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- Definition of perfect secrecy
- The one-time pad encryption scheme
- Necessity of long keys for perfect secrecy
- Computational secrecy and the derandomized one-time pad.
- Public key encryption
- A taste of advanced topics

“Human ingenuity cannot concoct a cipher which human ingenuity cannot resolve.”, Edgar Allen Poe, 1841

“I hope my handwriting, etc. do not give the impression I am just a crank or circle-squarer…. The significance of this conjecture [that certain encryption schemes are exponentially secure against key recovery attacks] .. is that it is quite feasible to design ciphers that are effectively unbreakable.”, John Nash, letter to the NSA, 1955.

"“Perfect Secrecy” is defined by requiring of a system that after a cryptogram is intercepted by the enemy the a posteriori probabilities of this cryptogram representing various messages be identically the same as the a priori probabilities of the same messages before the interception. It is shown that perfect secrecy is possible but requires, if the number of messages is finite, the same number of possible keys.", Claude Shannon, 1945

“We stand today on the brink of a revolution in cryptography.”, Whitfeld Diffie and Martin Hellman, 1976

Cryptography - the art or science of “secret writing” - has been around for several millenia, and for almost all of that time Edgar Allan Poe’s quote above held true. Indeed, the history of cryptography is littered with the figurative corpses of cryptosystems believed secure and then broken, and sometimes with the actual corpses of those who have mistakenly placed their faith in these cryptosystems.

Yet, something changed in the last few decades, which is the
“revolution” alluded to (and to a large extent initiated by) Diffie and
Hellman’s 1976 paper quoted above. New cryptosystems have been found
that have not been broken despite being subjected to immense efforts
involving both human ingenuity and computational power on a scale that
completely dwarves the “code breakers” of Poe’s time. Even more
amazingly, these cryptosystem are not only seemingly unbreakable, but
they also achieve this under much harsher conditions. Not only do
today’s attackers have more computational power but they also have more
data to work with. In Poe’s age, an attacker would be lucky if they got
access to more than a few encryptions of known messages. These days
attackers might have massive amounts of data- terabytes or more - at
their disposal. In fact, with *public key* encryption, an attacker can
generate as many ciphertexts as they wish.

The key to this success has been a clearer understanding of both how to
*define* security for cryptographic tools and how to relate this
security to *concrete computational problems*. Cryptography is a vast
and continuously changing topic, but we will touch on some of these
issues in this chapter.

A great many cryptosystems have been devised and broken throughout the ages. Let us recount just one such story. In 1587, Mary the queen of Scots, and the heir to the throne of England, wanted to arrange the assassination of her cousin, queen Elisabeth I of England, so that she could ascend to the throne and finally escape the house arrest under which she had been for the last 18 years. As part of this complicated plot, she sent a coded letter to Sir Anthony Babington.

Mary used what’s known as a *substitution cipher* where each letter is
transformed into a different obscure symbol (see
maryscottletterfig). At a first look, such a letter might seem
rather inscrutable- a meaningless sequence of strange symbols. However,
after some thought, one might recognize that these symbols *repeat*
several times and moreover that different symbols repeat with different
frequencies. Now it doesn’t take a large leap of faith to assume that
perhaps each symbol corresponds to a different letter and the more
frequent symbols correspond to letters that occur in the alphabet with
higher frequency. From this observation, there is a short gap to
completely breaking the cipher, which was in fact done by queen
Elisabeth’s spies who used the decoded letters to learn of all the
co-conspirators and to convict queen Mary of treason, a crime for which
she was executed. Trusting in superficial security measures (such as
using “inscrutable” symbols) is a trap that users of cryptography have
been falling into again and again over the years. (As in many things,
this is the subject of a great XKCD cartoon, see
XKCDnavajofig.)

Many of the troubles that cryptosystem designers faced over history (and
still face!) can be attributed to not properly defining or understanding
what are the goals they want to achieve in the first place. Let us focus
on the setting of *private key encryption*.*public key encryption* invented, see publickeyencdef.*sender*
(traditionally called “Alice”) wants to send a message (known also as a
*plaintext*) \(x\in \{0,1\}^*\) to a *receiver* (traditionally called
“Bob”). They would like their message to be kept secret from an
*adversary* who listens in or “eavesdrops” on the communication channel
(and is traditionally called “Eve”).

Alice and Bob share a *secret key* \(k \in \{0,1\}^*\). Alice uses the key
\(k\) to “scramble” or *encrypt* the plaintext \(x\) into a *ciphertext*
\(y\), and Bob uses the key \(k\) to “unscramble” or *decrypt* the
ciphertext \(y\) back into the plaintext \(x\). This motivates the following
definition:

Let \(L:\N \rightarrow \N\) be some function. A pair of polynomial-time
computable functions \((E,D)\) mapping strings to strings is a *valid
private key encryption scheme* (or *encryption scheme* for short) with
plaintext length function \(L(\cdot)\) if for every \(k\in \{0,1\}^n\) and
\(x \in \{0,1\}^{L(n)}\), \[
D(k,E(k,x))=x \;. \label{eqvalidenc}
\] We also require that our encryption schemes are *ciphertext length
regular* in the sense that all ciphertexts corresponding to keys of the
same length are of the same length: there is some function
\(C:\N \rightarrow \N\) such that for every \(k\in \{0,1\}^n\) and
\(x\in \{0,1\}^{L(n)}\), \(|E(k,x)|=C(n)\).*length function*
of \((E,D)\). The “ciphtertext length regularity” condition is added
for technical convenience and is not at all important. You can
ignore it in a first reading.

We will often write the first input (i.e., the key) to the encryption and decryption as a subscript and so can write \eqref{eqvalidenc} also as \(D_k(E_k(x))=x\).

encryptiondef says nothing about the *security* of \(E\) and
\(D\), and even allows the trivial encryption scheme that ignores the key
altogether and sets \(E_k(x)=x\) for every \(x\). Defining security is not a
trivial matter.

You would appreciate the subtleties of defining security of encryption
if at this point you take a five minute break from reading, and try
(possibly with a partner) to brainstorm on how you would mathematically
define the notion that an encryption scheme is *secure*, in the sense
that it protects the secrecy of the plaintext \(x\).

Throughout history, many attacks on cryptosystems are rooted in the
cryptosystem designers’ reliance on “security through obscurity”—
trusting that the fact their *methods* are not known to their enemy will
protect them from being broken. This is a faulty assumption - if you
reuse a method again and again (even with a different key each time)
then eventually your adversaries will figure out what you are doing. And
if Alice and Bob meet frequently in a secure location to decide on a new
method, they might as well take the opportunity to exchange their
secrets. These considerations led Auguste Kerckhoffs in 1883 to state
the following principle:

A cryptosystem should be secure even if everything about the system, except the key, is public knowledge.The actual quote is “Il faut qu’il n’exige pas le secret, et qu’il puisse sans inconvénient tomber entre les mains de l’ennemi” loosely translated as “The system must not require secrecy and can be stolen by the enemy without causing trouble”. According to Steve Bellovin the NSA version is “assume that the first copy of any device we make is shipped to the Kremlin”.

