==Phrack Inc.==
Volume 0x0e, Issue 0x44, Phile #0x0e of 0x13
==
==[ Secure Function Evaluation vs. Deniability ]==
==[ in OTR and similar protocols ]==
==
==[ greg ]==
==
[ Contents
1  Introduction
1.1  Prelude
2  Preliminaries
2.1  DiffieHellman
2.2  RSA
2.3  Oblivious Transfer
2.4  Secure Function Evaluation
3  OTR
4  The Attack
4.1  Sharing DiffieHellman Keys
4.2  Generating MAC and Encryption Keys
4.3  Sending and Receiving Messages
4.4  The Final Protocol
4.5  What's Left
5  References
6  Greetingz
[ 1  Introduction
Recent cryptographic primitives and protocols offer a wide range of
features besides confidentiality and integrity. There are many protocols
that have more advanced properties, such as forward secrecy, deniability or
anonymity. In this article, we're going to have a deeper look at
deniability in communication (e.g. messaging) protocols. One protocol that
claims to offer deniability is OTR. Although our construction can probably
be extended in a quite general way, we'll stick with OTR as an example
protocol. Our goal is to show the limits of deniability, especially in
protocols that offer message integrity features (as OTR does). We will do
this by constructing a protocol that enables each partner in a conversation
to cooperate with an observing party, such that he can prove the
authenticity of any message that was part of the conversation to the
observing party.
[ 1.1  Prelude
It was one of these days sitting together with bruhns and discussing stuff
(TM). Out of the sudden, he came up with the question: "You know, I'm
asking myself what a trusted timestamping service could be good for...?". I
told him "timestamps, most probably". He was like "Uhm, yes. And wouldn't
that affect the deniability of OTR somehow?". We discussed the matter for
quite a while and we finally agreed that a trusted timestamping service
itself wouldn't be enough to destroy the deniability of OTR. But our
interest remained...
[ 2  Preliminaries
In this section, we're going to give a quick overview of cryptographic
primitives we're gonna use. If you're already familiar with those, you can
happily skip our explanations and get to the real meat. The explanations
in this section will not contain all the mathematical background (i.e.
proofs ;) ), which is necessary to really *understand* what's going on.
We'd rather like to provide a highlevel overview of how all the individual
components and how they can be combined.
[ 2.1  Symmetric Operations
We'll keep this real short; you probably know the most common symmetric
crypto algorithms. We will be using symmetric block ciphers (such as AES)
and hash functions (SHA for instance). Also, we will need MAC functions in
the following sections. You might already know HMAC, which is a MAC scheme
based on hash functions. MACs (Message Authentication Codes) are used to
protect the integrity of messages. Being a symmetric primitive, creating
and verifying the MAC requires knowledge of the same key. If someone can
verify a MAC, they can also create one.
[ 2.1  DiffieHellman
The DiffieHellman scheme is one of the most widely used key establishment
protocols today. The basic idea is the following: Alice and Bob want to
securely establish a key over an insecure channel. DiffieHellman enables
them to do this. During such a keyexchange both parties publicly send some
values and after the communication is finished, both can compute a common
key, which can *not* be computed by anyone who wiretaps the communication.
[ 2.1.1 The Math behind it
Alice and Bob agree on a prime p and some "generator" g. We won't discuss
too many details of the mathematical background here (if you're interested
in math, refer to [1]), so it's sufficient to say that in practice, g will
often have the value 2 and the prime p will be large. In many cases, p and
g are fixed parameters, on which both parties rely. Before describing the
actual protocol, we want to show one interesting observation: Given some
number x, it's trivial to compute values y = g^x mod p ("square and
multiply" are the magic words). Given the value y however, it's not trivial
at all to compute the value of x ("discrete logarithm problem", if you're
interested). This property can be used to build a keyestablishment scheme
like this:
A  a = g^x mod p > B
A < b = g^y mod p  B
A picks a random x, computes a = g^x mod p and sends that value over to B.
B picks a random y, computes b = g^y mod p and sends that value over to A.
The values a and b are also referred to as DiffieHellman public keys.
A now performs the following computation:
(2.1.1) ka = b^x mod p
B does the same and computes
(2.1.2) kb = a^y mod p
We can observe that due to the equation
(2.1.3) ka = b^x mod p = (g^y)^x = g^(yx) = g^(xy) = (g^x)^y = a^y = kb
ka and kb are equal. So A and B have established a common key k
(k = ka = kb). As an attacker however neither knows x nor y, he cannot
perform the same computation. The attacker could try to obtain x from a,
but as we outlined above, this is (hopefully) computationally infeasible
for large primes p and good generators g. In case of an active attacker,
this scheme can be broken by a simple maninthemiddle attack, where the
attacker replaces Alice's and Bob's values by his own ones and then proxies
the traffic between both parties. This problem can be fixed by making use
of an authentication scheme: Alice and Bob need to "sign" the values that
they transfer, so that the attacker cannot modify them without destroying
the signature. There are many signature schemes out there (for instance
based on RSA, which is described below) and all of them come with
additional costs (you need to exchange public keys beforehand etc.). We
assume you know about all the higherlevel problems, such as key
distribution, revocations, trustmodels, etc. The basic principle of
DiffieHellman however stays the same  and that is what we're going to
focus on later in this article.
[ 2.2  RSA
Another gem of modern cryptography is the RSA crypto system. RSA is also
based on modular arithmetic, but it works in a different way than
DiffieHellman. Alice wants Bob to send her an encrypted message. However,
Alice and Bob have not exchanged any key material (if they had, Bob could
just make use of any blockcipher like AES to send encrypted data to
Alice). With RSA, Alice can send Bob a thing called her "public key". This
public key can be used by Bob to encrypt messages. However, nobody can
decrypt messages encrypted with Alice's public key without knowing another
piece of information called Alice's "secret key". As the name suggests,
Alice keeps her secret key secret. Therefore everybody can encrypt messages
for Alice, but nobody besides Alice can decrypt these messages.
