condition. The condition is that A extracts all the data stored inside a user's smart card, or/and eavesdrops on all the messages involved in an authentication-key agreement session [37,38].
6.4.1. Mutual authentication
Note that r1P can be only decrypted by the legal server. Thus, after Ui receives CHALLENGE message depending upon the result of decryption test, Ui verifies the legitimacy of S by checking the equivalence h(SK,r1P,r3)?=Auths. Simultaneously, only the legal Ui can derive r2P from CHALLENGE message. And hence, S assures that he is communicating with the legitimate Ui, employing checking the equivalence h(SK,r2P,r3+1)?=Authu after receiving RESPONSE message. Therefore, A cannot cheat any of the communicating entity. And thus the proposed protocol achieves proper mutual authentication.
6.4.2. Internal-privileged attack
Ui's password PWi is hashed by the random number r during registration phase. And it is not delivered to S. Therefore, an honest but curious insider has no ability, to know the real password PWi of Ui. In other words, the proposed protocol indicates good capability of defeating insider attack.
6.4.3. User anonymity
The identity IDi of Ui is disguised for dual protection, involved in all the transmitted messages. For one thing, the random number n is picked excepts the private key s of S, based on the ECDLP assumption. For another, the plain-text IDi is hashed by double times along with the secret parameters r1P, and r2P. And the two parameters are encrypted by nsP. That is, to identify two parameters are equate as determine ECDLP problem. The problem is one of the well-known difficult problems within polynomial time. In a word, the proposed protocol supports high user anonymity.
6.4.4. Perfect forward and backward secrecy
Suppose that perpetual privacy information s of S is compromised by A, he is incapable of computing the current as well as the future session keys. It is noteworthy that the session key is related with three important parameters, i.e, Ui's identity, two random numbers r1 and r2 generated by Ui and S, which are present in the form of r1P and r2P, respectively. With the purpose of acquiring IDi and r1P from intercepted REQUEST message {X5=EncnsP(IDi,r1P,sP)}, another random number n is also needed. Unfortunately, no one but S knows what is the real random number n. More seriously, r2P can not be derived through X6=Ench(IDi,s)P(r1P,r2P) without knowing IDi of Ui. Next, assume that the current session key is corrupted by A. He plans to deduce the next negotiatory key. The cause of failure of extraction the secret SK=h(r1P,r2P,IDi) lies in the irreversible property of hash function, he total has no way to compute the previous and future session keys thereby. In a conclusion, our protocol preserves perfect forward secrecy property.
6.4.5. Resist modification attacks
We assume A intercepts the REQUEST message {X5} submitted to S, he attempts to modify the parameter r1P and transmit the forged message {X5} to S. However, he has no way to attain IDi and n to compute the symmetric key nsP, and thus encrypt IDi. Even with the smart card security breached, he is not strong enough to gain the exact values of IDi, and s without the knowledge of PWi. On the other hand, without knowing IDi, n and s, it is impossible for A to learn r1P picked by Ui intercepting CHALLENGE message. That is, A could neither impersonate as an authorised user nor masquerade as a legitimate sever, through eavesdropping the raw messages to tamper with them. In a word, our proposal is is resistance to modification attacks.
6.4.6. Resist off-line password attack
We suppose A gets {r,X2,X3,X4,h()} and all the public messages {X5,X6,Auths,r3,Authu} from the communication channel. First and foremost, the password PWi of Ui do not involve in transcription between Ui and S. Hence, A has no way to ensure whether the guesses password is true or false. Secondly, to derive the correct password PWi from X3 is a computationally infeasible task for A, in condition that he is unaware of the identity IDi of Ui, and the private key s of S. Eventually, A is not entirely sure what are the real values of IDi, and PWi by X4, because the two personal information are hidden in two-layer hash function. In conclusion, the proposed protocol could thwart off-line password attack successfully.
6.4.7. Resist key-compromise impersonation attack
The whole design presumes that the privacy information s is considered as an open parameter. In this case, A has no way to impersonate as an authorised user thereby accessing service resources. The resources are provided by the secure server. Let's analyze the cases. Aiming at playing a valid user, the identity of the real user Ui is urgent required, since the server will detect the attack while checking Authu. To get the real value of IDi, another random number n, which only the legal server knows it, is also needed. Thus, it is not possible for A to try to impersonate as a legitimate user.
6.4.8. Session key agreement
Two parties independently negotiate an ephemeral session key SK=h(r1r2P,IDi). They keep communicate securely on the strength of it for a subsequent communication. Upon this ability, each entities can encrypt the following packets, to preserve the security of the handshake. Moreover, the negotiatory key is fresh for each session. The reason is that random numbers are different based on the property of hash function. As a consequence, deriving the session key through the eavesdropped information is a challenging task for A.
6.4.9. Stolen-verifier attacks
The service provider does not create the password verification table of the service requester. Even though the service provider's database is available by A, he still cannot steal and modify user passwords, and thus attain the authentication information of users. Thus it can be said that our proposal can withstand stolen-verifier attacks.
7.