Why is it OK to assume the key is secret and not the algorithm? Because
we can always choose a fresh key. But of course that won’t help us much
if our key is if we choose our key to be “1234” or “passw0rd!”. In fact,
if you use *any* deterministic algorithm to choose the key then
eventually your adversary will figure this out. Therefore for security
we must choose the key at *random* and can restate Kerckhoffs’s
principle as follows:

There is no secrecy without randomness

This is such a crucial point that is worth repeating:

There is no secrecy without randomness

At the heart of every cryptographic scheme there is a secret key, and the secret key is always chosen at random. A corollary of that is that to understand cryptography, you need to know probability theory.

If you think about encryption scheme security for a while, you might
come up with the following principle for defining security: *“An
encryption scheme is secure if it is not possible to recover the key \(k\)
from \(E_k(x)\)”*. However, a moment’s thought shows that the key is not
really what we’re trying to protect. After all, the whole point of an
encryption is to protect the confidentiality of the *plaintext* \(x\). So,
we can try to define that *“an encryption scheme is secure if it is not
possible to recover the plaintext \(x\) from \(E_k(x)\)”*. Yet it is not
clear what this means either. Suppose that an encryption scheme reveals
the first 10 bits of the plaintext \(x\). It might still not be possible
to recover \(x\) completely, but on an intuitive level, this seems like it
would be extremely unwise to use such an encryption scheme in practice.
Indeed, often even *partial information* about the plaintext is enough
for the adversary to achieve its goals.

The above thinking led Shannon in 1945 to formalize the notion of
*perfect secrecy*, which is that an encryption reveals absolutely
nothing about the message. There are several equivalent ways to define
it, but perhaps the cleanest one is the following:

A valid encryption scheme \((E,D)\) with length \(L(\cdot)\) is *perfectly
secret* if for every \(n\in \N\) and plaintexts \(x,x' \in \{0,1\}^{L(n)}\),
the following two distributions \(Y\) and \(Y'\) over \(\{0,1\}^*\) are
identical:

- \(Y\) is obtained by sampling \(k\sim \{0,1\}^n\) and outputting \(E_k(x)\).
- \(Y'\) is obtained by sampling \(k\sim \{0,1\}^n\) and outputting \(E_k(x')\).

This definition might take more than one reading to parse. Try to think of how this condition would correspond to your intuitive notion of “learning no information” about \(x\) from observing \(E_k(x)\), and to Shannon’s quote in the beginning of this chapter. In particular, suppose that you knew ahead of time that Alice sent either an encryption of \(x\) or an encryption of \(x'\). Would you learn anything new from observing the encryption of the message that Alice actually sent? It may help you to look at perfectsecfig.

To understand perfectsecrecy, suppose that Alice sends only
one of two possible messages: “attack” or “retreat”, which we denote by
\(x_0\) and \(x_1\) respectively, and that she sends each one of those
messages with probability \(1/2\). Let us put ourselves in the shoes of
*Eve*, the eavesdropping adversary. A priori we would have guessed that
Alice sent either \(x_0\) or \(x_1\) with probability \(1/2\). Now we observe
\(y=E_k(x_i)\) where \(k\) is a uniformly chosen key in \(\{0,1\}^n\). How
does this new information cause us to update our beliefs on whether
Alice sent the plaintext \(x_0\) or the plaintext \(x_1\)?

Before reading the next paragraph, you might want to try the analysis yourself. You may find it useful to look at the Wikipedia entry on Bayesian Inference or these MIT lecture notes.

Let us define \(p_0(y)\) to be the probability (taken over
\(k\sim \{0,1\}^n\)) that \(y=E_k(x_0)\) and similarly \(p_1(y)\) to be
\(\Pr_{k \sim \{0,1\}^n}[y=E_k(x_1)]\). Note that, since Alice chooses the
message to send at random, our a priori probability for observing \(y\) is
\(\tfrac{1}{2}p_0(y) + \tfrac{1}{2}p_1(y)\). However, as per
perfectsecrecy, the perfect secrecy condition guarantees that
\(p_0(y)=p_1(y)\)! Let us denote the number \(p_0(y)=p_1(y)\) by \(p\). By the
formula for conditional probability, the probability that Alice sent the
message \(x_0\) conditioned on our observation \(y\) is simply*Bayes’
rule* which, although a simple restatement of the formula for
conditional probability, is an extremely important and widely used
tool in statistics and data analysis.

Since the probability that \(i=0\) and \(y\) is the ciphertext \(E_k(0)\) is equal to \(\tfrac{1}{2}\cdot p_0(y)\), and the a priori probability of observing \(y\) is \(\tfrac{1}{2}p_0(y) + \tfrac{1}{2}p_1(y)\), we can rewrite \eqref{bayeseq} as \[ \Pr[i=0 | y=E_k(x_i)] = \frac{\tfrac{1}{2}p_0(y)}{\tfrac{1}{2}p_0(y)+\tfrac{1}{2}p_1(y)} = \frac{p}{p +p} = \frac{1}{2} \] using the fact that \(p_0(y)=p_1(y)=p\). This means that observing the ciphertext \(y\) did not help us at all! We still would not be able to guess whether Alice sent “attack” or “retreat” with better than 50/50 odds!

This example can be vastly generalized to show that perfect secrecy is
indeed “perfect” in the sense that observing a ciphertext gives Eve *no
additional information* about the plaintext beyond her a priori
knowledge.

*Perfect secrecy* is an extremely strong condition, and implies that an
eavesdropper does not learn *any* information from observing the
ciphertext. You might think that an encryption scheme satisfying such a
strong condition will be impossible, or at least extremely complicated,
to achieve. However it turns out we can in fact obtain perfectly secret
encryption scheme fairly easily. Such a scheme for two-bit messages is
illustrated in onetimepadtwofig

In fact, this can be generalized to any number of bits:

There is a perfectly secret valid encryption scheme \((E,D)\) with \(L(n)=n\).

Our scheme is the one-time pad also known as the “Vernam Cipher”, see onetimepadfig. The encryption is exceedingly simple: to encrypt a message \(x\in \{0,1\}^n\) with a key \(k \in \{0,1\}^n\) we simply output \(x \oplus k\) where \(\oplus\) is the bitwise XOR operation that outputs the string corresponding to XORing each coordinate of \(x\) and \(k\).

