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\noindent {\large Essential Coding Theory \hfill
Madhu Sudan\\[-.1in]
6.896 \hfill \mbox{}\\[-.1in]
Due: Wednesday, October 30, 2002\hfill \mbox{}\\[.4in]}
{\LARGE \centering Problem Set 3 \\[.4in] \par}
\noindent {\bf Instructions:} See PS1.
\begin{enumerate}
\item {\sf Asymptotics of codes:} Given $\epsilon > 0$ express
the rate of the best family of binary codes of relative distance
$\frac12 - \epsilon$,
you can (a) construct, and (b) show the existence of. Express
the rate in big-Oh notation (i.e., $O(\epsilon^d)$ implies there
exist constants $c$ and $\epsilon_0$ such that for all $\epsilon
< \epsilon_0$, the rate of the code of relative distance
$\frac12-\epsilon$ is at least $c \epsilon^d$.)
How constructive are your codes in Part (a)?
{\sf
\begin{quote}
Obviously one should be doing Part (b) first. The Gilbert-Varshamov
bound gives a rate of $R = 1 - H(\frac12 - \epsilon)$.
Using $H(\frac12-\epsilon) = 1 - \Theta(\epsilon^2)$, we get
$R = \Theta(\epsilon^2)$ for a randomly chosen code.
As mentioned in lectures, this is shown to be optimal by the
Linear Programming bound, so can't be improved.
For constructive results, we have two options:
Option 1 - Concatenation of Reed-Solomon code with "greedily chosen
code": Here for an appropriate integer $\ell$, we pick an outer
RS code of relative distance $1 - \epsilon/2$ and block length
$2^{\ell}$ over an alphabet of size $q = 2^{\ell}$. The rate
of this code is $\epsilon/2$. We then
concatenate it with a greedily chosen binary code of relative
distance $\frac12 - \frac{\epsilon}2$, message length $\ell$ and
block length $\ell/\Theta(\epsilon^2)$. The concatenated code
has block length $n = q^2/\Theta(\epsilon^2)$, rate $\Theta(\epsilon^3)$
and relative distance $(1 - \epsilon/2)(\frac12 - \frac\epsilon2)
\geq \frac12 - \epsilon$. The construction time of this code
is dominated by the time taken to construct the inner code
which is $2^{\ell/\Theta(\epsilon^2)} = n^{\Theta(1/\epsilon^2)}$.
Option 2 - Concatenation of AG codes with Hadamard code:
In this case we pick an appropriately large constant $q$ so that
an outer AG code of relative distance $1 - 2\epsilon$ and rate $\epsilon$
of arbitrarily long block length $\ell$
can be found. Note that this requires $q = O(\epsilon^2)$ (since
an AG code needs $1 - R - \delta \geq \frac{1}{\sqrt{q}-1}$).
Now concatenate this code with an inner Hadamard code with
$q$ messages and block length roughly $q$. The rate of this
code is $\log q/q = \Theta(\epsilon^{-2} \log \frac1\epsilon)$.
Concatenating the two gives a code of block length $n = \ell
\epsilon^{-2}$,
rate $\Omega(\epsilon^{-3} \log\frac1\epsilon)$ and relative distance
$\frac12 - \epsilon$. So the code has slightly better rate (by a
logarithmic factor. Furthermore it can be constructed in time that
is a fixed polynomial in $n$ ($\Theta(n^2)$ as per latest results).
The literature has a third code with such rate and distance, due
to Alon, Bruck, Naor, Naor and Roth. We may encounter this code later.
\end{quote}
}
\item {\sf Variants of RS codes:}
The two parts of this question consider variants
of Reed-Solomon codes over $\F_q$, obtained by evaluations of
polynomials at $n$ distinct points $\alpha_1,\ldots,\alpha_n \in \F_q$.
The message will be speified by a sequence of coefficients
$c_0,\ldots,c_{k-1} \in \F_q$ and its encoding will be the
evaluation of a polynomial $p(x)$ at $\alpha_1,\ldots,\alpha_n$.
What will be different is the definition of $p(x)$ given
$c_0,\ldots,c_{k-1}$. Give exact bounds on the distance of the
resulting code. (Note, the distance may be a function of the
set $\{\alpha_1,\ldots.\alpha_n\}$.)
\begin{enumerate}
\item $p(x) = \sum_{i=0}^{k-1} c_i x^{i + \ell}$, where
$\ell$ is some non-negative integer.