[ 2.2.1 More Math
Alice wants to receive messages from Bob, so she first needs to generate an
RSA keypair. Alice does the following: She picks two primes p and q and
computes
(2.2.1) N = p * q
She picks a value e (in practice, e = 65537 is a common choice) and
computes
(2.2.2) d = e^1 mod (p1)(q1) (i.e. e*d = 1 mod (p1)(q1))
This computation can be performed efficiently using the extended euclidean
algorithm (but again, we won't dive into all the mathematical details too
much). Alice keeps all the values besides N and e secret.
A  N = p * q, e > B
A < c = m^e mod N  B
Alice now sends over N and e to Bob. Bob uses N and e to encrypt his
message m as follows:
(2.2.3) c = m^e mod N
Then, Bob sends the ciphertext c over to Alice. Alice can use d to decrypt
the ciphertext:
(2.2.4) m = c^d mod N
This works due to the way e and d are chosen in equation (2.2.2).
To decrypt the ciphertext, an attacker could of course try to compute d.
But computing d is hard without knowing p and q. And obtaining p and q from
N is assumed to be an infeasible problem for large values of N.
The tuple (N, e) is commonly called an RSA public key, whereas (N, d) is
called private key. We can view an RSA instance (with fixed keys) as a set
of two functions, f and f^1, where f is the function that encrypts data
using the public key and f^1 is the function that decrypts data using the
private key. We'll call such functions oneway functions.
Instead of encrypting data with the receiver's public key, we can also use
RSA as a signature scheme. The signer of a message first uses a hash
function on his message. He then encrypts the hash value with his private
key. This signature can be verified using the signer's public key: the
verifier uses the public key to decrypt the hash value, computes the hash
of the message he received and then compares the hashes. An attacker will
not be able to produce such a signature, because he doesn't know the
signer's private key.
Please be aware that (like all the other algorithms described in this
document), RSA should in practice not be used as described above. In
particular, we did not describe how to correctly convert messages into
numbers (RSA operates on natural numbers, remember?) and how to securely
pad plaintexts. Depending on the respective security goals, there are a
number of possible padding schemes (such as OAEP+), but we're not going to
describe them here in detail.
[ 2.3  Oblivious Transfer
Oblivious transfer is a real funny primitive. Suppose, Bob knows two values
x0 and x1. Alice wants to obtain one of those values, but she doesn't want
to tell Bob which value she wants. Now Bob could of course tell Alice both
values (that way he wouldn't know, which one Alice was interested in).
However, Bob wants to make some money and so he takes $1k per value. Poor
Alice however only has $1k, so she can't afford to buy both values from
Bob. This problem can be solved with an oblivious transfer. An oblivious
transfer is a cryptographic protocol, so it requires a number of messages
to be exchanged between Alice and Bob. After the messages are exchanged,
Alice will receive the value she wanted and Bob won't know which value that
was.
[ 2.3.1 Math Voodoo
There are a number of protocols for performing an oblivious transfer, based
on different cryptographic assumptions. We are going to describe one
classical example here, which can be implemented using a publickey
cryptosystem (such as RSA). More details of this construction can be found
in [7].
The system works like this: Bob picks oneway functions f, f^1 and sends f
over to Alice. Along with f, he sends two random values r0 and r1. You can
think of f and f^1 as RSA functions using fixed keys (as described above).
A < f, r0, r1  B
A  z = f(k) XOR rb > B
Alice wants to receive value xb (b = 0 or 1) from Bob. She picks a random
k, computes f(k) and XORs it with r0 if she wants to receive x0 or with r1
if she wants to receive x1. The XOR operation is sometimes also called
"blinding". Depending on the cryptosystem that is used to obtain f and
f^1, there might be more appropriate choices then just using XOR. For RSA,
it would be natural to use integer addition and subtraction (modulo N)
instead of the XOR operation.
Alice now sends the result z to Bob. Bob performs some computations:
(2.3.1) k0 = f^1(z XOR r0)
(2.3.1) k1 = f^1(z XOR r1)
One of the k values will be Alice's, but Bob doesn't know which one. The
other value will be junk, but it's important to note that this junk value
cannot be computed by Alice (she doesn't know f^1). Now Bob simply does
the following:
A < x0 XOR k0, x1 XOR k1  B
Depending on which k value is the one that Alice actually knows, she can
decrypt x0 or x1. And that's it: Alice now knows the value she wanted to
receive and one junk value, which doesn't tell her anything. Bob however
doesn't know which of the k values was the one that Alice picked, so he he
cannot tell, which value Alice wanted to receive.
Let's try it out:
Say Bob hast two values x0 = 7 and x1 = 1. He is willing to share one with
Alice. First he generates f and f^1. To do that, he just uses RSA. He
picks two prime numbers p = 5 and q = 11 and gets N = 55. Also, he picks
e = 3 as encryption exponent (don't do that at home, kids!). The
decryption exponent would then be d = 27 (you can compute that using the
euclidean algorithm or alternatively you could just believe us). Bob now
can send out (N, e) = (55, 3) to Alice, along with some random values
(r0, r1) = (4, 9).
Suppose Alice wants to retrieve the value of x1. First of all, she picks a
random k, let's say k = 6. She encrypts it using the public key Bob sent
(i.e. she applies Bob's oneway function): f(6) = 6^3 mod 55 = 51. She
computes z = f(k) + r1 = 51 + 9 mod 55 = 5, which she sends to Bob.