For two binary strings \(a\) and \(b\) of the same length \(n\), we define \(a \oplus b\) to be the string \(c \in \{0,1\}^n\) such that \(c_i = a_i + b_i \mod 2\) for every \(i\in [n]\). The encryption scheme \((E,D)\) is defined as follows: \(E_k(x) = x\oplus k\) and \(D_k(y)= y \oplus k\). By the associative law of addition (which works also modulo two), \(D_k(E_k(x))=(x\oplus k) \oplus k = x \oplus (k \oplus k) = x \oplus 0^n = x\), using the fact that for every bit \(\sigma \in \{0,1\}\), \(\sigma + \sigma \mod 2 = 0\) and \(\sigma + 0 = \sigma \mod 2\). Hence \((E,D)\) form a valid encryption.

To analyze the perfect secrecy property, we claim that for every \(x\in \{0,1\}^n\), the distribution \(Y_x=E_k(x)\) where \(k \sim \{0,1\}^n\) is simply the uniform distribution over \(\{0,1\}^n\), and hence in particular the distributions \(Y_{x}\) and \(Y_{x'}\) are identical for every \(x,x' \in \{0,1\}^n\). Indeed, for every particular \(y\in \{0,1\}^n\), the value \(y\) is output by \(Y_x\) if and only if \(y = x \oplus k\) which holds if and only if \(k= x \oplus y\). Since \(k\) is chosen uniformly at random in \(\{0,1\}^n\), the probability that \(k\) happens to equal \(k \oplus y\) is exactly \(2^{-n}\), which means that every string \(y\) is output by \(Y_x\) with probability \(2^{-n}\).

The argument above is quite simple but is worth reading again. To
understand why the one-time pad is perfectly secret, it is useful to
envision it as a bipartite graph as we’ve done in
onetimepadtwofig. (In fact the encryption scheme of
onetimepadtwofig is precisely the one-time pad for \(n=2\).) For
every \(n\), the one-time pad encryption scheme corresponds to a bipartite
graph with \(2^n\) vertices on the “left side” corresponding to the
plaintexts in \(\{0,1\}^n\) and \(2^n\) vertices on the “right side”
corresponding to the ciphertexts \(\{0,1\}^n\). For every \(x\in \{0,1\}^n\)
and \(k\in \{0,1\}^n\), we connect \(x\) to the vertex \(y=E_k(x)\) with an
edge that we label with \(k\). One can see that this is the complete
bipartite graph, where every vertex on the left is connected to *all*
vertices on the right. In particular this means that for every left
vertex \(x\), the distribution on the ciphertexts obtained by taking a
random \(k\in \{0,1\}^n\) and going to the neighbor of \(x\) on the edge
labeled \(k\) is the uniform distribution over \(\{0,1\}^n\). This ensures
the perfect secrecy condition.

So, does onetimepad give the final word on cryptography, and means that we can all communicate with perfect secrecy and live happily ever after? No it doesn’t. While the one-time pad is efficient, and gives perfect secrecy, it has one glaring disadvantage: to communicate \(n\) bits you need to store a key of length \(n\). In contrast, practically used cryptosystems such as AES-128 have a short key of \(128\) bits (i.e., \(16\) bytes) that can be used to protect terabytes or more of communication! Imagine that we all needed to use the one time pad. If that was the case, then if you had to communicate with \(m\) people, you would have to maintain (securely!) \(m\) huge files that are each as long as the length of the maximum total communication you expect with that person. Imagine that every time you opened an account with Amazon, Google, or any other service, they would need to send you in the mail (ideally with a secure courier) a DVD full of random numbers, and every time you suspected a virus, you’d need to ask all these services for a fresh DVD. This doesn’t sound so appealing.

This is not just a theoretical issue. The Soviets have used the one-time
pad for their confidential communication since before the 1940’s. In
fact, even before Shannon’s work, the U.S. intelligence already knew in
1941 that the one-time pad is in principle “unbreakable” (see page 32 in
the Venona
document). However,
it turned out that the hassle of manufacturing so many keys for all the
communication took its toll on the Soviets and they ended up reusing the
same keys for more than one message. They did try to use them for
completely different receivers in the (false) hope that this wouldn’t be
detected. The Venona
Project of the U.S. Army
was founded in February 1943 by Gene Grabeel (see
genegrabeelfig), a former home economics teacher from Madison
Heights, Virgnia and Lt. Leonard Zubko. In October 1943, they had their
breakthrough when it was discovered that the Russians were reusing their
keys.

Unfortunately it turns out that (as shown by Shannon) that such long
keys are *necessary* for perfect secrecy:

For every perfectly secret encryption scheme \((E,D)\) the length function \(L\) satisfies \(L(n) \leq n\).

The idea behind the proof is illustrated in longkeygraphfig. If the number of keys is smaller than the number of messages then the neighborhoods of all vertices in the corresponding graphs cannot be identical.

Let \(E,D\) be a valid encryption scheme with messages of length \(L\) and key of length \(n<L\). We will show that \((E,D)\) is not perfectly secret by providing two plaintexts \(x_0,x_1 \in \{0,1\}^L\) such that the distributions \(Y_{x_0}\) and \(Y_{x_1}\) are not identical, where \(Y_x\) is the distribution obtained by picking \(k \sim \{0,1\}^n\) and outputting \(E_k(x)\). We choose \(x_0 = 0^L\). Let \(S_0 \subseteq \{0,1\}^*\) be the set of all ciphertexts that have nonzero probability of being output in \(Y_{x_0}\). That is, \(S=\{ y \;|\; \exists_{k\in \{0,1\}^n} y=E_k(x_0) \}\). Since there are only \(2^n\) keys, we know that \(|S_0| \leq 2^n\).

We will show the following claim:

**Claim I:** There exists some \(x_1 \in \{0,1\}^L\) and \(k\in \{0,1\}^n\)
such that \(E_k(x_1) \not\in S_0\).

Claim I implies that the string \(E_k(x_1)\) has positive probability of
being output by \(Y_{x_1}\) and zero probability of being output by
\(Y_{x_0}\) and hence in particular \(Y_{x_0}\) and \(Y_{x_1}\) are not
identical. To prove Claim I, just choose a fixed \(k\in \{0,1\}^n\). By
the validity condition, the map \(x \mapsto E_k(x)\) is a one to one map
of \(\{0,1\}^L\) to \(\{0,1\}^*\) and hence in particular the *image* of
this map: the set \(I = \{ y \;|\; \exists_{x\in \{0,1\}^L} y=E_k(x) \}\)
has size at least (in fact exactly) \(2^L\). Since \(|S_0| = 2^n < 2^L\),
this means that \(|I|>|S_0|\) and so in particular there exists some
string \(y\) in \(I \setminus S_0\). But by the definition of \(I\) this means
that there is some \(x\in \{0,1\}^L\) such that \(E_k(x) \not\in S_0\) which
concludes the proof of Claim I and hence of longkeysthm.

To sum up the previous episodes, we now know that:

- It is possible to obtain a perfectly secret encryption scheme with key length the same as the plaintext.

and

- It is not possible to obtain such a scheme with key that is even a single bit shorter than the plaintext.