\item $p(x) = \sum_{i=0}^{k-1} c_i x^{2i}$.
\end{enumerate}
\begin{quote}
{\sf
\begin{enumerate}
\item
For a message $c_0,\ldots,c_{k-1}$, let $C(x) = \sum_{i=0}^{k-1}
c_i x^i$ be the standard Reed-Solomon polynomial associate with
this message. Then in the first part, the encoding of this
message is the sequence $\langle \alpha_i^{\ell} C(\alpha_i)
\rangle_{i=1}^n$. If $\ell = 0$, this is just the Reed-Solomon code
and its distance equals $n-k+1$. Else consider a non-zero codeword of
minimum weight. For every $i$ s.t. $\alpha_i^{\ell} C(\alpha_i) = 0$,
either $\alpha_i = 0$ or $C(\alpha_i) = 0$. There are at most
$k-1$ locations for which the latter can hold (for a non-zero
message) and at most one coordinate for which the former holds.
Thus the code has distance at least $n-k$. Furthermore equality
holds iff there exists an $i$ s.t. $\alpha_i = 0$, else it distance
is $n-k+1$.
\item Let $C$ be as above. Then here the encoding is the evaluation
of $C$ at $\beta_1,\ldots,\beta_n$ where $\beta_i = \alpha_i^2$.
This is essentially just a RS code, except we don't know that the
$\beta_i$'s are distinct. Indeed every $\beta_i$ could potentially
appear twice in the sequence. Let $S$ be the set of $\beta_i$'s
that appear twice, and let $T$ be the set of $\beta_i$'s that appear
exactly once. Let $s = |S|$ and $t = |T|$ (so $2s+t = n$).
Then the minimum distance of the code is at most $n -
(k-1 + \min\{s,k-1\})$. (Consider $C$ which is zero on all of
$S$ and $k - 1 - s$ points of $T$ if $s < k-1$ or on $k-1$ points
of $S$ is $s \geq k-1$ - this achieves the bound.)
Now to get some bounds on $s$ for given $q$ and $n$. If $q$ is a
power of two, then $s = 0$, since the map $\alpha \mapsto \alpha^2$
is a bijection with the map $\beta \mapsto \beta^{q/2}$ inverting it.
For other fields there are at exactly $q/2$ distinct squares in the
field. So this implies $n - q/2 \leq s \leq n/2$ and any $s$
in this range is achievable.
\end{enumerate}
}
\end{quote}
\item {\sf Hadamard matrices:} Recall that an $n \times n$
matrix $H$ all
of whose entries are from $\{+1,-1\}$ is a Hadamard matrix if
$H \cdot H^T = n \cdot I$ where the matrix product is over the reals
and $I$ is the $n \times n$ identity matrix.
\begin{enumerate}
\item Show that if there is an $n \times n$ Hadamard matrix
then $n$ is either $1$ or $2$ or a multiple of $4$.
\begin{quote}
{\sf
Let $\vec a$, $\vec b$, $\vec c$ be three distinct rows of a
Hadamard matrix. (So we are assuming $n \geq 3$.)
For $i,j \in \{1,-1\}$, let $S_{i,j} = \{k | a_k = i\cdot b_k
\mbox{ and } a_k = j\cdot c_k\}$. Let
$\alpha = |S_{1,1}|$,
$\beta = |S_{1,-1}|$,
$\gamma = |S_{-1,1}|$,
and $\delta = |S_{-1,-1}|$.
Then $\alpha + \beta$ counts the number of coordinates where
$\vec a$ equals $\vec b$ and so $\alpha + \beta = n/2$.
Similarly $\alpha + \gamma$ counts the number of coordinates where
$\vec a$ equals $\vec c$ and so $\alpha + \gamma = n/2$.
Finally, $\alpha + \delta$ counts the number of coordinates where
$\vec b$ equals $\vec c$ and so $\alpha + \delta = n/2$.
Oh, and of course, $\alpha + \beta + \gamma + \delta = n$.
Solving the $4 \times 4$ linear system above, we get
$\alpha = \beta = \gamma = \delta = n/4$. Since each is an integer,
we have $n$ must be a multiple of $4$.
}
\end{quote}
\item Given an $n \times n$ Hadamard matrix $H_n$ and an
$m \times m$ Hadamard matrix $H_m$, construct an
$(nm) \times (nm)$ Hadamard matrix.
\begin{quote}
{\sf
Let $\F$ be any field (say rationals, for this problem).