Bob now determines his candidates for k (i.e. k0 and k1) by computing:
k0 = f^1(z  r0) = (5  4)^27 mod 55 = 1
k1 = f^1(z  r1) = (5  9)^27 mod 55 = 6 < Alice's k, but Bob doesn't
know that
Bob then sends to Alice: x0 + k0 = 7 + 1 and x1 + k1 = 1 + 6.
Alice receives the two values 8 and 7. She knows that Bob's second value
was x1 + k. As she is interested in x1, she takes that value and computes
x1 = (x1 + k1)  k = 7  6 = 1 (observe that k = k1, which only Alice
knows). Now Alice could try to cheat and to also obtain x0. But to do that,
she would need to know the value that Bob computed for k0, which she won't
be able to compute without knowing f^1 (i.e. the secret exponent d in our
case).
[ 2.4  Secure Function Evaluation
Secure function evaluation is another real gem of modern cryptography. A
classical example is the 0day problem. Two hackers A and B have a certain
number of 0day exploits each. They want to determine who is more elite, but
they are so paranoid that they don't even want the other to know how many
0days they have. So, A knows some number x, B knows y and both want to
compute the function f(x, y) = { 1 if x > y, 1 if y > x and 0 otherwise}.
Secure Function Evaluation solves this problem. Again, both parties
exchange a number of messages and after this is done, both of them know the
result of the function without having learned anything about the input of
the other party. And instead of the function shown above, they could just
arbitrarily agree on any function to be evaluated.
One interesting practical application of SFE is to perform mutual
authentication based on a shared secret. Two parties knowing a shared
secret can jointly compute the function
f(x, y) = {1 if x = y, 0 otherwise}. Interestingly, the OTR protocol makes
use of such a SFE scheme for authentication.
[ 2.4.1 More Voodoo
Suppose there is a function f(x, y), which two parties want to compute
(this is actually secure twoparty computation, which is not the most
general case  for our purpose however it is sufficient). Both want to
share the result and both want to keep their own inputs safe. There are
several constructions that allow us to perform SFE. We'll discuss only one
of them here: Yao's garbled circuits [3]. As the name suggests, the
function to be evaluated by both parties first has to be transformed to a
boolean circuit. For many functions, this is generally not a problem. The
next step is to "garble" the circuit. The main idea behind this garbling
process is that we want everyone to be able to evaluate the circuit, while
nobody should see what he actually evaluates. Therefore, we will try to
hide all the bits that "flow" through the circuit. For hiding the bits, we
could make use of a block cipher. However, we have to take care that the
circuit can still be evaluated! Therefore, we will also have to modify all
the gates in the circuit somehow, so that they are able to work with
"garbled" inputs. Now one could imagine that such a modification is a hard
task. Fortunately, there's a simple trick: All the gates in our circuit are
very small boolean functions (by small we mean that they don't have many
inputs). We can therefore replace every gate by its truth table. The truth
table simply maps the input bits to the respective output bits. For a
simple NAND gate, the truth table would look like this:
\a
b\ 1 0
+
1  0 1
0  1 1
Now that we have replaced every gate by its truth table, we will just have
to modify the truth tables, so that they reflect the fact that all the bit
values are garbled. The trick here is the following: Instead of the real
values of the input bits (1 or 0), we pick random cryptographic keys (say
128 bits long). We will then use those keys to encrypt the values in the
truth table. Instead of the input values for the gate, we will then use the
random keys (i.e. instead of 1 or 0, we just pick two random bitstrings per
wire).
As an example, consider a NAND gate again. We choose four keys ka0, ka1,
kb0 and kb1. Those are the keys for the respective input values of the gate
(i.e. a=0, a=1, b=0, b=1). Also, we pick an encryption function E and a
decryption function D. For simplicity, we assume that if a wrong key is
supplied to D, it will signal this (e.g. return an error code) instead of
providing a junk plaintext. We now perform the following transformation on
the truth table of our gate:
\a \ a
b\ 1 0 b \  1 0
+ > +
1  0 1 1  E_ka1(E_kb1(0)) E_ka0(E_kb1(1))
0  1 1 0  E_ka1(E_kb0(1)) E_ka0(E_kb0(1))
The elements in the truth table are double encrypted using the two keys
that belong to the values of a or b, respectively. When evaluating the
circuit, you only know the keys that correspond to the correct input values
(so for example you know ka0 and kb1 but no other key). By simply trying to
decrypt every value in the table, it is easy to find the according output
value (only one decryption will succeed).
The next question would then be: How to garble a whole circuit? It's not
much different. Assume that two gates are connected like this:
Out

++
 G1 
++
 
++ In_3
G2 
++
 \
In_1 In_2
We have inputs In_1, In_2, In_3 and one output value Out. G2's output is
connected to one of the input wires of G1. In the truth table of G2, we
therefore put the *key* that corresponds to the respective input value of
G1 (so instead of doubleencrypting G2's output value 1 or 0, we
doubleencrypt the respective key for one of G1's input pins). The gate G1
can be garbled as described above. The keys for the input wires In_1, In_2
and In_3 are assumed to be already known by the party evaluating the
circuit. G2 can now easily be evaluated and yields the missing key for
evaluating the gate G1. However, during the evaluation of the circuit, no
intermediate values (like the real output of G2) are disclosed to the
evaluating party.
Let's try that in practice:
Say Alice and Bob want to evaluate a function. The following protocol can
be used: A prepares a garbled circuit and hardcodes her input values into
the circuit. She sends the result to B. B now needs to retrieve the keys
for his input values from A. But beware of two limitations here:
1) B doesn't want to disclose his input values to A (obviously).