How does this mesh with the fact that, as we’ve already seen, people
routinely use cryptosystems with a 16 bytes key but many terabytes of
plaintext? The proof of longkeysthm does give in fact a way to
break all these cryptosystems, but an examination of this proof shows
that it only yields an algorithm with time *exponential in the length of
the key*. This motivates the following relaxation of perfect secrecy to
a condition known as *“computational secrecy”*. Intuitively, an
encryption scheme is computationally secret if no polynomial time
algorithm can break it. The formal definition is below:

Let \((E,D)\) be a valid encryption scheme where for keys of length \(n\),
the plaintexts are of length \(L(n)\) and the ciphertexts are of length
\(m(n)\). We say that \((E,D)\) is *computationally secret* if for every
polynomial \(p:\N \rightarrow \N\), and large enough \(n\), if \(P\) is an
\(m(n)\)-input and single output NAND program of at most \(p(n)\) lines, and
\(x_0,x_1 \in \{0,1\}^{L(n)}\) then \[
\left| \E_{k \sim \{0,1\}^n} [P(E_k(x_0))] - \E_{k \sim \{0,1\}^n} [P(E_k(x_1))] \right| < \tfrac{1}{p(n)} \label{eqindist}
\]

compsecdef requires a second or third read and some practice
to truly understand. One excellent exercise to make sure you follow it
is to see that if we allow \(P\) to be an *arbitrary* function mapping
\(\{0,1\}^{m(n)}\) to \(\{0,1\}\), and we replace the condition in
\eqref{eqindist} that the lefhand side is smaller than
\(\tfrac{1}{p(L(n))}\) with the condition that it is equal to \(0\) then we
get the perfect secrecy condition of perfectsecrecy. Indeed if
the distributions \(E_k(x_0)\) and \(E_k(x_1)\) are identical then applying
any function \(P\) to them we get the same expectation. On the other hand,
if the two distributions above give a different probability for some
element \(y^*\in \{0,1\}^{m(n)}\), then the function \(P(y)\) that outputs
\(1\) iff \(y=y^*\) will have a different expectation under the former
distribution than under the latter.

compsecdef raises two natural questions:

- Is it strong enough to ensure that a computationally secret encryption scheme protects the secrecy of messages that are encrypted with it?
- It is weak enough that, unlike perfect secrecy, it is possible to obtain a computationally secret encryption scheme where the key is much smaller than the message?

To the best of our knowledge, the answer to both questions is *Yes*.
Regarding the first question, it is not hard to show that if, for
example, Alice uses a computationally secret encryption algorithm to
encrypt either “attack” or “retreat” (each chosen with probability
\(1/2\)), then as long as she’s restricted to polynomial-time algorithms,
an adversary Eve will not be able to guess the message with probability
better than, say, \(0.51\), even after observing its encrypted form. (We
omit the proof, but it is an excellent exercise for you to work it out
on your own.)

To answer the second question we will show that under the same
assumption we used for derandomizing \(\mathbf{BPP}\), we can obtain a
computationally secret cryptosystem where the key is almost
*exponentially* smaller than the plaintext.

It turns out that if pseudorandom generators exist as in the optimal PRG conjecture, then there exists a computationally secret encryption scheme with keys that are much shorter than the plaintext. The construction below is known as a stream cipher, though perhaps a better name is the “derandomized one-time pad”. It is widely used in practice with keys on the order of a few tens or hundreds of bits protecting many terabytes or even petabytes of communication.

Suppose that the optimal PRG conjecture is true. Then for every constant \(a\in \N\) there is a computationally secret encryption scheme \((E,D)\) with plaintext length \(L(n)\) at least \(n^a\).

The proof is illustrated in derandonetimepadfig. We simply take the one-time pad on \(L\) bit plaintexts, but replace the key with \(G(k)\) where \(k\) is a string in \(\{0,1\}^n\) and \(G:\{0,1\}^n \rightarrow \{0,1\}^L\) is a pseudorandom generator.

Since an exponential function of the form \(2^{\delta n}\) grows faster than any polynomial of the form \(n^a\), under the optimal PRG conjecture we can obtain a polynomial-time computable \((2^{\delta n},2^{-\delta n})\) pseudorandom generator \(G:\{0,1\}^n \rightarrow \{0,1\}^L\) for \(L = n^a\). We now define our encryption scheme as follows: given key \(k\in \{0,1\}^n\) and plaintext \(x\in \{0,1\}^L\), the encryption \(E_k(x)\) is simply \(x \oplus G(k)\). To decrypt a string \(y \in \{0,1\}^m\) we output \(y \oplus G(k)\). This is a valid encryption since \(G\) is computable in polynomial time and \((x \oplus G(k)) \oplus G(k) = x \oplus (G(k) \oplus G(k))=x\) for every \(x\in \{0,1\}^L\).

Computational secrecy follows from the condition of a pseudorandom
generator. Suppose, towards a contradiction, that there is a polynomial
\(p\), NAND program \(Q\) of at most \(p(L)\) lines and
\(x,x' \in \{0,1\}^{L(n)}\) such that \[
\left| \E_{k \sim \{0,1\}^n}[ Q(E_k(x))] - \E_{k \sim \{0,1\}^n}[Q(E_k(x'))] \right| > \tfrac{1}{p(L)}
\] which by the definition of our encryption scheme means that \[
\left| \E_{k \sim \{0,1\}^n}[ Q(G(k) \oplus x)] - \E_{k \sim \{0,1\}^n}[Q(G(k) \oplus x')] \right| > \tfrac{1}{p(L)} \;. \label{eqprgsecone}
\] Now since (as we saw in the security analysis of the one-time pad),
the distribution \(r \oplus x\) and \(r \oplus x'\) are identical, where
\(r\sim \{0,1\}^L\), it follows that \[
\E_{r \sim \{0,1\}^L} [ Q(r \oplus x)] - \E_{r \sim \{0,1\}^L} [ Q(r \oplus x')] = 0 \;. \label{eqprgsectwo}
\] By plugging \eqref{eqprgsectwo} into \eqref{eqprgsecone} we
can derive that \[
\left| \E_{k \sim \{0,1\}^n}[ Q(G(k) \oplus x)] - \E_{r \sim \{0,1\}^L} [ Q(r \oplus x)] + \E_{r \sim \{0,1\}^L} [ Q(r \oplus x')] - \E_{k \sim \{0,1\}^n}[Q(G(k) \oplus x')] \right| > \tfrac{1}{p(L)} \;. \label{eqprgsethree}
\] (Please make sure that you can see why this is true.) Now we can use
the *triangle inequality* that \(|A+B| \leq |A|+|B|\) for every two
numbers \(A,B\), applying it for
\(A= \E_{k \sim \{0,1\}^n}[ Q(G(k) \oplus x)] - \E_{r \sim \{0,1\}^L} [ Q(r \oplus x)]\)
and
\(B= \E_{r \sim \{0,1\}^L} [ Q(r \oplus x')] - \E_{k \sim \{0,1\}^n}[Q(G(k) \oplus x')]\)
to derive \[
\left| \E_{k \sim \{0,1\}^n}[ Q(G(k) \oplus x)] - \E_{r \sim \{0,1\}^L} [ Q(r \oplus x)] \right| + \left| \E_{r \sim \{0,1\}^L} [ Q(r \oplus x')] - \E_{k \sim \{0,1\}^n}[Q(G(k) \oplus x')] \right| > \tfrac{1}{p(L)} \;. \label{eqprgsefour}
\] In particular, either the first term or the second term of the
lefthand-side of \eqref{eqprgsefour} must be at least
\(\tfrac{1}{2p(L)}\). Let us assume the first case holds (the second case
is analyzed in exactly the same way). Then we get that \[
\left| \E_{k \sim \{0,1\}^n}[ Q(G(k) \oplus x)] - \E_{r \sim \{0,1\}^L} [ Q(r \oplus x)] \right| > \tfrac{1}{2p(L)} \;. \label{distingprgeq}
\] But if we now define the NAND program \(P_x\) that on input
\(r\in \{0,1\}^L\) outputs \(Q(r \oplus x)\) then (since XOR of \(L\) bits can
be computed in \(O(L)\) lines), we get that \(P_x\) has \(p(L)+O(L)\) lines
and by \eqref{distingprgeq} it can distinguish between an input of
the form \(G(k)\) and an input of the form \(r \sim \{0,1\}^k\) with
advantage better than \(\tfrac{1}{2p(L)}\). Since a polynomial is
dominated by an exponential, if we make \(L\) large enough, this will
contradict the \((2^{\delta n},2^{-\delta n})\) security of the
pseudorandom generator \(G\).