For vectors $\vec a \in \F^n$ and $\vec b \F^m$,
let $\vec a \otimes \vec b \in \F^{nm}$
denote their outer product (aka tensor product), namely
the vector whose $ij$-th coordinate is $a_i \cdot b_j$.
Note that
if $\vec a$, $\vec b$ are $+1/-1$ vectors, then so is
$\vec a \otimes \vec b$.
Furthermore, if $\vec a, \vec c \in \F^n$ and
$\vec b, \vec d \in \F^m$,
$$\langle \vec a \otimes \vec b, \vec c \otimes \vec d
\rangle = \sum_{ij} a_i b_j c_i d_j = (\sum_i a_i c_i )(\sum_j b_j d_j)
= \langle \vec a, \vec c \rangle \cdot \langle \vec b, \vec d \rangle.$$
We show how to use tensor products to build a big Hadamard matrix
from two smaller ones.
Let $\vec{u_1}\ldots,\vec{u_n}$ be the rows of $H_n$
and
let $\vec{v_1}\ldots,\vec{v_m}$ be the rows of $H_m$.
By the condition $H_n H_n^T = nI$, we have $\langle
\vec{u_i},\vec{u_j}\rangle = 0$ if $i \ne j$. (Similarly
for the $\vec{v_i}$'s.)
Let $H_{nm}$ be the matrix whose rows are
$\vec{u_i} \otimes \vec{v_j}$ for all $i \in [n]$, $j \in [m]$.
As noted above, this is a $+1/-1$ matrix. Thus the diagonal
entries of $H_{nm} H^T_{nm}$ are all $nm$ as required. Now
consider the off-diagonal entry
$(H_{nm} H^T_{nm})_{(ij),(kl)} =
\langle \vec{u_i} \otimes \vec{v_j},
\vec{u_k} \otimes \vec{v_l} \rangle
=
\langle \vec{u_i},\vec{u_k} \rangle \cdot
\langle \vec{v_j},\vec{v_l} \rangle$.
Since at least one of the conditions $i \ne k$ or
$j \ne l$ holds, we have the above inner product is zero.
This proves the off-diagonal entries are zero as required.
}
\end{quote}
\item (Not to be turned in) Let $q$ be a prime power equivalent
to $3$ modulo $4$. Let $H = \{h_{ij}\}$ be the $(q+1) \times (q+1)$
matrix with $h_{ij} = 1$ if $i = 1$ or $j = 1$ or $i = j$,
and $h_{ij} = (j-i)^{(q-1)/2}$ otherwise. Verify that $H$ is
a Hadamard matrix. (The purpose of this exercise is point out that
Hadamard matrices of many size, and not just powers of $2$, exist.)
\begin{quote}
{\sf
Turns out there were some typoes in the above question so
it was not easy to ``verify''. Will post revised question + answer
shortly. (Some you nevertheless responded with ``verified''!)}
\end{quote}
\end{enumerate}
\item Let $C$ be an infinite family of binary codes obtained
by concatenation of two infinite families of codes $C_1$ and $C_2$.
(The $i$th code of $C$ is obtained by concatenating the $i$th
code of $C_1$ with the $i$th code in $C_2$. The block lengths of
the codes in $C_1$ and $C_2$ tend to infinity as $i \to \infty$.)
Give an upper bound on the rate of $C$ as a function of its
minimum distance.
\begin{quote}
{\sf
Informally, If $C_1$ has rate $R_1$ and rel. distance $\delta_1$
then by the Singleton bound $R_1 + \delta_1 \leq 1$. Similarly
if $C_2$ has rate $R_2$ and rel. distance $\delta_1$, then
$R_2 + 2\delta_2 \leq 1$ (by the Plotkin bound, since $C_2$ is
binary). The concatenated code has rel. distance $\delta = \delta_1
\delta_2$ and rate $R \leq (1 - \delta_1)(1 - 2\delta_2)$.
Setting $\delta_2 = \delta/\delta_1$ and maximizing over
$\delta_1$ we get
$R \leq \max_{\delta \leq \delta_1 \leq 1} \{ 1 - \delta_1 -
2\delta/\delta_1 + 2\delta\} = 1 - 2\sqrt{2\delta} + 2\delta$.
A ``Mathematica'' plot reveals this is not as good as the GV bound.
