# Hamming’s Code

## Or how to detect and correct errors

Last time we made a quick tour through the main theorems of Claude Shannon, which essentially solved the following two problems about communicating over a digital channel.

1. What is the best encoding for information when you are guaranteed that your communication channel is error free?
2. Are there any encoding schemes that can recover from random noise introduced during transmission?

The answers to these questions were purely mathematical theorems, of course. But the interesting shortcoming of Shannon’s accomplishment was that his solution for the noisy coding problem (2) was nonconstructive. The question remains: can we actually come up with efficiently computable encoding schemes? The answer is yes! Marcel Golay was the first to discover such a code in 1949 (just a year after Shannon’s landmark paper), and Golay’s construction was published on a single page! We’re not going to define Golay’s code in this post, but we will mention its interesting status in coding theory later. The next year Richard Hamming discovered another simpler and larger family of codes, and went on to do some of the major founding work in coding theory. For his efforts he won a Turing Award and played a major part in bringing about the modern digital age. So we’ll start with Hamming’s codes.

We will assume some basic linear algebra knowledge, as detailed our first linear algebra primer. We will also use some basic facts about polynomials and finite fields, though the lazy reader can just imagine everything as binary $\{ 0,1 \}$ and still grok the important stuff.

Richard Hamming, inventor of Hamming codes. [image source]

## What is a code?

The formal definition of a code is simple: a code $C$ is just a subset of $\{ 0,1 \}^n$ for some $n$. Elements of $C$ are called codewords.

This is deceptively simple, but here’s the intuition. Say we know we want to send messages of length $k$, so that our messages are in $\{ 0,1 \}^k$. Then we’re really viewing a code $C$ as the image of some encoding function $\textup{Enc}: \{ 0,1 \}^k \to \{ 0,1 \}^n$. We can define $C$ by just describing what the set is, or we can define it by describing the encoding function. Either way, we will make sure that $\textup{Enc}$ is an injective function, so that no two messages get sent to the same codeword. Then $|C| = 2^k$, and we can call $k = \log |C|$ the message length of $C$ even if we don’t have an explicit encoding function.

Moreover, while in this post we’ll always work with $\{ 0,1 \}$, the alphabet of your encoded messages could be an arbitrary set $\Sigma$. So then a code $C$ would be a subset of tuples in $\Sigma^n$, and we would call $q = |\Sigma|$.

So we have these parameters $n, k, q$, and we need one more. This is the minimum distance of a code, which we’ll denote by $d$. This is defined to be the minimum Hamming distance between all distinct pairs of codewords, where by Hamming distance I just mean the number of coordinates that two tuples differ in. Recalling the remarks we made last time about Shannon’s nonconstructive proof, when we decode an encoded message $y$ (possibly with noisy bits) we look for the (unencoded) message $x$ whose encoding $\textup{Enc}(x)$ is as close to $y$ as possible. This will only work in the worst case if all pairs of codewords are sufficiently far apart. Hence we track the minimum distance of a code.

So coding theorists turn this mess of parameters into notation.

Definition: A code $C$ is called an $(n, k, d)_q$-code if

• $C \subset \Sigma^n$ for some alphabet $\Sigma$,
• $k = \log |C|$,
• $C$ has minimum distance $d$, and
• the alphabet $\Sigma$ has size $q$.

The basic goals of coding theory are:

1. For which values of these four parameters do codes exist?
2. Fixing any three parameters, how can we optimize the other one?

In this post we’ll see how simple linear-algebraic constructions can give optima for one of these problems, optimizing $k$ for $d=3$, and we’ll state a characterization theorem for optimizing $k$ for a general $d$. Next time we’ll continue with a second construction that optimizes a different bound called the Singleton bound.

## Linear codes and the Hamming code

A code is called linear if it can be identified with a linear subspace of some finite-dimensional vector space. In this post all of our vector spaces will be $\{ 0,1 \}^n$, that is tuples of bits under addition mod 2. But you can do the same constructions with any finite scalar field $\mathbb{F}_q$ for a prime power $q$, i.e. have your vector space be $\mathbb{F}_q^n$. We’ll go back and forth between describing a binary code $q=2$ over $\{ 0,1 \}$ and a code in $\mathbb{F}_q^n$. So to say a code is linear means:

• The zero vector is a codeword.
• The sum of any two codewords is a codeword.
• Any scalar multiple of a codeword is a codeword.