The two most widely used forms of (private key) encryption schemes in
practice are *stream ciphers* and *block ciphers*. (To make things more
confusing, a block cipher is always used in some mode of
operation
and some of these modes effectively turn a block cipher into a stream
cipher.) A block cipher can be thought as a sort of a “random invertible
map” from \(\{0,1\}^n\) to \(\{0,1\}^n\), and can be used to construct a
pseudorandom generator and from it a stream cipher, or to encrypt data
directly using other modes of operations. There are a great many other
security notions and considerations for encryption schemes beyond
computational secrecy. Many of those involve handling scenarios such as
*chosen plaintext*, *man in the middle*, and *chosen ciphertext*
attacks, where the adversary is not just merely a passive eavesdropper
but can influence the communication in some way. While this chapter is
meant to give you some taste of the ideas behind cryptography, there is
much more to know before applying it correctly to obtain secure
applications, and a great many people have managed to get it wrong.

We’ve also mentioned before that an efficient algorithm for \(\mathbf{NP}\) could be used to break all cryptography. We now give an example of how this can be done:

Suppose that \(\mathbf{P}=\mathbf{NP}\). Then there is no computationally secret encryption scheme with \(L(n) > n\). Furthermore, for every valid encryption scheme \((E,D)\) with \(L(n) > n+100\) there is a polynomial \(p\) such that for every large enough \(n\) there exist \(x_0,x_1 \in \{0,1\}^{L(n)}\) and a \(p(n)\)-line NAND program \(EVE\) s.t. \[ \Pr_{i \sim \{0,1\}, k \sim \{0,1\}^n}[ EVE(E_k(x_i))=i ] \geq 0.99 \]

Note that the “furthermore” part is extremely strong. It means that if
the plaintext is even a little bit larger than the key, then we can
already break the scheme in a very strong way. That is, there will be a
pair of messages \(x_0\), \(x_1\) (think of \(x_0\) as “sell” and \(x_1\) as
“buy”) and an efficient strategy for Eve such that if Eve gets a
ciphertext \(y\) then she will be able to tell whether \(y\) is an
encryption of \(x_0\) or \(x_1\) with probability very close to \(1\).

The proof follows along the lines of longkeysthm but this time paying attention to the computational aspects. If \(\mathbf{P}=\mathbf{NP}\) then for every plaintext \(x\) and ciphertext \(y\), we can efficiently tell whether there exists \(k\in \{0,1\}^n\) such that \(E_k(x)=y\). So, to prove this result we need to show that if the plaintexts are long enough, there would exist a pair \(x_0,x_1\) such that the probability that a random encryption of \(x_1\) also is a valid encryption of \(x_0\) will be very small. The details of how to show this are below.

We focus on showing only the “furthermore” part since it is the more interesting and the other part follows by essentially the same proof. Suppose that \((E,D)\) is such an encryption, let \(n\) be large enough, and let \(x_0 = 0^{L(n)}\). For every \(x\in \{0,1\}^{L(n)}\) we define \(S_x\) to be the set of all valid encryption of \(x\). That is \(S_x = \{ y \;|\; \exists_{k\in \{0,1\}^n} y=E_k(x) \}\). As in the proof of longkeysthm, since there are \(2^n\) keys \(k\), \(|S_x| \leq 2^n\) for every \(x\in \{0,1\}^{L(n)}\). We denote by \(S_0\) the set \(S_{x_0}\). We define our algorithm \(EVE\) to output \(0\) on input \(y\in \{0,1\}^*\) if \(y\in S_0\) and to output \(1\) otherwise. This can be implemented in polynomial time if \(\mathbf{P}=\mathbf{NP}\), since the key \(k\) can serve the role of an efficiently verifiable solution. (Can you see why?) Clearly \(\Pr[ EVE(E_k(x_0))=0 ] =1\) and so in the case that \(EVE\) gets an encryption of \(x_0\) then she guesses correctly with probability \(1\). The remainder of the proof is devoted to showing that there exists \(x_1 \in \{0,1\}^{L(n)}\) such that \(\Pr[ EVE(E_k(x_1))=0 ] \leq 0.01\), which will conclude the proof by showing that \(EVE\) guesses wrongly with probability at most \(\tfrac{1}{2}0 + \tfrac{1}{2}0.01 < 0.01\).

Consider now the following probabilistic experiment (which we define solely for the sake of analysis). We consider the sample space of choosing \(x\) uniformly in \(\{0,1\}^{L(n)}\) and define the random variable \(Z_k(x)\) to equal \(1\) if and only if \(E_k(x)\in S_0\). For every \(k\), the map \(x \mapsto E_k(x)\) is one-to-one, which means that the probability that \(Z_k=1\) is equal to the probability that \(x \in E_k^{-1}(S_0)\) which is \(\tfrac{|S_0|}{2^{L(n)}}\). So by the linearity of expectation \(\E[\sum_{k \in \{0,1\}^n} Z_k] \leq \tfrac{2^n|S_0|}{2^{L(n)}} \leq \tfrac{2^{2n}}{2^{L(n)}}\).