Some analytic confirmation of this fact can be found when $\delta$
approaches $0$ - where the GV bound gives a rate of $1 - O(\delta
\log \frac1\delta)$, while this bound is $1 - \Omega(\sqrt{\delta})$.
On the other hand,
when $\delta = \frac12 - \epsilon$ and $\epsilon \to 0$,
this bound does not beat the GV bound asymptotically. (Both are
$\Theta(\epsilon^2)$.) However replacing the Plotkin bound above
with the LP bound will allow us to prove an upper bound on the rate
of concatenated codes of $O(\epsilon^3)$ which is again worse than
the GV bound.
To be more formal with any of the above, we need to take into
account the fact that specific members of these families are not
subject exactly to the Singleton bound/Plotkin bound, and furthermore
their rates may vary. To deal with all this formally, let
$R_{i,j}$ (resp. $\delta_{i,j}$) denote the rate (resp. relative
distance) of the $j$th code in the $i$th family. Since the block
lengths of the code $C_1$ tends to infinity, we have: For every
$\epsilon > 0$, there exists a $j_0$ such that for every $j \geq
j_0$, it is the case that $R_{1,j} + \delta_{1,j} \leq 1 + \epsilon$
(Singleton bound using $n_j \geq \frac1\epsilon$). Now since
the length of the codes $C_2$ tend to infinity, once again we have:
For every $\epsilon > 0$, there exists a $j'_0$ such that for all
$j \geq j'_0$, $R_{2,j}+2\delta_{2,j} \leq 1 + 2\epsilon$.
Combining the two, we have for every $j \geq \max\{j_0,j'_0\}$
the rate of the $j$th concatenated code $R_j = R_{1,j}\cdot R_{2,j}
\leq (1 - \delta_{1,j})(1 - 2\delta_{2,j}) + O(\epsilon)$.
Using $\delta_{1,j}\cdot \delta_{2,j} \geq \delta$ gives
$R_j \leq 1 - 2\sqrt{2\delta} + 2\delta + O(\epsilon)$.
Letting $\epsilon \to 0$ gives the bound of the first para,
formally.
}
\end{quote}
\item Consider the following simple
edit distance between strings: $x \in \Sigma^n$ is at distance
$d$ from $y \in \Sigma^m$ if $y$ can be obtained from $x$ by
first deleting upto $d$ coordinates of $x$ and getting an
intermediate string $z \in \Sigma^{\ell}$ where $\ell \geq n-d$,
and then inserting up to $d$ characters into $z$ (at arbitrary
locations) to get $y$. What are the analogs of the
Singleton bound, the Hamming (packing) bound on codes, and the
Gilbert-Varshamov bounds for this measure of distance?
\begin{quote}
{\sf
Both the Singleton and Hamming bounds obviously hold for
edit distance codes as well (since the edit distance is
upper bounded by the Hamming distance). The issue is: Are
they tight?
The Singleton bound is essentially tight. Consider an
$[n,k,n-k+1]$ RS code
$C$
over a $q$-ary alphabet with $q \gg n$. Now consider the
code $C'$ over the alphabet $\Sigma = \F_q \times [n]$,
whose codewords are strings of the form $(c'_1,\ldots,c'_n)$
where $c'_i = (i,c_i)$ and $(c_1,\ldots,c_n)$ is a codeword of
$C$. This code has $q^k = |\Sigma|^{k'}$ codewords with
edit distance $n-k+1$, where $k' = k \log q/(\log q + \log n)$.
As $q \to \infty$, this quantity approaches $k$, indicating that
these codes are achieving the Singleton bound $(d \geq n - k + 1)$.
As for the Hamming bound, obviously it is not tight since it
is not tight even for Hamming distance. However I don't believe
it is tight even for fixed $d$, unlike the case of the
Hamming distance.
I don't have a proof. Attempts welcome.
Finally the G-V bound in this case would be obtained by upper
bounding the volume of the ball of radius $d$. An easy bound
is roughly $\left({n \choose d}\right)^2 2^d$ (for a binary
alphabet). This would lead to the asymptotic bound saying there
exists codes of rate $R$ and distance $\delta$ for
$R = 1 - \delta - 2H(\delta)$.
Again, my guess would be that the G-V bound is close to being right
(and in particular, one can't have codes of relative distance
$\delta = .4$, say.) Proofs/counterexamples welcome.
}
\end{quote}
\end{enumerate}
\end{document}