Linear codes are the simplest kinds of codes, but already they give a rich variety of things to study. The benefit of linear codes is that you can describe them in a lot of different and useful ways besides just describing the encoding function. We’ll use two that we define here. The idea is simple: you can describe everything about a linear subspace by giving a basis for the space.

Definition: generator matrix of a $(n,k,d)_q$-code $C$ is a $k \times n$ matrix $G$ whose rows form a basis for $C$.

There are a lot of equivalent generator matrices for a linear code (we’ll come back to this later), but the main benefit is that having a generator matrix allows one to encode messages $x \in \{0,1 \}^k$ by left multiplication $xG$. Intuitively, we can think of the bits of $x$ as describing the coefficients of the chosen linear combination of the rows of $G$, which uniquely describes an element of the subspace. Note that because a $k$-dimensional subspace of $\{ 0,1 \}^n$ has $2^k$ elements, we’re not abusing notation by calling $k = \log |C|$ both the message length and the dimension.

For the second description of $C$, we’ll remind the reader that every linear subspace $C$ has a unique orthogonal complement $C^\perp$, which is the subspace of vectors that are orthogonal to vectors in $C$.

Definition: Let $H^T$ be a generator matrix for $C^\perp$. Then $H$ is called a parity check matrix.

Note $H$ has the basis for $C^\perp$ as columns. This means it has dimensions $n \times (n-k)$. Moreover, it has the property that $x \in C$ if and only if the left multiplication $xH = 0$. Having zero dot product with all columns of $H$ characterizes membership in $C$.

The benefit of having a parity check matrix is that you can do efficient error detection: just compute $yH$ on your received message $y$, and if it’s nonzero there was an error! What if there were so many errors, and just the right errors that $y$ coincided with a different codeword than it started? Then you’re screwed. In other words, the parity check matrix is only guarantee to detect errors if you have fewer errors than the minimum distance of your code.

So that raises an obvious question: if you give me the generator matrix of a linear code can I compute its minimum distance? It turns out that this problem is NP-hard in general. In fact, you can show that this is equivalent to finding the smallest linearly dependent set of rows of the parity check matrix, and it is easier to see why such a problem might be hard. But if you construct your codes cleverly enough you can compute their distance properties with ease.

Before we do that, one more definition and a simple proposition about linear codes. The Hamming weight of a vector $x$, denoted $wt(x)$, is the number of nonzero entries in $x$.

Proposition: The minimum distance of a linear code $C$ is the minimum Hamming weight over all nonzero vectors $x \in C$.

Proof. Consider a nonzero $x \in C$. On one hand, the zero vector is a codeword and $wt(x)$ is by definition the Hamming distance between $x$ and zero, so it is an upper bound on the minimum distance. In fact, it’s also a lower bound: if $x,y$ are two nonzero codewords, then $x-y$ is also a codeword and $wt(x-y)$ is the Hamming distance between $x$ and $y$.

$\square$

So now we can define our first code, the Hamming code. It will be a $(n, k, 3)_2$-code. The construction is quite simple. We have fixed $d=3, q=2$, and we will also fix $l = n-k$. One can think of this as fixing $n$ and maximizing $k$, but it will only work for $n$ of a special form.

We’ll construct the Hamming code by describing a parity-check matrix $H$. In fact, we’re going to see what conditions the minimum distance $d=3$ imposes on $H$, and find out those conditions are actually sufficient to get $d=3$. We’ll start with 2. If we want to ensure $d \geq 2$, then you need it to be the case that no nonzero vector of Hamming weight 1 is a code word. Indeed, if $e_i$ is a vector with all zeros except a one in position $i$, then $e_i H = h_i$ is the $i$-th row of $H$. We need $e_i H \neq 0$, so this imposes the condition that no row of $H$ can be zero. It’s easy to see that this is sufficient for $d \geq 2$.

Likewise for $d \geq 3$, given a vector $y = e_i + e_j$ for some positions $i \neq j$, then $yH = h_i + h_j$ may not be zero. But because our sums are mod 2, saying that $h_i + h_j \neq 0$ is the same as saying $h_i \neq h_j$. Again it’s an if and only if. So we have the two conditions.