We will now use the following extremely simple but useful fact known as
the *averaging principle* (see also averagingprinciplerem):
for every random variable \(Z\), if \(\E[Z]=\mu\), then with positive
probability \(Z \leq \mu\). (Indeed, if \(Z>\mu\) with probability one, then
the expected value of \(Z\) will have to be larger than \(\mu\), just like
you can’t have a class in which all students got A or A- and yet the
overall average is B+.) In our case it means that with positive
probability \(\sum_{k\in \{0,1\}^n} Z_k \leq \tfrac{2^{2n}}{2^{L(n)}}\).
In other words, there exists some \(x_1 \in \{0,1\}^{L(n)}\) such that
\(\sum_{k\in \{0,1\}^n} Z_k(x_1) \leq \tfrac{2^{2n}}{2^{L(n)}}\). Yet this
means that if we choose a random \(k \sim \{0,1\}^n\), then the
probability that \(E_k(x_1) \in S_0\) is at most
\(\tfrac{1}{2^n} \cdot \tfrac{2^{2n}}{2^{L(n)}} = 2^{n-L(n)}\). So, in
particular if we have an algorithm \(EVE\) that outputs \(0\) if \(x\in S_0\)
and outputs \(1\) otherwise, then \(\Pr[ EVE(E_k(x_0))=0]=1\) and
\(\Pr[EVE(E_k(x_1))=0] \leq 2^{n-L(n)}\) which will be smaller than
\(2^{-10} < 0.01\) if \(L(n) \geq n+10\).

In retrospect breakingcryptowithnp is perhaps not surprising.
After all, as we’ve mentioned before it is known that the Optimal PRG
conjecture (which is the basis for the derandomized one-time pad
encryption) is *false* if \(\mathbf{P}=\mathbf{NP}\) (and in fact even if
\(\mathbf{NP}\subseteq \mathbf{BPP}\) or even
\(\mathbf{NP} \subseteq \mathbf{P_{/poly}}\)).

People have been dreaming about heavier than air flight since at least the days of Leonardo Da Vinci (not to mention Icarus from the greek mythology). Jules Verne wrote with rather insightful details about going to the moon in 1865. But, as far as I know, in all the thousands of years people have been using secret writing, until about 50 years ago no one has considered the possibility of communicating securely without first exchanging a shared secret key.

Yet in the late 1960’s and early 1970’s, several people started to
question this “common wisdom”. Perhaps the most surprising of these
visionaries was an undergraduate student at Berkeley named Ralph Merkle.
In the fall of 1974 Merkle wrote in a project
proposal for his computer security course
that while “it might seem intuitively obvious that if two people have
never had the opportunity to prearrange an encryption method, then they
will be unable to communicate securely over an insecure channel… I
believe it is false”. The project proposal was rejected by his professor
as “not good enough”. Merkle later submitted a paper to the
communication of the ACM where he apologized for the lack of references
since he was unable to find any mention of the problem in the scientific
literature, and the only source where he saw the problem even *raised*
was in a science fiction story. The paper was rejected with the comment
that “Experience shows that it is extremely dangerous to transmit key
information in the clear.” Merkle showed that one can design a protocol
where Alice and Bob can use \(T\) invocations of a hash function to
exchange a key, but an adversary (in the random oracle model, though he
of course didn’t use this name) would need roughly \(T^2\) invocations to
break it. He conjectured that it may be possible to obtain such
protocols where breaking is *exponentially harder* than using them, but
could not think of any concrete way to doing so.

We only found out much later that in the late 1960’s, a few years before Merkle, James Ellis of the British Intelligence agency GCHQ was having similar thoughts. His curiosity was spurred by an old World-War II manuscript from Bell labs that suggested the following way that two people could communicate securely over a phone line. Alice would inject noise to the line, Bob would relay his messages, and then Alice would subtract the noise to get the signal. The idea is that an adversary over the line sees only the sum of Alice’s and Bob’s signals, and doesn’t know what came from what. This got James Ellis thinking whether it would be possible to achieve something like that digitally. As Ellis later recollected, in 1970 he realized that in principle this should be possible, since he could think of an hypothetical black box \(B\) that on input a “handle” \(\alpha\) and plaintext \(x\) would give a “ciphertext” \(y\) and that there would be a secret key \(\beta\) corresponding to \(\alpha\), such that feeding \(\beta\) and \(y\) to the box would recover \(x\). However, Ellis had no idea how to actually instantiate this box. He and others kept giving this question as a puzzle to bright new recruits until one of them, Clifford Cocks, came up in 1973 with a candidate solution loosely based on the factoring problem; in 1974 another GCHQ recruit, Malcolm Williamson, came up with a solution using modular exponentiation.

But among all those thinking of public key cryptography, probably the
people who saw the furthest were two researchers at Stanford, Whit
Diffie and Martin Hellman. They realized that with the advent of
electronic communication, cryptography would find new applications
beyond the military domain of spies and submarines, and they understood
that in this new world of many users and point to point communication,
cryptography will need to scale up. Diffie and Hellman envisioned an
object which we now call “trapdoor permutation” though they called “one
way trapdoor function” or sometimes simply “public key encryption”.
Though they didn’t have full formal definitions, their idea was that
this is an injective function that is easy (e.g., polynomial-time) to
*compute* but hard (e.g., exponential-time) to *invert*. However, there
is a certain *trapdoor*, knowledge of which would allow polynomial time
inversion. Diffie and Hellman argued that using such a trapdoor
function, it would be possible for Alice and Bob to communicate securely
*without ever having exchanged a secret key*. But they didn’t stop
there. They realized that protecting the *integrity* of communication is
no less important than protecting its *secrecy*. Thus they imagined that
Alice could “run encryption in reverse” in order to certify or *sign*
messages.

At the point, Diffie and Hellman were in a position not unlike
physicists who predicted that a certain particle should exist but
without any experimental verification. Luckily they met Ralph
Merkle, and his ideas about a
probabilistic *key exchange protocol*, together with a suggestion from
their Stanford colleague John
Gill, inspired them to come up
with what today is known as the Diffie Hellman Key
Exchange
(which unbeknownst to them was found two years earlier at GCHQ by
Malcolm Williamson). They published their paper “New Directions in
Cryptography”
in 1976, and it is considered to have brought about the birth of modern
cryptography.

The Diffie-Hellman Key Exchange is still widely used today for secure
communication. However, it still felt short of providing Diffie and
Hellman’s elusive trapdoor function. This was done the next year by
Rivest, Shamir and Adleman who came up with the RSA trapdoor function,
which through the framework of Diffie and Hellman yielded not just
encryption but also signatures.