2) A doesn't want to tell B the keys for both input values, because
then B would be able to reverseengineer the circuit and to obtain
A's input values.
You've probably already seen the solution: B uses an oblivious transfer to
obtain the keys for his input values from A. For every bit b of his input
values, Bob will obliviously obtain the correct key k_b0 or k_b1 like this:
A  f, r0, r1 > B
A < z = f(k) XOR rb  B
A  k_b0 XOR k0, k_b1 XOR k1 > B
B is now able to evaluate the whole circuit. Depending on how A built the
circuit, the output truth tables could contain real or garbled values.
Using some simple tricks, we can even split the output between A and B (so
that A gets some part of the result and B gets another part). We'll detail
on that later. Now there are some problems when one party isn't honest.
Alice for instance could just prepare a malicious circuit that leaks
information about Bob's secret inputs. There are ways to prevent such
attacks ("cut and choose", zero knowledge proofs, etc), but we won't
provide the details here. A more detailed description (along with a
security proof) can be found in [3].
[ 3  OTR
For those who are not familiar with the OTR protocol, this section might
provide some help. OTR features a number of cryptographic properties,
including confidentiality, integrity, forward secrecy and deniability.
There are two major phases of the protocol: initial key exchange and
message exchange. The initial key exchange is based on the DiffieHellman
protocol. It is referred to as AKE (Authenticated Key Exchange). To defend
against active attackers, a publickey signature scheme (DSA in this
particular case) is used. The DSA master keys have to be exchanged
beforehand (OTR also offers to authenticate DSA keys using the SMP
protocol, but that's not interesting in our case).
All the cryptographic details are provided in [2]; it's not particularly
helpful to repeat them here. Keeping in mind that OTR's key exchange is
based on DiffieHellman combined with some symmetric crypto and a signature
scheme will suffice. After the keyexchange phase, each party will have a
number of symmetric keys for encryption and authentication. Those are
derived from the DiffieHellman master key by hashing it in various ways
(encryption and MAC key will obviously be different).
The messages are encrypted using AES in CTR mode, and each message is MACed
using the symmetric key material. That offers us confidentiality and
integrity. It's important to note that *only* symmetric keys are used for
the actual payload crypto. The DSA master keys are only used in the initial
keyexchange phase.
The next feature we're going to look at is forward secrecy. Forward secrecy
means that even if the (DSA) key of a participant is disclosed, past
conversations cannot be compromised. Forward secrecy in OTR is established
by the DiffieHellman protocol: after a conversation ends, both parties can
safely wipe the DiffieHellman key that they generated. There is no way for
an attacker (and not even for the conversation partners) to recompute that
key afterwards: to do that, one would either need to know the private
exponent of one party (which is of course also wiped from memory) or one
would need to derive the key from the public information exchanged between
both parties, which is infeasible (hopefully; that's what DiffieHellman
relies on in the first place).
Having understood how OTR provides forward secrecy, we can move on to
deniability. During the conversation, both parties can be sure that the
messages they receive are authentic and not modified by an attacker. It is
immediately clear that the message authenticity can not be verified without
the MAC key. If one of the conversation partners wants to convince a third
party that a message is authentic, this conversation partner implicitly
proofs his knowledge of the MAC key to the third party. But then again, the
third party can not be sure that the conversation partner didn't fake the
message (he can do this as he knows the MAC key). This is what we call weak
deniability [4]. Obviously, OTR offers weak deniability, as message
authentication is performed using only symmetric primitives. But OTR offers
even more: In every message, the sending party includes a new
DiffieHellman key exchange proposal. The proposal is also covered by the
MAC to rule out MITM attacks. So both parties frequently generate new key
material. And this lets us do a nice trick: as soon as they generate new
MAC keys they publicly disclose the old MAC keys. The old keys aren't used
anymore, so this is safe. But as the MAC keys are public, *everybody* could
create fake messages and compute proper MACs for those. This is what we
call strong deniability. OTR ships with a toolkit containing software for
actually forging messages. Depending on how much you already know (only the
MAC keys, MAC and encryption keys, MAC keys and some message plaintext),
you can use different tools to forge messages. If you know parts of the
plaintext and the MAC keys, you can exploit the fact that AES is used in
CTR mode to directly modify the known parts of the plaintext. If there is
no known plaintext, the otr_remac tool might helpful: Every message
contains a new DiffieHellman key exchange proposal in plaintext (but
covered by the MAC). Now you can simply replace that proposal by one that
you generated (e.g. using the otr_sesskeys tool) and compute a new MAC for
the packet. That allows you to easily fake the rest of the conversation:
You know your own private DiffieHellman key, so you can generate a
plausible set of MAC and encryption keys and just use that one. It will
look legitimate because the modified packet (containing your key exchange
data) still has a valid MAC.
[ 4  The Attack
The deniability of OTR stems from the fact that a third party does not know
whether a message has been sent during a conversation (and before the MAC
keys were disclosed) or was generated afterwards (when the MAC keys were
public). An obvious way to attack OTR's deniability would therefore be to
just monitor all the OTR traffic between A and B. If one party now decides
to disclose the MAC and encryption keys used for a particular message, the
authenticity of that message can be verified. And as the message has been
recorded during the conversation (i.e. before the MAC keys were public),
the recording party knows that it was not generated afterwards.
