\resetseq
\head{PRFs are MACs, \seq}
- If $F_K\colon \{0, 1\}^* \to \{0, 1\}^L$ is a $(t, q, \epsilon)$-secure
- PRF, then it's also a $(t', q_T, q_V, \epsilon')$-secure MAC, with $q = q_T
- + q_V + 1$, $t = t' + O(q)$, and $\epsilon' \le \epsilon + (q_V + 1)
- 2^{-L}$. The constant hidden by the $O(\cdot)$ is small and depends on the
- model of computation.
+ If $F_K\colon \{0, 1\}^* \to \{0, 1\}^L$ is a PRF then it's also a MAC.
+ Quantitatively:
+ \[ \InSec{suf-cma}(F; t, q_T, q_V) \le
+ \InSec{prf}(F; t + O(q), q_T + q_V) +
+ \frac(q_V + 1}{2^L}. \]
+ The constant hidden by the $O(\cdot)$ is small and depends on the model of
+ computation.
Suppose $A$ can break $F$ used as a MAC in time $t$ and with $q_T$ and
$q_V$ queries to its tagging and verification oracles respectively.
If $\InSec{uh}(H) \le \epsilon$ then we say that $H$ is
\emph{$\epsilon$-almost universal}. Note that the concept of
almost-universality is not quantified by running times.
+
+ Such families definitely exist: we don't need to make intractability
+ assumptions. We'll see some examples later.
If $H$ is $1/|{\ran H}|$-almost universal, then we say that $H$ is
\emph{universal}. Sometimes it's said that this is the best possible
\end{slide}
\begin{slide}
- \topic{the collision game}
- \head{Universal hashing, \seq: the collision game}
-
- Suppose that, instead of merely a pair $(x, y)$, our adversary was allowed
- to return a \emph{set} $Y$ of $q$ elements, and measure the probability
- that $H_K(x) = H_K(y)$ for some $x \ne y$ with $x, y \in Y$, and for $K
- \inr \keys H$.
-
- Let $\InSec{uh-set}(H; q)$ be maximum probability achievable for sets $Y$
- with $|Y| \le q$. Then
- \[ \InSec{uh-set}(H; q) \le \frac{q(q - 1)}{2} \cdot \InSec{uh}(H) .\]
+ \topic{domain and range extension of a PRF}
+ \head{Universal hashing, \seq: PRF domain extension}
+
+ Suppose that $F\colon \keys F \times \{0, 1\}^L \to \{0, 1\}^l$ is a PRF.
+ We'd like to have a PRF for a bigger domain $\{0, 1\}^n$.
+
+ If $H\colon \{0, 1\}^k \times \{0, 1\}^n \to \{0, 1\}^L$ is an
+ almost-universal hash function. Then we can define $F'$ by
+ \[ F'_{K, h}(x) = F_K(H_h(x)). \]
+ Now we can prove that
+ \[ \InSec{prf}(F'; t, q) \le
+ \InSec{prf}(F; t, q) +
+ \frac{q(q - 1)}{2} \InSec{ah}(H). \]
+ Immediately this tells us that $F'$ is a MAC for $n$-bit messages.
\end{slide}
\begin{proof}
- This is rather tedious. We use the dynamic view. Suppose $A$ returns $(x,
- Y)$ with $|Y| = q$, and succeeds with probability $\epsilon$. Consider
+ Suppose $A$ attacks $F'$. Consider the distinguisher $B$ which attacks
+ $F$:
\begin{program}
- Adversary $A'$: \+ \\
- $(x, Y) \gets A$; \\
- $y \getsr Y$; \\
- \RETURN $(x, Y \setminus \{y\})$;
+ Distinguisher $B^{F(\cdot)}$: \+ \\
+ $h \getsr \{0, 1\}^k$; \\
+ \RETURN $A^{F(H_h(\cdot))}()$;
\end{program}
- The worst that can happen is that $A'$ accidentally removes the one
- colliding element from $Y$. This occurs with probability $2/q$. So
- \[ \Succ{uh-set}{H}(A') \ge \frac{q - 2}{q} \Succ{uh-set}{H}(A). \]
- Rearranging and maximizing gives
- \[ \InSec{uh-set}(H; q) \le
- \frac{q}{q - 2} \cdot \InSec{uh-set}(H; q - 1). \]
- Note that $\InSec{uh-set}(H; 2) = \InSec{uh}(H)$ is our initial notion. A
- simple inductive argument completes the proof.