A *public key encryption* consists of a triple of algorithms:

- The
*key generation algorithm*, which we denote by \(KeyGen\) or \(KG\) for short, is a randomized algorithm that outputs a pair of strings \((e,d)\) where \(e\) is known as the*public*(or*encryption*) key, and \(d\) is known as the*private*(or*decryption*) key. The key generation algorithm gets as input \(1^n\) (i.e., a string of ones of length \(n\)). We refer to \(n\) as the*security parameter*of the scheme. The bigger we make \(n\), the more secure the encryption will be, but also the less efficient it will be. - The
*encryption algorithm*, which we denote by \(E\), takes the encryption key \(e\) and a plaintext \(x\), and outputs the ciphertext \(y=E_e(x)\). - The
*decryption algorithm*, which we denote by \(D\), takes the decryption key \(d\) and a ciphertext \(y\), and outputs the plaintext \(x=D_d(y)\).

We now make this a formal definition:

A *computationally secret public key encryption* with plaintext length
\(L:\N \rightarrow \N\) is a triple of randomized polynomial-time
algorithms \((KG,E,D)\) that satisfy the following conditions:

- For every \(n\), if \((e,d)\) is output by \(KG(1^n)\) with positive probability, and \(x\in \{0,1\}^{L(n)}\), then \(D_d(E_e(x))=x\) with probability one.
- For every polynomial \(p\), and sufficiently large \(n\), if \(P\) is a NAND program of at most \(p(n)\) lines then for every \(x,x'\in \{0,1\}^{L(n)}\), \(\left| \E[ P(e,E_e(x))] - \E[P(e,E_e(x'))] \right| < 1/p(n)\), where this probability is taken over the coins of \(KG\) and \(E\).

Note that we allowed \(E\) and \(D\) to be *randomized* as well. In fact, it
turns out that it is *necessary* for \(E\) to be randomized to obtain
computational secrecy. It also turns out that, unlike the private key
case, obtaining public key encryption for a *single bit of plaintext* is
sufficient for obtaining public key encryption for arbitrarily long
messages. In particular this means that we cannot obtain a perfectly
secret public key encryption scheme.

We will not give full constructions for public key encryption schemes in this chapter, but will mention some of the ideas that underlie the most widely used schemes today. These generally belong to one of two families:

*Group theoretic constructions*based on problems such as*integer factoring*, and the*discrete logarithm*over finite fields or elliptic curves.*Lattice/coding based constructions*based on problems such as the*closest vector in a lattice*or*bounded distance decoding*.

Group theory based encryptions are more widely implemented, but the
lattice/coding schemes are recently on the rise, particularly because
encryption schemes based on the former can be broken by *quantum
computers*, which we’ll discuss later in this course.

As just one example of how public key encryption schemes are constructed, let us now describe the Diffie-Hellman key exchange. We describe the Diffie-Hellman protocol in a somewhat of an informal level, without presenting a full security analysis.

The computational problem underlying the Diffie Hellman protocol is the
*discrete logarithm problem*. Let’s suppose that \(g\) is some integer. We
can compute the map \(x \mapsto g^x\) and also its *inverse*
\(y \mapsto \log_g y\).*binary search*: start with
some interval \([x_{min},x_{max}]\) that is guaranteed to contain
\(\log_g y\). We can then test whether the inteveral’s midpoint
\(x_{mid}\) satisfies \(g^{x_{mid}} > y\), and based on that halve the
size of the interval.*modular
arithmetic* and work modulo some prime number \(p\). If \(p\) has \(n\) binary
digits, \(g\) is in \([p]\), it is known how to compute the map
\(x \mapsto g^x \mod p\) in time polynomial in \(n\). (This is not trivial,
and is a great exercise for you to work this out.*discrete logarithm*). In fact, there is no known polynomial-time
algorithm for computing the map \((g,x,p) \mapsto \log_g x \mod p\), where
we define \(\log_g x \mod p\) as the number \(a \in [p]\) such that
\(g^a = x \mod p\).

The Diffie-Hellman protocol for Bob to send a message to Alice is as follows:

**Alice:**Chooses \(p\) to be a random \(n\) bit long prime (which can be done by choosing random numbers and running a primality testing algorithm on them), and \(g\) and \(a\) at random in \([p]\). She sends to Bob the triple \((p,g,g^a \mod p)\).**Bob:**Given the triple \((p,g,h)\), Bob sends a message \(x \in \{0,1\}^L\) to Alice by choosing \(b\) at random in \([p]\), and sending to Alice the pair \((g^b \mod p, rep(h^b \mod p) \oplus x)\) where \(rep:[p] \rightarrow \{0,1\}^*\) is some “representation function” that maps \([p]\) to \(\{0,1\}^L\).The function \(rep\) does not need to be one-to-one and you can think of \(rep(z)\) as simply outputting \(L\) of the bits of \(z\) in the natural binary representation. The function \(rep\) does need to satisfy certain technical conditions which we omit in this description. **Alice:**Given \(g',z\), Alice recovers \(x\) by outputting \(rep(g'^a \mod p) \oplus z\).

The correctness of the protocol follows from the simple fact that \((g^a)^b = (g^b)^a\) for every \(g,a,b\) and this still holds if we work modulo \(p\). Its security relies on the computational assumption that computing this map is hard, even in a certain “average case” sense (this computational assumption is known as the Decisional Diffie Hellman assumption). The Diffie-Hellman key exchange protocol can be thought of as a public key encryption where the Alice’s first message is the public key, and Bob’s message is the encryption.

One can think of the Diffie-Hellman protocol as being based on a
“trapdoor pseudorandom generator” whereas the triple \(g^a,g^{b},g^{ab}\)
looks “random” to someone that doesn’t know \(a\), but someone that does
know \(a\) can see that raising the second element to the \(a\)-th power
yields the third element. The Diffie-Hellman protocol can be described
abstractly in the context of any finite Abelian
group for which we can
efficiently compute the group operation. It has been implemented on
other groups than numbers modulo \(p\), and in particular Elliptic Curve
Cryptography
(ECC) is
obtained by basing the Diffie Hellman on elliptic curve groups which
gives some practical advantages.*quantum computer*. We will discuss
quantum computing later in this course.

There is a great deal to cryptography beyond just encryption schemes,
and beyond the notion of a passive adversary. A central objective is
*integrity* or *authentication*: protecting communications from being
modified by an adversary. Integerity is often more fundamental than
secrecy: whether it is a software update or viewing the news, you might
often not care about the communication being secret as much as that it
indeed came from its claimed source. *Digital signature schemes* are the
analog of public key encryption for authentication, and are widely used
(through the idea of
certificates) to
provide a foundations of trust in the digital world.

Similarly, even for encryption, we often need to ensure security against
*active attacks*, and so notions such as non-malleability and adaptive
chosen
ciphertext
security have been proposed. An encryption scheme is only as secure as
the secret key, and mechanisms to make sure the key is generated
properly, and is protected against refresh or even compromise (i.e.,
forward secrecy) have
been studied as well. Hopefully this chapter provides you with some
appreciation for cryptography as an intellectual field, but does not
imbue you with a false self of confidence in implementing it.