Let's look at a reallife example to shed some more light on what we're
doing. Imagine two hackers A and B who want to talk about serious stuff
(TM) using OTR. Both of them are slightly paranoid and don't trust each
other. In particular, Bob fears that Alice might backstab him. However, as
OTR is deniable, Bob assumes that even if Alice discloses the contents of
their conversation, he could still plausibly argue that Alice just made it
up to discredit him. So Bob ignores his paranoia and tells Alice his
secrets. Alice indeed plans to backstab Bob. Her first plan is simple: She
will just submit all the encrypted and authenticated messages to the
police. The police will later be able to state in court that Alice didn't
fake the messages after the conversation. She however quickly realizes that
this approach is inherently flawed: Bob could argue that Alice just sent
fake messages to the police (as Alice knows all the keys she could generate
such fake messages). Alice knows that this problem could be fixed if the
Police sniffed all the traffic themselves. But she also knows that this is
going to be difficult, so she comes up with a second idea: Why not use a
trusted third party?
Instead of submitting her messages to the police, she will just disclose
her private DSA key to her lawyer. Then, during her conversation with Bob,
she will use her lawyer as a proxy (i.e. she will let *him* do the
crypto). This way the lawyer can be sure that the conversation is
authentic. The judges will trust Alice's lawyer in court (at least they'll
trust him more than they trust Alice), so her problem is solved. Alice's
setup would look like this:
++
 Alice 
++
^
 NonOTR (maybe SSL)
v
++
 Lawyer  trust ++
 Speaks for  <>  Police / Court 
 Alice  ++
++
^
 OTR (Bob thinks he talks to Alice)
v
++
 Bob 
++
But now Alice realizes that she doesn't trust her lawyer enough to give him
her private DSA key: He could misuse it to impersonate her. Also, Alice
doubts that her lawyer's words would be trusted enough in court.
This example shows the problems that Alice has when she wants to break the
deniability of OTR. Her problems can be summarized as follows (we'll now
call the police the "observing party" and the lawyer will be called
"trusted third party"):
a) The observing party needs to sniff the network traffic. That implies
quite a privileged network position, as the traffic needs to be sniffed
passively, i.e. without the help of A or B. Because if A or B would
send their traffic to the observing party, A or B might just insert
bogus messages into their "sniff" stream and the observing party
couldn't be sure about the authenticity. Even worse, paranoid A and B
could use an anonymizing network, so that sniffing their traffic would
be a nontrivial task.
b) Also, the authenticity of a message can only be proven to the observing
party, but not to anybody else (as anybody else didn't sniff the traffic
and the observing party could just have cut some pieces or inserted new
ones).
Problem b) is not that much of importance. Just imagine the observing party
as the police, the judges or even Fnord. You should always assume that the
observing party is exactly the guys you wanna protect yourself against. If
you think that the police probably won't even get all the crypto stuff and
therefore just believe any plaintext logfile you show them, that's OK
(you're probably right). There might however be agencies that would not
really trust plaintext logs. And those agencies might be very interested in
the contents of some OTR conversations.
Problem a) remains open. Obviously, neither A nor B really trust the
observing party. If we had a trusted third party, we actually could mount
an attack against OTR's deniability, just as described in the lawyer
example above. Well, lucky us, neither A, nor B, nor the observing party
trust anybody and therefore, there will be no trusted third party ;)
Really? Interestingly, a trusted third party can be emulated using secure
function evaluation. This is what we didn't tell in the section above: You
can view a secure function evaluation scheme as a replacement for a trusted
third party. So instead of letting a third party compute some function
f(x, y), A and B can perform the computation on their own and still get the
same result: both players only receive f(x, y) but A doesn't see y and B
doesn't see x. So the main idea of our attack is: Emulate a trusted third
party using secure function evaluation. The setup that Alice now plans is
the following:
++
 Alice <+
++ 
^ 
  SFE Voodoo for emulating the lawyer
 
 v
 ++
  Police / Court 
 ++

 OTR

v
++
 Bob 
++
Our central idea is the following: A can send all the messages she received
from B to the observing party (the police in the figure above, but that
could really be everyone). The messages are still encrypted, so this is not
a problem. To make sure that the messages are not faked by A, we need to
make sure that A cannot produce valid MACs without the help of the
observing party. We therefore share the MAC key between A and the observing
party. Every time, A wants to validate or produce a MAC, she has to
cooperate with the observing party. Later on, A can reveal the encryption
key for any message to the observing party, which can be sure that the
message is authentic.
In the following section (4.1  4.3), we will provide a highlevel overview
of the attack. In section 4.4, you can find the actual protocol that Alice
and the observing party use.
[ 4.1  Sharing DiffieHellman Keys
OTR uses DiffieHellman to establish sortlived MAC and encryption keys.
The first part of our exercise is therefore to build a DiffieHellman
compatible 3party protocol that allows for sharing the generated key
between two parties. The following protocol between Alice (A), Bob (B)
and the observing party (O) works:
O  g^o > A  (g^o)^a > B
O < g^a  A
A < g^b  B
All computations are done modulo some prime p and g is a generator of a
sufficiently large subgroup of Z_p*, just as DiffieHellman mandates.
B will now compute g^oab as key. However, neither A nor O can reproduce
that key. If A wanted to compute it, she would need to know O's secret
exponent o. Similar for O. We can therefore say that the key k is shared
between O and A, in the sense that A and O need to cooperate in order to
actually use it.