• No row of $H$ may be zero.
• All rows of $H$ must be distinct.

That is, any parity check matrix with those two properties defines a distance 3 linear code. The only question that remains is how large can $n$  be if the vectors have length $n-k = l$? That’s just the number of distinct nonzero binary strings of length $l$, which is $2^l - 1$. Picking any way to arrange these strings as the rows of a matrix (say, in lexicographic order) gives you a good parity check matrix.

Theorem: For every $l > 0$, there is a $(2^l - 1, 2^l - l - 1, 3)_2$-code called the Hamming code.

Since the Hamming code has distance 3, we can always detect if at most a single error occurs. Moreover, we can correct a single error using the Hamming code. If $x \in C$ and $wt(e) = 1$ is an error bit in position $i$, then the incoming message would be $y = x + e$. Now compute $yH = xH + eH = 0 + eH = h_i$ and flip bit $i$ of $y$. That is, whichever row of $H$ you get tells you the index of the error, so you can flip the corresponding bit and correct it. If you order the rows lexicographically like we said, then $h_i = i$ as a binary number. Very slick.

Before we move on, we should note one interesting feature of linear codes.

Definition: A code is called systematic if it can be realized by an encoding function that appends some number $n-k$ “check bits” to the end of each message.

The interesting feature is that all linear codes are systematic. The reason is as follows. The generator matrix $G$ of a linear code has as rows a basis for the code as a linear subspace. We can perform Gaussian elimination on $G$ and get a new generator matrix that looks like $[I \mid A]$ where $I$ is the identity matrix of the appropriate size and $A$ is some junk. The point is that encoding using this generator matrix leaves the message unchanged, and adds a bunch of bits to the end that are determined by $A$. It’s a different encoding function on $\{ 0,1\}^k$, but it has the same image in $\{ 0,1 \}^n$, i.e. the code is unchanged. Gaussian elimination just performed a change of basis.

If you work out the parameters of the Hamming code, you’ll see that it is a systematic code which adds $\Theta(\log n)$ check bits to a message, and we’re able to correct a single error in this code. An obvious question is whether this is necessary? Could we get away with adding fewer check bits? The answer is no, and a simple “information theoretic” argument shows this. A single index out of $n$ requires $\log n$ bits to describe, and being able to correct a single error is like identifying a unique index. Without logarithmically many bits, you just don’t have enough information.

## The Hamming bound and perfect codes

One nice fact about Hamming codes is that they optimize a natural problem: the problem of maximizing $d$ given a fixed choice of $n$, $k$, and $q$. To get this let’s define $V_n(r)$ denote the volume of a ball of radius $r$ in the space $\mathbb{F}_2^n$. I.e., if you fix any string (doesn’t matter which) $x$, $V_n(r)$ is the size of the set $\{ y : d(x,y) \leq r \}$, where $d(x,y)$ is the hamming distance.

There is a theorem called the Hamming bound, which describes a limit to how much you can pack disjoint balls of radius $r$ inside $\mathbb{F}_2^n$.

Theorem: If an $(n,k,d)_2$-code exists, then

$\displaystyle 2^k V_n \left ( \left \lfloor \frac{d-1}{2} \right \rfloor \right ) \leq 2^n$

Proof. The proof is quite simple. To say a code $C$ has distance $d$ means that for every string $x \in C$ there is no other string $y$ within Hamming distance $d$ of $x$. In other words, the balls centered around both $x,y$ of radius $r = \lfloor (d-1)/2 \rfloor$ are disjoint. The extra difference of one is for odd $d$, e.g. when $d=3$ you need balls of radius 1 to guarantee no overlap. Now $|C| = 2^k$, so the total number of strings covered by all these balls is the left-hand side of the expression. But there are at most $2^n$ strings in $\mathbb{F}_2^n$, establishing the desired inequality.

$\square$

Now a code is called perfect if it actually meets the Hamming bound exactly. As you probably guessed, the Hamming codes are perfect codes. It’s not hard to prove this, and I’m leaving it as an exercise to the reader.