+ Suppose $F$ is an oracle for $F_K(\cdot)$. Then $A$ plays its standard
+ attack game and all is well. Conversely, suppose $F$ is a random
+ function. Then the only way $A$ can distinguish its oracle $F(H_h(\cdot))$
+ from a random function is if it issues two queries $x_i \ne x_j$ such that
+ $H_h(x_i) = H_h(x_j)$. But for each pair $(i, j)$ this happens with
+ probability at most $\InSec{ah}(H)$. The stated bound follows.
\end{proof}
-\begin{slide}
- \topic{a MAC}
- \head{Universal hashing, \seq: a MAC}
-
- Suppose that $H\colon \{0, 1\}^k \times \{0, 1\}^* \to \{0, 1\}^l$ is an
- almost universal hash function, and $F\colon \{0, 1\}^{k'} \times \{0,
- 1\}^l \to \{0, 1\}^L$ is a PRF\@. Define a MAC $\Xid{\mathcal{M}}{UH}^{H,
- F} = (\Xid{T}{UH}^{H, F}, \Xid{V}{UH}^{H, F})$ where:
- \begin{eqnarray*}[rl]
- \Xid{T}{UH}^{H, F}_{K, K'}(m) &= F_{K'}(H_K(m)) \\
- \Xid{V}{UH}^{H, F}_{K, K'}(m, \tau) &= \begin{cases}
- 1 & if $\tau = F_{K'}(H_K(m))$ \\
- 0 & otherwise
- \end{cases}.
- \end{eqnarray*}
- We have
- \begin{eqnarray*}[Ll]
- \InSec{suf-cma}(\Xid{\mathcal{M}}{UH}^{H, F}; t, q_T, q_V) \\
- & \le
- (q_V + 1) \biggl(\InSec{prf}(F; t, q_T + 1) + \frac{1}{2^L} +
- \frac{q_T(q_T - 1)}{2} \cdot \InSec{uh}(H)\biggr).
- \end{eqnarray*}
-\end{slide}
-
-\begin{proof}
- We shall prove the result for $q_V = 0$ and $q_T = q$, and appeal to the
- earlier result on verification oracles.
-
- Suppose $A$ attacks the scheme $\Xid{\mathcal{M}}{UH}^{H, F}$ in time $t$,
- issuing $q$ tagging queries. Consider a distinguisher $D$, constructed
- from a forger $A$:
- \begin{program}
- Distinguisher $D^{F(\cdot)}$: \+ \\
- $K \getsr \{0, 1\}^k$; \\
- $\Xid{T}{list} \gets \emptyset$; \\
- $(m, \tau) \gets A^{\id{tag}(K, \cdot)}$; \\
- \IF $m \notin \Xid{T}{list} \land \tau = F(H_K(m))$
- \THEN \RETURN $1$; \\
- \ELSE \RETURN $0$; \- \\[\smallskipamount]
- Oracle $\id{tag}(K, m)$: \+ \\
- $\Xid{T}{list} \gets \Xid{T}{list} \cup \{m\}$; \\
- \RETURN $F(H_K(m))$; \- \\[\smallskipamount]
- \end{program}
- Note that $A$ isn't provided with a verification oracle: that's because we
- didn't allow it any verification queries.
-
- We can see immediately that
- \[ \Pr[K \getsr \{0, 1\}^{k'} : D^{F_K(\cdot)} = 1] =
- \Succ{suf-cma}{\Xid{\mathcal{M}}{UH}^{H, F}}(A). \]%
-
- We must now find an upper bound for $\Pr[F \getsr \Func{l}{L} :
- D^{F(\cdot)}]$. Suppose that the adversary returns the pair $(m^*,
- \tau^*)$, and that its tagging oracle queries and their answers are $(m_i,
- \tau_i)$ for $0 \le i < q$. Consider the event $C$ that $H_K(m) =
- H_K(m')$ for some $m \ne m'$, with $m, m' \in \{m^*\} \cup \{\,m_i \mid 0
- \le i < q\,\}$.
-
- If $C$ doesn't occur, then $F$ has not been queried before at $H_K(m)$, but
- there's a $2^{-L}$ probability that the adversary guesses right anyway. If
- $C$ does occur, then we just assume that the adversary wins, even though it
- might not have guessed the right tag.
-
- By our result on the collision game, $\Pr[C] \le q \cdot \InSec{uh}(H)$.
- Then
- \[ \Succ{prf}{F}(D) \ge
- \Succ{suf-cma}{\Xid{\mathcal{M}}{UH}^{H, F}}(A) -
- \frac{1}{2^L} - \frac{q(q - 1)}{2} \cdot \InSec{uh}(H). \]%
- The result follows.