[ 4.2  Generating MAC and Encryption Keys
Now that we have established a shared DiffieHellman key, we need to
securely derive the MAC and encryption keys from it. Let's assume we have a
circuit C, which takes the shared DiffieHellman key k as input and returns
the corresponding MAC and encryption keys as output. This circuit follows
immediately from the OTR specification. Before we can evaluate the circuit,
we first need to compute k (which neither A nor O know at this time). So
the overall function that A and O want to compute is:
f(a, o) = C(((g^b)^a)^o mod p)
We can transform this function to a new circuit and evaluate it together
(i.e. A and O evaluate the circuit). After the evaluation, A could get the
encryption keys. But that's not a good idea, because the OTR spec mandates
that MAC_key = hash(encryption_key). If A knew the encryption key, she
could compute the according MAC key. Also, it would be bad if O would get
the MAC keys, because then O could impersonate A and B. Therefore, we'll
slightly modify the circuit, so that A may pick a random bit string, which
the circuit XORs to the MAC key and to the encryption key (assuming the
random string is long enough for both keys). The "blinded" MAC and
encryption keys are then provided to A and O, the bitmask remains in A's
memory. If they want to use one of the keys for something, they will
evaluate a circuit that first XORs both halves together and then does the
actual computation using the key. At no point in time, A or O actually
learn the MAC or the encryption key.
Now that we know how to generate all the symmetric key material, we are
able to perform the full initial key exchange phase of OTR.
[ 4.3  Sending and Receiving Messages
When A receives a message from B, she cannot immediately decrypt it because
she doesn't know the decryption key. Also, verifying and sending messages
needs O's cooperation.
1) Message Decryption
If A wants to decrypt one of B's messages, she cooperates with O. Both
parties will jointly evaluate a decryption circuit. The circuit will be
built in such a way that only Alice will learn the result (i.e. Alice
will again provide a random bitstring as input the the circuit, which is
XORed to the result).
2) Message Verification
If A wants to verify one of B's messages, she has to cooperate with O.
A and O will jointly evaluate some sort of HMAC circuit, in order to
find out whether a message is authentic or not. We can design the
message verification function in such a way that O will immediately
learn the encrypted message and the MAC verification result. This
enables A to afterwards reveal the encryption key for a particular
message, so that O will be convinced A didn't fake it.
3) Message Creation
When A wants to create a message, she encrypts it together with O, just
as described in 1). In order to compute a MAC for the message, A and O
again cooperate. As each message has to contain a new DiffieHellman
public key, A and O will jointly compute such a key using the scheme
outlined above.
[ 4.4  The Final Protocol
In this section we'll describe our final protocol. It offers the following
features:
* We have three parties: A, B and O. A and O cooperate to backstab B. B is
not able to deny any of his messages towards O.
* O will not learn any message plaintext, unless A explicitly tells O the
respective keys.
* O is not able to impersonate neither A nor B.
* No trust relation between A and O is required.
* A does not have to disclose a whole conversation to O; it is possible to
only disclose selected messages.
* B does not notice that A and O cooperate.
[ 4.4.1  Initial KeyExchange
This section describes OTR's authenticated keyexchange (AKE).
Bob starts the key exchange by picking a random r and x and sending
AES_r(g^x), HASH(g^x) to Alice. That's the regular OTR protocol. Alice
then does a DiffieHellman keyexchange with O as outlined in section
4.1. We assume that A and O communicate over a secure channel.
O A < AES_r(g^x), HASH(g^x)  B
O < g^a  A
O  g^o > A
Now Alice sends her DiffieHellman public key to Bob. Note that she doesn't
know the private exponent of the key: she knows only a and g^ao, but
neither ao nor o.
A  g^ao > B
Bob has already computed the common key k (which Alice can't do) and uses
it to derive encryption keys c and c' and MAC keys m1, m1', m2, m2' (see
the OTR specs [2] for details) by hashing k in various ways. Bob builds
the following messages:
M_B = MAC_m1(g^x, g^ao, pub_B, keyid_B)
X_B = pub_B, keyid_B, sig_B(M_B)
Where pub_B is Bob's public DSA key and keyid_B is an identifier for Bob's
DiffieHellman proposal g^x. sig_B is created using Bob's DSA key. Using
the already derived symmetric keys, he sends AES_c(X_B),MAC_m2(AES_c(X_B))
over to Alice.
A < r, AES_c(X_B),MAC_m2(AES_c(X_B))  B
Alice is now supposed to also derive all the symmetric keys
and to use them to decrypt and verify the stuff that Bob sent. But Alice
cannot do that, so she cooperates with O. O sends her a garbled circuit C1,
which will compute
C1(o, a, mask) = (c, c') XOR mask
Alice randomly chooses mask, so only she will learn c and c'. In a number
of oblivious transfers, Alice receives the keys for her input values from
O.
O  C1 > A\
\
O  > A \
O <  A 
O  > A  Compute c, c' using SFE. Only A
.  receives the values.
. 
. /
O < eval(C1)  A /
O  (c,c') XOR mask > A/
Now Alice is finally able to decrypt the stuff that Bob sent her. She does
so and gets X_B. Currently, she is not able to verify the MAC_m2() value
Bob sent  she'll do that later. First she sends sig_B(M_B) over to O.
O < sig_B(M_B)  A
In order to actually verify sig_B(M_B), A and O first need to compute M_B.
As described above, M_B = MAC_m1(g^x, g^ao, pub_B, keyid_B). In order to
compute that MAC, both parties again need to cooperate. O creates a circuit
C2, which computes:
C2(o, a, pub_B, keyid_B) = MAC_m1(g^x, g^ao, pub_B, keyid_B)
Alice again uses oblivious transfers to obtain the keys for her secret
input value a, evaluates the circuit and both parties obtain the result
M_B.
O  C2 > A\
\
O  > A \
O <  A 
O  > A  Compute M_B using SFE. A and O
.  receive the value.
. 