The obvious follow-up question is whether there are any other perfect codes. The answer is yes, some of which are nonlinear. But some of them are “trivial.” For example, when $d=1$ you can just use the identity encoding to get the code $C = \mathbb{F}_2^n$. You can also just have a code which consists of a single codeword. There are also some codes that encode by repeating the message multiple times. These are called “repetition codes,” and all three of these examples are called trivial (as a definition). Now there are some nontrivial and nonlinear perfect codes I won’t describe here, but here is the nice characterization theorem.

Theorem [van Lint ’71, Tietavainen ‘73]: Let $C$ be a nontrivial perfect $(n,d,k)_q$ code. Then the parameters must either be that of a Hamming code, or one of the two:

• A $(23, 12, 7)_2$-code
• A $(11, 6, 5)_3$-code

The last two examples are known as the binary and ternary Golay codes, respectively, which are also linear. In other words, every possible set of parameters for a perfect code can be realized as one of these three linear codes.

So this theorem was a big deal in coding theory. The Hamming and Golay codes were both discovered within a year of each other, in 1949 and 1950, but the nonexistence of other perfect linear codes was open for twenty more years. This wrapped up a very neat package.

Next time we’ll discuss the Singleton bound, which optimizes for a different quantity and is incomparable with perfect codes. We’ll define the Reed-Solomon and show they optimize this bound as well. These codes are particularly famous for being the error correcting codes used in DVDs. We’ll then discuss the algorithmic issues surrounding decoding, and more recent connections to complexity theory.

Until then!

# Linear Programming and the Simplex Algorithm

In the last post in this series we saw some simple examples of linear programs, derived the concept of a dual linear program, and saw the duality theorem and the complementary slackness conditions which give a rough sketch of the stopping criterion for an algorithm. This time we’ll go ahead and write this algorithm for solving linear programs, and next time we’ll apply the algorithm to an industry-strength version of the nutrition problem we saw last time. The algorithm we’ll implement is called the simplex algorithm. It was the first algorithm for solving linear programs, invented in the 1940’s by George Dantzig, and it’s still the leading practical algorithm, and it was a key part of a Nobel Prize. It’s by far one of the most important algorithms ever devised.

As usual, we’ll post all of the code written in the making of this post on this blog’s Github page.

## Slack variables and equality constraints

The simplex algorithm can solve any kind of linear program, but it only accepts a special form of the program as input. So first we have to do some manipulations. Recall that the primal form of a linear program was the following minimization problem.

$\min \left \langle c, x \right \rangle \\ \textup{s.t. } Ax \geq b, x \geq 0$

where the brackets mean “dot product.” And its dual is

$\max \left \langle y, b \right \rangle \\ \textup{s.t. } A^Ty \leq c, y \geq 0$

The linear program can actually have more complicated constraints than just the ones above. In general, one might want to have “greater than” and “less than” constraints in the same problem. It turns out that this isn’t any harder, and moreover the simplex algorithm only uses equality constraints, and with some finicky algebra we can turn any set of inequality or equality constraints into a set of equality constraints.

We’ll call our goal the “standard form,” which is as follows:

$\max \left \langle c, x \right \rangle \\ \textup{s.t. } Ax = b, x \geq 0$

It seems impossible to get the usual minimization/maximization problem into standard form until you realize there’s nothing stopping you from adding more variables to the problem. That is, say we’re given a constraint like:

$\displaystyle x_7 + x_3 \leq 10,$

we can add a new variable $\xi$, called a slack variable, so that we get an equality:

$\displaystyle x_7 + x_3 + \xi = 10$

And now we can just impose that $\xi \geq 0$. The idea is that $\xi$ represents how much “slack” there is in the inequality, and you can always choose it to make the condition an equality. So if the equality holds and the variables are nonnegative, then the $x_i$ will still satisfy their original inequality. For “greater than” constraints, we can do the same thing but subtract a nonnegative variable. Finally, if we have a minimization problem “$\min z$” we can convert it to $\max -z$.

So, to combine all of this together, if we have the following linear program with each kind of constraint,

We can add new variables $\xi_1, \xi_2$, and write it as

By defining the vector variable $x = (x_1, x_2, x_3, \xi_1, \xi_2)$ and $c = (-1,-1,-1,0,0)$ and $A$ to have $-1, 0, 1$ as appropriately for the new variables, we see that the system is written in standard form.