Beyond encryption and signature schemes, cryptographers have managed to obtain objects that truly seem paradoxical and “magical”. We briefly discuss some of these objects. We do not give any details, but hopefully this will spark your curiosity to find out more.

On October 31, 1903, the mathematician Frank Nelson Cole, gave an hourlong lecture to a meeting of the American Mathematical Society where he did not speak a single word. Rather, he calculated on the board the value \(2^{67}-1\) which is equal to \(147,573,952,589,676,412,927\), and then showed that this number is equal to \(193,707,721 \times 761,838,257,287\). Cole’s proof showed that \(2^{67}-1\) is not a prime, but it also revealed additional information, namely its actual factors. This is often the case with proofs: they teach us more than just the validity of the statements.

In *Zero Knowledge Proofs* we try to achieve the opposite effect. We
want a proof for a statement \(X\) where we can *rigorously show* that the
proofs reveals *absolutely no additional information about \(X\)* beyond
the fact that it is true. This turns out to be an extremely useful
object for a variety of tasks including authentication, secure
protocols, voting, anonymity in
cryptocurrencies, and more.
Constructing these objects relies on the theory of \(\mathbf{NP}\)
completeness. Thus this theory that originally was designed to give a
*negative result* (show that some problems are hard) ended up yielding
*positive applications*, enabling us to achieve tasks that were not
possible otherwise.

Suppose that we are given a bit-by-bit encryption of a string
\(E_k(x_0),\ldots,E_k(x_{n-1})\). By design, these ciphertexts are
supposed to be “completely unscrutable” and we should not be able to
extract any information about \(x_i\)’s from it. However, already in 1978,
Rivest, Adleman and Dertouzos observed that this does not imply that we
could not *manipulate* these encryptions. For example, it turns out the
security of an encryption scheme does not immediately rule out the
ability to take a pair of encryptions \(E_k(a)\) and \(E_k(b)\) and compute
from them \(E_k(a NAND b)\) *without knowing the secret key \(k\)*. But do
there exist encryption schemes that allow such manipulations? And if so,
is this a bug or a feature?

Rivest et al already showed that such encryption schemes could be
*immensely* useful, and their utility has only grown in the age of cloud
computing. After all, if we can compute NAND then we can use this to run
any algorithm \(P\) on the encrypted data, and map
\(E_k(x_0),\ldots,E_k(x_{n-1})\) to \(E_k(P(x_0,\ldots,x_{n-1}))\). For
example, a client could store their secret data \(x\) in encrypted form on
the cloud, and have the cloud provider perform all sorts of computation
on these data without ever revealing to the provider the private key,
and so without the provider *ever learning any information* about the
secret data.

The question of *existence* of such a scheme took much longer time to
resolve. Only in 2009 Craig Gentry gave the first construction of an
encryption scheme that allows to compute a universal basis of gates on
the data (known as a *Fully Homomorphic Encryption scheme* in crypto
parlance). Gentry’s scheme left much to be desired in terms of
efficiency, and improving upon it has been the focus of an intensive
research program that has already seen significant improvements.

Cryptography is about enabling mutually distrusting parties to achieve a
common goal. Perhaps the most general primitive achieving this objective
is secure multiparty
computation.
The idea in secure multiparty computation is that \(n\) parties interact
together to compute some function \(F(x_0,\ldots,x_{n-1})\) where \(x_i\) is
the private input of the \(i\)-th party. The crucial point is that there
is *no commonly trusted party or authority* and that nothing is revealed
about the secret data beyond the function’s output. One example is an
*electronic voting protocol* where only the total vote count is
revealed, with the privacy of the individual voters protected, but
without having to trust any authority to either count the votes
correctly or to keep information confidential. Another example is
implementing a second price (aka Vickrey)
auction where \(n-1\)
parties submit bids to an item owned by the \(n\)-th party, and the item
goes to the highest bidder but at the price of the *second highest bid*.
Using secure multiparty computation we can implement second price
auction in a way that will ensure the secrecy of the numerical values of
all bids (including even the top one) except the second highest one, and
the secrecy of the identity of all bidders (including even the second
highest bidder) except the top one. We emphasize that such a protocol
requires no trust even in the auctioneer itself, that will also not
learn any additional information. Secure multiparty computation can be
used even for computing *randomized* processes, with one example being
playing Poker over the net without having to trust any server for
correct shuffling of cards or not revealing the information.

Most of the exercises have been written in the summer of 2018 and haven’t yet been fully debugged. While I would prefer people do not post online solutions to the exercises, I would greatly appreciate if you let me know of any bugs. You can do so by posting a GitHub issue about the exercise, and optionally complement this with an email to me with more details about the attempted solution.

Much of this text is taken from my lecture notes on cryptography.

Shannon’s manuscript was written in 1945 but was classified, and a partial version was only published in 1949. Still it has revolutionized cryptography, and is the forerunner to much of what followed.

John Nash made seminal contributions in mathematics and game theory, and was awarded both the Abel Prize in mathematics and the Nobel Memorial Prize in Economic Sciences. However, he has struggled with mental illness throughout his life. His biography, A Beautiful Mind was made into a popular movie. It is natural to compare Nash’s 1955 letter to the NSA to Gödel’s letter to von Neumann we mentioned before. From the theoretical computer science point of view, the crucial difference is that while Nash informally talks about exponential vs polynomial computation time, he does not mention the word “Turing Machine” or other models of computation, and it is not clear if he is aware or not that his conjecture can be made mathematically precise (assuming a formalization of “sufficiently complex types of enciphering”).

The definition of computational secrecy we use is the notion of
*computational indistinguishability* (known to be equivalent to
*semantic security*) that was given by Goldwasser and Micali in 1982.

Although they used a different terminology, Diffie and Hellman already
made clear in their paper that their protocol can be used as a public
key encryption, with the first message being put in a “public file”. In
1985, ElGamal showed how to obtain a *signature scheme* based on the
Diffie Hellman ideas, and since he described the Diffie-Hellman
encryption scheme in the same paper, it is sometimes also known as
ElGamal encryption.

Zero-knowledge proofs were constructed by Goldwasser, Micali, and Rackoff in 1982, and their wide applicability was shown (using the theory of \(\mathbf{NP}\) completeness) by Goldreich, Micali, and Wigderson in 1986.

Two party and multiparty secure computation protocols were constructed (respectively) by Yao in 1982 and Goldreich, Micali, and Wigderson in 1987. The latter work gave a general transformation from security against passive adversaries to security against active adversaries using zero knowledge proofs.

Some topics related to this chapter that might be accessible to advanced students include: (to be completed)

Copyright 2018, Boaz Barak.

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