. /
O < eval(C2)  A /
O  M_B > A/
Now that both have computed M_B, they first check the signature sig_B(M_B),
just as the OTR protocol mandates. If A and O are convinced that sig_B(M_B)
is OK, they can verify the MAC_m2(...) that B sent earlier. Again, they
perform some SFE voodoo to do that. The observing party prepares a circuit
C3, which computes:
C3(o, a, AES_c(X_B)) = MAC_m2'(AES_c(X_B))
A again uses oblivious transfers to obtain the keys for her input values
and the result is shared between both parties.
O  C3 > A\
\
O  > A \
O <  A 
O  > A  Compute MAC_m2'(AES_c(X_B)) using
.  SFE. Both receive the result.
. 
. /
O < eval(C3)  A /
O  MAC > A/
Now A and O are convinced that the key exchange with B succeeded. But they
still need to convince B that everything is OK. In particular, OTR mandates
that A should compute
M_A = MAC_m1'(g^ao, g^x, pub_A, keyid_A)
X_A = pub_A, keyid_A, sig_A(M_A)
and then send AES_c'(X_A), MAC_m2'(AES_c'(X_A)) over to B. Computing the
AES part can be done by A, because A knows the key c'. But for computing
the MAC, A and O again need to cooperate. First, A sends AES_c'(X_A) over
to O. Then O prepares a circuit C4, which computes:
C4(o, a, AES_c'(X_A)) = MAC_m2'(AES_c'(X_A))
Using oblivious transfers, Alice obtains the keys for her inputs from O.
After evaluating the circuit, A and O obtain MAC_m2'(AES_c'(X_A)).
O < AES_c'(X_A)  A\
O  C4 > A \
\
O  > A 
O <  A  Compute MAC_m2'(AES_c'(X_A)). Both
O  > A  parties receive the value.
. 
. /
O < eval(C4)  A /
O  MAC_m2'(AES_c'(X_A)) > A/
That's it. A can now send all the required values to B.
 AES_c'(X_A), MAC_m2'(AES_c'(X_A)) > B
B verifies all the stuff (just like A did but without the SFE) and the
key exchange is done.
[ 4.4.2  Message Exchange
Once they have exchanged their initial key material, Alice and Bob can
exchange actual messages. Suppose, Alice wants to send a message to Bob;
we'll restrict ourselves to that scenario. Receiving messages works
similar.
Alice now does the following (from the OTR protocol spec [2]):
Picks the most recent of her own DiffieHellman encryption keys that Bob
has acknowledged receiving (by using it in a Data Message, or failing
that, in the AKE). Let key_A be that key, and let keyid_A be its serial
number.
If the above key is Alice's most recent key, she generates a new
DiffieHellman key (next_dh), to get the serial number keyid_A+1.
To do this, Alice again needs to cooperate with the observing party.
The steps are exactly the same as we have already seen in the initial
keyexchange:
O < g^a  A
O  g^o > A
Alice now uses g^ao as next_dh. When she computed next_dh, Alice
picks the most recent of Bob's DiffieHellman encryption keys that she has
received from him (either in a Data Message or in the AKE). Let key_B be
that key, and let keyid_B be its serial number.
Now Alice would actually need to use DiffieHellman to compute a fresh
shared key with Bob, which she can use to derive the encryption and MAC
key. But as she doesn't really know the private exponent (she knows g^ao,
a and g^a, but not ao), she again needs to cooperate with O. So here we go:
O prepares a circuit C1:
C1(o, a, mask) = (ek, mk) XOR mask
The circuit will compute both, ek and mk (the encryption and MAC keys),
blinded with some value chosen by Alice. The result will be supplied only
to the observing party. Alice will keep the value of mask. In a number of
oblivious transfers, Alice receives the keys for her input values from O.
O  C1 > A\
\
O  > A \
O <  A 
O  > A  Compute (ek, mk) XOR mask using SFE.
.  Only O receives the result.
. 
. /
O < eval(C1)  A /
Alice now picks a value ctr, so that (key_A, key_B, ctr) is unique. The
ctr value is needed, because AES is going to be used in counter mode to
encrypt Alice's payload. The next step for Alice is to encrypt her message.
As she doesn't know the encryption key, O prepares a circuit C2 for her:
C2(ek_o, ek_a, ctr, msg) = AESCTR_ek,ctr(msg)
The inputs ek_o and ek_a denote O's and A's knowledge about ek, which is
ek XOR mask in O's case and mask in A's case. The result of the circuit
will only be provided to A (i.e. A just doesn't send it over to O). In a
number of oblivious transfers, Alice receives the keys for her input
values from O.
O  C2 > A\
\
O  > A \
O <  A 
O  > A  Encrypt msg using SFE. Only A
.  receives the result.
. /
. /
Now Alice can compute:
T_A = (keyid_A, keyid_B, next_dh, ctr, AESCTR_ek,ctr(msg))
T_A already contains Alice's message, but she still needs to MAC it. This
is again done by A and O together. O prepares a circuit C3:
C3(mk_o, mk_a, T_A) = MAC_mk(T_A)
O  C3 > A\
\
O  > A \
O <  A  Compute MAC_mk(T_A). Both
O  > A  parties receive the value.
. 
. /
O < eval(C3)  A /
O  MAC_mk(T_A) > A/
Please be aware that Alice will keep T_A secret. Although T_A doesn't
contain any plaintext, Alice does not want to disclose it to the observing
party. If she did, then her own deniability would also be gone.