This is the kind of tedious transformation we can automate with a program. Assuming there are $n$ variables, the input consists of the vector $c$ of length $n$, and three matrix-vector pairs $(A, b)$ representing the three kinds of constraints. It’s a bit annoying to describe, but the essential idea is that we compute a rectangular “identity” matrix whose diagonal entries are $\pm 1$, and then join this with the original constraint matrix row-wise. The reader can see the full implementation in the Github repository for this post, though we won’t use this particular functionality in the algorithm that follows.

There are some other additional things we could do: for example there might be some variables that are completely unrestricted. What you do in this case is take an unrestricted variable $z$ and replace it by the difference of two unrestricted variables $z' - z''$.  For simplicity we’ll ignore this, but it would be a fruitful exercise for the reader to augment the function to account for these.

## What happened to the slackness conditions?

The “standard form” of our linear program raises an obvious question: how can the complementary slackness conditions make sense if everything is an equality? It turns out that one can redo all the work one did for linear programs of the form we gave last time (minimize w.r.t. greater-than constraints) for programs in the new “standard form” above. We even get the same complementary slackness conditions! If you want to, you can do this entire routine quite a bit faster if you invoke the power of Lagrangians. We won’t do that here, but the tool shows up as a way to work with primal-dual conversions in many other parts of mathematics, so it’s a good buzzword to keep in mind.

In our case, the only difference with the complementary slackness conditions is that one of the two is trivial: $\left \langle y^*, Ax^* - b \right \rangle = 0$. This is because if our candidate solution $x^*$ is feasible, then it will have to satisfy $Ax = b$ already. The other one, that $\left \langle x^*, A^Ty^* - c \right \rangle = 0$, is the only one we need to worry about.

Again, the complementary slackness conditions give us inspiration here. Recall that, informally, they say that when a variable is used at all, it is used as much as it can be to fulfill its constraint (the corresponding dual constraint is tight). So a solution will correspond to a choice of some variables which are either used or not, and a choice of nonzero variables will correspond to a solution. We even saw this happen in the last post when we observed that broccoli trumps oranges. If we can get a good handle on how to navigate the set of these solutions, then we’ll have a nifty algorithm.

Let’s make this official and lay out our assumptions.

## Extreme points and basic solutions

Remember that the graphical way to solve a linear program is to look at the line (or hyperplane) given by $\langle c, x \rangle = q$ and keep increasing $q$ (or decreasing it, if you are minimizing) until the very last moment when this line touches the region of feasible solutions. Also recall that the “feasible region” is just the set of all solutions to $Ax = b$, that is the solutions that satisfy the constraints. We imagined this picture:

The constraints define a convex area of “feasible solutions.” Image source: Wikipedia.

With this geometric intuition it’s clear that there will always be an optimal solution on a vertex of the feasible region. These points are called extreme points of the feasible region. But because we will almost never work in the plane again (even introducing slack variables makes us relatively high dimensional!) we want an algebraic characterization of these extreme points.

If you have a little bit of practice with convex sets the correct definition is very natural. Recall that a set $X$ is convex if for any two points $x, y \in X$ every point on the line segment between $x$ and $y$ is also in $X$. An algebraic way to say this (thinking of these points now as vectors) is that every point $\delta x + (1-\delta) y \in X$ when $0 \leq \delta \leq 1$. Now an extreme point is just a point that isn’t on the inside of any such line, i.e. can’t be written this way for $0 < \delta < 1$. For example,

A convex set with extremal points in red. Image credit Wikipedia.

Another way to say this is that if $z$ is an extreme point then whenever $z$ can be written as $\delta x + (1-\delta) y$ for some $0 < \delta < 1$, then actually $x=y=z$. Now since our constraints are all linear (and there are a finite number of them) they won’t define a convex set with weird curves like the one above. This means that there are a finite number of extreme points that just correspond to the intersections of some of the constraints. So there are at most $2^n$ possibilities.

Indeed we want a characterization of extreme points that’s specific to linear programs in standard form, “$\max \langle c, x \rangle \textup{ s.t. } Ax=b, x \geq 0$.” And here is one.