Performance considerations and functionality comparison
This part evaluates and examines the performance of our proposed scheme, and compare it with five related protocols relies basically on the ECC, one-way hash functions. In order to evaluate the entire computing cost for each protocol more accurately, the arithmetic mean for each cryptographic computation timings after running 1000 times are shown in Table 4. The processor is Intelr Pentiumr CPU G3250, 3.20GHz with 4.0GB of RAM running Windows 10. We use the jPBC library primitives timings. The processor is Intelr Pentiumr CPU G3250, 3.20GHz with 4.0GB of RAM running Windows 10. We use the jPBC library(2.0.0) [41], a Java port of the PBC library written in C [39,40], the Java Development Kit used is the Oracle jdk 1.8.0 65. We used the Type A curves with the prime order q defined as E(Fq):y2=x3+x, hash function as SHA-3 [19,42], and symmetric encryption algorithm as AES [43]. The calculation expense and execution time of the registration and key agreement phases with the revelent protocols [24,27-30,32] are listed in Table 5. Also, we compare the security attributes with the related protocols [24,27-30,32] (Table 6). The entire running time of our AKAP is lower than Tu et al. [24]. And the time is lower than Yeh, Chen and Shih's protocol[28] too. Table 5 has shown the results. However, one thing is ignored by these protocols [24,27-30,32] from Table 6. It is to analyze whether the design has ability to conquer key compromise impersonation attack. Additionally, the related protocols seem not to be considered as a SIP authenticated key agreement, which can be perfect or ideal. The reason is that the protocol lacks some essential security properties. In general, our proposal takes a better tradeoff between computational cost and security attributes, while comparing with the protocols [24,27-30,32].
Table 4.
jPBC library primitives timings.
Operation
thtp
tsm
tpa
th
tinv
tsym
Aritmetic mean
10.8966ms
10.5129ms
0.4338ms
0.0359ms
0.0428ms
0.1755ms
Note: thtp: executing a hash to point operation; tsm: executing an elliptic curve scalar multiplication; tpa: executing an elliptic curve point addition; th: executing a hash function operation; tinv: executing a modular inversion; tsym: executing a symmetric encryption/decryption
Despite the role that AKAP plays in ensuring the security of communication in an open network, designing secure and efficient protocols, including in a VoIP environment, remains challenging. For example, in this paper, we revisited and revealed vulnerabilities in the design of Zhang, Tang and Zhu's protocol [32]. We also presented an improved protocol and demonstrated its correctness and security, as well as demonstrating its utility in terms of performance efficiency. Future research will include implementing a prototype of the protocol and evaluating it in a real-world deployment.
Acknowledgments
The authors would like to thank all the editors and anonymous reviewers for their helpful advice. This paper is supported by the National Natural Science Foundation of China (No. 61802276), and the Fundamental Research Funds for the Central Universities of China (No.3122021027).
Conflict of interest
The authors declare that they have no known competing financial interest or personal relationship that could have appeared to influence the work reported in this paper.
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Note: thtp: executing a hash to point operation; tsm: executing an elliptic curve scalar multiplication; tpa: executing an elliptic curve point addition; th: executing a hash function operation; tinv: executing a modular inversion; tsym: executing a symmetric encryption/decryption
Note: thtp: executing a hash to point operation; tsm: executing an elliptic curve scalar multiplication; tpa: executing an elliptic curve point addition; th: executing a hash function operation; tinv: executing a modular inversion; tsym: executing a symmetric encryption/decryption
Figure 1. A cutoff area Ω (i.e., the shaded region) is delimited in a homogeneously-stressed infinite medium
Figure 2. A single inclusion is taken as the whole cutoff area Ω in a plane stress specimen
Figure 3. Three local near-defect FE meshes for the specimen in Figure 2
Figure 4. A plane stress specimen with the cutoff area W containing a family of parallel cracks
Figure 5. The normalized effective elastic moduli versus the crack density parameter f (Problem 2.1)
Figure 6. The results of Weff/Λ0 (evaluated with extended MCCT) versus λmax, along with the fitted curves (f = 0.367)(Problem 2.2)
Figure 7. The normalized effective coefficients versus the crack density parameter f(Problem 2.2)
Figure 8. A plane stress specimen with the cutoff area Ω containing a central elliptical void
Figure 9. The deformed mesh for the specimen in Figure 8 for λmax = 3. (a) uniaxial (x1-direction), (b) biaxial
Figure 10. The distribution of maximum principal Cauchy stress in the near-void area for λmax = 3. (a) uniaxial (x1 -direction) (b) biaxial
Figure 11. The results of Weff/Λ0 versus λmax, along with the fitted curves (Case Ⅲ) (Problem 3.2)
Figure 12. (a) A body, of infinite extent in the x1- direction, contains a thin weak material layer. A representative column is chosen for FE analysis. (b) A square RAE is delimited for modeling the microstructure of a material layer containing periodically-distributed identical elliptical voids
Figure 13. (a) The FE mesh for a representative column (d= h0/20). (b)-(d) The deformed FE mesh (no void, circular void, and elliptical void)
Figure 14. The reactions at the upper boundary versus the applied displacements
Figure 15. The variations of R1/R1, ref and R2/R1, ref with respect to h
Figure 16. The variations of R1/R1, ref and R2/R1, ref, at (u1, u2) = (0.1, 0.01) h0, with respect to d