Also, the OTR protocol mandates that Alice should send her old MAC keys in
plaintext to Bob, so that they can be considered public. If A and O wanted
to, they could do that (by computing the old MAC key again and sharing the
result). But as long as Bob doesn't check what Alice sent, she can just
send garbage. Indeed, in its current version (libotr 3.2.0), the OTR
implementation doesn't check the disclosed MAC keys. Consider the excerpt
from proto.c, line 657:
 snip 
/* Just skip over the revealed MAC keys, which we don't need. They
* were published for deniability of transcripts. */
bufp += reveallen; lenp = reveallen;
 snap 
So Alice can safely send:
A T_A,MAC_mk(T_A),oldmackeys=foobar> B
[ 4.5  What's Left
We have seen that in a scenario where at least one party cooperates with
the attacker, deniability is nontrivial. Our construction can be extended
and adopted and we conjecture that it quite generally applies to deniable
messaging protocols.
Regarding performance: Yeah, we know that all the SFE voodoo can be quite
expensive. Especially modular exponentiation in circuits is not really
cheap. However, there are ways to optimize the basic scheme we have
outlined here. If you're interested in that, you might wanna read [5] as an
introduction. Also, refer to section 4.5.2, which outlines one particular
optimization of our DiffieHellmanscheme. Regarding network latency: When
looking at all the crypto protocols outlined in this article (especially at
oblivious transfers), you will notice that often multiple messages need to
be exchanged. If you need 3 messages for one oblivious transfer and you
want to perform 128 oblivious transfers (for some 128bit crypto key or
so), then you end up with 384 messages being exchanged. In terms of network
latency, that might be troublesome. However, there are two things that help
us: first, we can perform oblivious transfers in parallel (i.e. still
exchange three messages but every message now contains data for 128
oblivious transfers). We can also precompute many values and exchange them
before they are really needed (random values for instance).
[ 4.5.1  FAQ
Q: This is all bullshit! I could just share my private keys with the
police, and that would also kill deniability!
Yep. And the police would then be able to impersonate you. One of our
key points is that you don't need to trust the observing party, neither
need they to trust you.
A: But the observing party won't be able to prove anything in court!
Well, yes and no. In a constitutional state you'd need to actually prove
stuff in court. Unfortunately, such states are rare. But even if you
live in such a state, then the observing party could be the judge.
Q: But all the conversations that I had before my peer cooperated with the
observing party are deniable, right?
A: Yes, unless the observing party sniffed your traffic (if you used a
decent anonymizer, this is unlikely).
Q: Wait, the observing party so far only learned that *somebody* has sent
a message. But how do they know it was the person that I tell them it
was?
Good question. This knowledge is generated during the initial key
exchange of OTR. To be precise, the observing party and the backstabber
both learn the identity of the conversation peer when he signs his
keyexchange proposal with his DSA key. The observing party also sees
that and as they track all subsequent keyexchanges, they can build a
"chain of evidence".
Q: But doesn't [4] already kill the deniability of OTR?
A: Ha, even better question! At least it attacks the strong deniability of
OTR. However, our scheme also attacks the weak deniability. Furthermore,
the attacker in [4] has far more capabilities than in our model. In [4],
the attacker is able to arbitrarily read and modify network traffic. In
our model, the attacker can rely on the cooperation with one of the two
conversation partners.
Q: OK, I'm convinced. Is there any implementation?
A: You're welcome to build one ;) See section 4.5.2 for details.
[ 4.5.2  How to Implement?
If you want to implement the scheme outlined above, first of all, you need
some framework for secure function evaluation. There are a number of
implementations out there, for instance Fairplay [6] or TASTY [5]. Once you
got your SFE framework running, you need to implement all the functions
that need to be computed jointly. The DiffieHellman stuff is probably most
efficient when implemented using a homomorphic cryptosystem (such as
RSA or ElGamal maybe). Now you may ask: how does a multiplicatively
homomorphic scheme help us computing DH keys? Well. There's some nice
optimization, which basically reduces the modular exponentiation to a
modular multiplication:
Alice picks some random j and sends g^(ab+bj) over to the observing party.
The observing party sends g^o.
A < g^b  B
O < g^(ab+j)  A
O  g^o > A
Note that Bob cannot compute g^abo, because Alice's value is "blinded" with
j. Alice cannot do so neither; she doesn't know o. Bob however can compute
g^(abo+jo). Alice can compute g^jo and also g^jo, because she knows j. If
Alice would send g^jo to O, then O could compute
g^(abo+jo) * g^jo = g^abo
This is only one modular multiplication. So instead of doing a whole
modular exponentiation, the circuit that Alice and the observing party
jointly compute does roughly the following:
C(o, a) = derive_keys(o*a)
Where the function derive_keys() is the OTR key derivation function
(hashing the common key in different ways to generate symmetric key
material), O's input value will look like g^(abo+jo) and A's input value
will look like g^jo.
All the symmetric operations (hashes and block ciphers) should probably be
implemented as circuits, for instance using Fairplay. Both SFE schemes
(circuits and homomorphic crypto) can be combined using the TASTY approach.
[ 5  References
[1] http://wwwee.stanford.edu/~hellman/publications/24.pdf
[2] http://www.cypherpunks.ca/otr/Protocolv23.0.0.html
[3] http://eprint.iacr.org/2004/175.pdf
[4] http://www.jbonneau.com/OTR_analysis.pdf
[5] http://eprint.iacr.org/2010/365.pdf
[6] http://www.pinkas.net/PAPERS/MNPS.pdf
[7] http://tinyurl.com/84z7wpu
[ 6  Greetingz
First of all I have to give a big shout to bruhns, who developed this stuff
together with me! There's this one person, which I'd like to say thanks for
everything (and that's quite a lot). Unfortunately, i cannot name this
person here. 291646a6d004d800b1bc61ba945c9cb46422f8ac. Also a big thanks to
Phrack staff for reading through all this and supplying me with real
helpful feedback!
Greetingz go out to the following awesome people in no particular order:
ths, fabs, joern, nowin, trapflag, jenny, twice#11
[ EOF