Definition: Let $A$ be an $m \times n$ matrix with $n \geq m$. A solution $x$ to $Ax=b$ is called basic if at most $m$ of its entries are nonzero.

The reason we call it “basic” is because, under some mild assumptions we describe below, a basic solution corresponds to a vector space basis of $\mathbb{R}^m$. Which basis? The one given by the $m$ columns of $A$ used in the basic solution. We don’t need to talk about bases like this, though, so in the event of a headache just think of the basis as a set $B \subset \{ 1, 2, \dots, n \}$ of size $m$ corresponding to the nonzero entries of the basic solution.

Indeed, what we’re doing here is looking at the matrix $A_B$ formed by taking the columns of $A$ whose indices are in $B$, and the vector $x_B$ in the same way, and looking at the equation $A_Bx_B = b$. If all the parts of $x$ that we removed were zero then this will hold if and only if $Ax=b$. One might worry that $A_B$ is not invertible, so we’ll go ahead and assume it is. In fact, we’ll assume that every set of $m$ columns of $A$ forms a basis and that the rows of $A$ are also linearly independent. This isn’t without loss of generality because if some rows or columns are not linearly independent, we can remove the offending constraints and variables without changing the set of solutions (this is why it’s so nice to work with the standard form).

Moreover, we’ll assume that every basic solution has exactly $m$ nonzero variables. A basic solution which doesn’t satisfy this assumption is called degenerate, and they’ll essentially be special corner cases in the simplex algorithm. Finally, we call a basic solution feasible if (in addition to satisfying $Ax=b$) it satisfies $x \geq 0$. Now that we’ve made all these assumptions it’s easy to see that choosing $m$ nonzero variables uniquely determines a basic feasible solution. Again calling the sub-matrix $A_B$ for a basis $B$, it’s just $x_B = A_B^{-1}b$. Now to finish our characterization, we just have to show that under the same assumptions basic feasible solutions are exactly the extremal points of the feasible region.

Proposition: A vector $x$ is a basic feasible solution if and only if it’s an extreme point of the set $\{ x : Ax = b, x \geq 0 \}$.

Proof. For one direction, suppose you have a basic feasible solution $x$, and say we write it as $x = \delta y + (1-\delta) z$ for some $0 < \delta < 1$. We want to show that this implies $y = z$. Since all of these points are in the feasible region, all of their coordinates are nonnegative. So whenever a coordinate $x_i = 0$ it must be that both $y_i = z_i = 0$. Since $x$ has exactly $n-m$ zero entries, it must be that $y, z$ both have at least $n-m$ zero entries, and hence $y,z$ are both basic. By our non-degeneracy assumption they both then have exactly $m$ nonzero entries. Let $B$ be the set of the nonzero indices of $x$. Because $Ay = Az = b$, we have $A(y-z) = 0$. Now $y-z$ has all of its nonzero entries in $B$, and because the columns of $A_B$ are linearly independent, the fact that $A_B(y-z) = 0$ implies $y-z = 0$.

In the other direction, suppose  that you have some extreme point $x$ which is feasible but not basic. In other words, there are more than $m$ nonzero entries of $x$, and we’ll call the indices $J = \{ j_1, \dots, j_t \}$ where $t > m$. The columns of $A_J$ are linearly dependent (since they’re $t$ vectors in $\mathbb{R}^m$), and so let $\sum_{i=1}^t z_{j_i} A_{j_i}$ be a nontrivial linear combination of the columns of $A$. Add zeros to make the $z_{j_i}$ into a length $n$ vector $z$, so that $Az = 0$. Now

$A(x + \varepsilon z) = A(x - \varepsilon z) = Ax = b$

And if we pick $\varepsilon$ sufficiently small $x \pm \varepsilon z$ will still be nonnegative, because the only entries we’re changing of $x$ are the strictly positive ones. Then $x = \delta (x + \varepsilon z) + (1 - \delta) \varepsilon z$ for $\delta = 1/2$, but this is very embarrassing for $x$ who was supposed to be an extreme point. $\square$

Now that we know extreme points are the same as basic feasible solutions, we need to show that any linear program that has some solution has a basic feasible solution. This is clear geometrically: any time you have an optimum it has to either lie on a line or at a vertex, and if it lies on a line then you can slide it to a vertex without changing its value. Nevertheless, it is a useful exercise to go through the algebra.