-\end{proof}
+\begin{exercise}
+ Show how to extend the \emph{range} of a PRF. Specifically, suppose you're
+ given a PRF $F\colon \keys{F} \times \{0, 1\}^L \to \{0, 1\}^t$, but want
+ an $l$-bit output. Show how to achieve this.
+\answer
+ There are two obvious techniques.
+ \begin{enumerate}
+ \item Suppose we have a PRG $G\colon \{0, 1\}^t \to \{0, 1\}^l$. Then we
+ can easily use $G \compose F$; it's easy to see how this is secure.
+ Indeed, we can use $F$ itself as a PRG, by defining $G(x) = F_x(0) \cat
+ F_x(1) \cat \cdots$.
+ \item Let $m = \lceil \log_2 (l/t) \rceil$. Use the construction from the
+ slide to build a PRF $F'$ with domain $\{0, 1\}^{L+m}$. Then we define
+ $F''_K(x) = F'_K(x \cat 0) \cat F'_K(x \cat 1) \cat \cdots \cat F'_K(x
+ \cat m - 1)$.
+ \end{enumerate}
+\end{exercise}
\begin{slide}
\topic{almost XOR-universality}
\end{slide}
\begin{slide}
+ \head{Almost XOR-universality, \seq: examples of AXU hash functions}
+
+ Let $\F = \gf{2^l}$ be a finite field. Given a message $m = (m_0, \ldots,
+ m_{n-1} \in \F^n$, we can hash it in several ways.
+ \begin{itemize}
+ \item Inner-product: choose $k = (k_0, \ldots, k_{n-1}) \in \F^n$.
+ \[ \Xid{H}{ip}_{k}(m) = m \cdot k = \sum_{0\le i<n} k_i m_i. \]
+ This hash is XOR-universal: $\InSec{axu}(\Xid{H}{ip}) = 1/2^l$.
+ \item Polynomial evaluation: choose $x \in \F$.
+ \[ \Xid{H}{pe}_{x}(m) = \sum_{0\le i<n} m_i x^{i+1}. \]
+ Here we have only $\InSec{axu}(\Xid{H}{pe}) = n/2^l$.
+ \end{itemize}
+\end{slide}
+
+\begin{proof}
+ \begin{itemize}
+ \item Suppose we have messages $m \ne m'$ and $\delta \in \F$. Then $m
+ \cdot k + m' \cdot k = \delta$. We must have some $m_j \ne m'_j$. It
+ follows that
+ \[ k_j = \biggl( \delta + \sum_{\begin{script}
+ 0\le i<n \\
+ i \ne j
+ \end{script}
+ k_i (m_i + m'_i)
+ \biggr) \biggm/ (m'_i - m_j) \]
+ But we choose $k$ uniformly at random, so the probability that this holds
+ is $2^{-l}$.
+ \item Again, we have distinct messages and some $\delta \in \F$. We have a
+ polynomial equation
+ \[ \delta + \sum_{0\le i <n} (m_i - m'_i) x^{i+1} = 0 \]
+ Since there is some $m_j \ne m'_j$, this has degree at most $n$, so
+ there are at most $n$ distinct roots $x \in \F$. But we choose $x$
+ uniformly at random, so $x$ is one of these roots with probability at
+ most $n/2^l$.
+ \end{itemize}
+\end{proof}
+
+\begin{exercise}
+ Suppose $H$ is an $\epsilon$-AXU hash function with $l$-bit output. Let
+ $G_h(x)$ be the first $t$ bits of $H_h(x)$. Show that $G$ is
+ $2^{l-t}\epsilon$-AXU.
+\answer
+ Let $m \ne m'$ be distinct messages and let $\delta \in \{0, 1\}^t$ be some
+ XOR difference. Then
+ \begin{eqnarray*}
+ \Pr[h \gets \keys{G} : G_h(m) \xor G_h(m') = \delta]
+ &= \sum_{\delta' \in \{0, 1\}^{l-t}}
+ \Pr[h \gets \keys{H} : H_h(m) \xor H_h(m') = \delta \cat \delta']
+ &\le \sum_{\delta' \in \{0, 1\}^{l-t}} \InSec{axu}(H)
+ &= 2^{l-t} \InSec{axu}(H).
+ \end{eqnarray*}
+\end{exercise}
+
+\begin{slide}
\topic{a better MAC}
\head{Almost XOR-universality, \seq: a better MAC}
~mdw / doc / ips / blobdiff