Theorem. Whenever a linear program is feasible and bounded, it has a basic feasible solution.

Proof. Let $x$ be an optimal solution to the LP. If $x$ has at most $m$ nonzero entries then it’s a basic solution and by the non-degeneracy assumption it must have exactly $m$ nonzero entries. In this case there’s nothing to do, so suppose that $x$ has $r > m$ nonzero entries. It can’t be a basic feasible solution, and hence is not an extreme point of the set of feasible solutions (as proved by the last theorem). So write it as $x = \delta y + (1-\delta) z$ for some feasible $y \neq z$ and $0 < \delta < 1$.

The only thing we know about $x$ is it’s optimal. Let $c$ be the cost vector, and the optimality says that $\langle c,x \rangle \geq \langle c,y \rangle$, and $\langle c,x \rangle \geq \langle c,z \rangle$. We claim that in fact these are equal, that $y, z$ are both optimal as well. Indeed, say $y$ were not optimal, then

$\displaystyle \langle c, y \rangle < \langle c,x \rangle = \delta \langle c,y \rangle + (1-\delta) \langle c,z \rangle$

Which can be rearranged to show that $\langle c,y \rangle < \langle c, z \rangle$. Unfortunately for $x$, this implies that it was not optimal all along:

$\displaystyle \langle c,x \rangle < \delta \langle c, z \rangle + (1-\delta) \langle c,z \rangle = \langle c,z \rangle$

An identical argument works to show $z$ is optimal, too. Now we claim we can use $y,z$ to get a new solution that has fewer than $r$ nonzero entries. Once we show this we’re done: inductively repeat the argument with the smaller solution until we get down to exactly $m$ nonzero variables. As before we know that $y,z$ must have at least as many zeros as $x$. If they have more zeros we’re done. And if they have exactly as many zeros we can do the following trick. Write $w = \gamma y + (1- \gamma)z$ for a $\gamma \in \mathbb{R}$ we’ll choose later. Note that no matter the $\gamma$, $w$ is optimal. Rewriting $w = z + \gamma (y-z)$, we just have to pick a $\gamma$ that ensures one of the nonzero coefficients of $z$ is zeroed out while maintaining nonnegativity. Indeed, we can just look at the index $i$ which minimizes $z_i / (y-z)_i$ and use $\delta = - z_i / (y-z)_i$. $\square$.

So we have an immediate (and inefficient) combinatorial algorithm: enumerate all subsets of size $m$, compute the corresponding basic feasible solution $x_B = A_B^{-1}b$, and see which gives the biggest objective value. The problem is that, even if we knew the value of $m$, this would take time $n^m$, and it’s not uncommon for $m$ to be in the tens or hundreds (and if we don’t know $m$ the trivial search is exponential).

So we have to be smarter, and this is where the simplex tableau comes in.

## The simplex tableau

Now say you have any basis $B$ and any feasible solution $x$. For now $x$ might not be a basic solution, and even if it is, its basis of nonzero entries might not be the same as $B$. We can decompose the equation $Ax = b$ into the basis part and the non basis part:

$A_Bx_B + A_{B'} x_{B'} = b$

and solving the equation for $x_B$ gives

$x_B = A^{-1}_B(b - A_{B'} x_{B'})$

It may look like we’re making a wicked abuse of notation here, but both $A_Bx_B$ and $A_{B'}x_{B'}$ are vectors of length $m$ so the dimensions actually do work out. Now our feasible solution $x$ has to satisfy $Ax = b$, and the entries of $x$ are all nonnegative, so it must be that $x_B \geq 0$ and $x_{B'} \geq 0$, and by the equality above $A^{-1}_B (b - A_{B'}x_{B'}) \geq 0$ as well. Now let’s write the maximization objective $\langle c, x \rangle$ by expanding it first in terms of the $x_B, x_{B'}$, and then expanding $x_B$.

\displaystyle \begin{aligned} \langle c, x \rangle & = \langle c_B, x_B \rangle + \langle c_{B'}, x_{B'} \rangle \\ & = \langle c_B, A^{-1}_B(b - A_{B'}x_{B'}) \rangle + \langle c_{B'}, x_{B'} \rangle \\ & = \langle c_B, A^{-1}_Bb \rangle + \langle c_{B'} - (A^{-1}_B A_{B'})^T c_B, x_{B'} \rangle \end{aligned}

If we want to maximize the objective, we can just maximize this last line. There are two cases. In the first, the vector $c_{B'} - (A^{-1}_B A_{B'})^T c_B \leq 0$ and $A_B^{-1}b \geq 0$. In the above equation, this tells us that making any component of $x_{B'}$ bigger will decrease the overall objective. In other words, $\langle c, x \rangle \leq \langle c_B, A_B^{-1}b \rangle$. Picking $x = A_B^{-1}b$ (with zeros in the non basis part) meets this bound and hence must be optimal. In other words, no matter what basis $B$ we’ve chosen (i.e., no matter the candidate basic feasible solution), if the two conditions hold then we’re done.

Now the crux of the algorithm is the second case: if the conditions aren’t met, we can pick a positive index of $c_{B'} - (A_B^{-1}A_{B'})^Tc_B$ and increase the corresponding value of $x_{B'}$ to increase the objective value. As we do this, other variables in the solution will change as well (by decreasing), and we have to stop when one of them hits zero. In doing so, this changes the basis by removing one index and adding another. In reality, we’ll figure out how much to increase ahead of time, and the change will correspond to a single elementary row-operation in a matrix.

Indeed, the matrix we’ll use to represent all of this data is called a tableau in the literature. The columns of the tableau will correspond to variables, and the rows to constraints. The last row of the tableau will maintain a candidate solution $y$ to the dual problem. Here’s a rough picture to keep the different parts clear while we go through the details.

But to make it work we do a slick trick, which is to “left-multiply everything” by $A_B^{-1}$. In particular, if we have an LP given by $c, A, b$, then for any basis it’s equivalent to the LP given by $c, A_B^{-1}A, A_{B}^{-1} b$ (just multiply your solution to the new program by $A_B$ to get a solution to the old one). And so the actual tableau will be of this form.

When we say it’s in this form, it’s really only true up to rearranging columns. This is because the chosen basis will always be represented by an identity matrix (as it is to start with), so to find the basis you can find the embedded identity sub-matrix. In fact, the beginning of the simplex algorithm will have the initial basis sitting in the last few columns of the tableau.

Let’s look a little bit closer at the last row. The first portion is zero because $A_B^{-1}A_B$ is the identity. But furthermore with this $A_B^{-1}$ trick the dual LP involves $A_B^{-1}$ everywhere there’s a variable. In particular, joining all but the last column of the last row of the tableau, we have the vector $c - A_B^T(A_B^{-1})^T c$, and setting $y = A_B^{-1}c_B$ we get a candidate solution for the dual. What makes the trick even slicker is that $A_B^{-1}b$ is already the candidate solution $x_B$, since $(A_B^{-1}A)_B^{-1}$ is the identity. So we’re implicitly keeping track of two solutions here, one for the primal LP, given by the last column of the tableau, and one for the dual, contained in the last row of the tableau.

I told you the last row was the dual solution, so why all the other crap there? This is the final slick in the trick: the last row further encodes the complementary slackness conditions. Now that we recognize the dual candidate sitting there, the complementary slackness conditions simply ask for the last row to be non-positive (this is just another way of saying what we said at the beginning of this section!). You should check this, but it gives us a stopping criterion: if the last row is non-positive then stop and output the last column.

## The simplex algorithm

Now (finally!) we can describe and implement the simplex algorithm in its full glory. Recall that our informal setup has been:

1. Find an initial basic feasible solution, and set up the corresponding tableau.
2. Find a positive index of the last row, and increase the corresponding variable (adding it to the basis) just enough to make another variable from the basis zero (removing it from the basis).
3. Repeat step 2 until the last row is nonpositive.
4. Output the last column.

This is almost correct, except for some details about how increasing the corresponding variables works. What we’ll really do is represent the basis variables as pivots (ones in the tableau) and then the first 1 in each row will be the variable whose value is given by the entry in the last column of that row. So, for example, the last entry in the first row may be the optimal value for $x_5$, if the fifth column is the first entry in row 1 to have a 1.

As we describe the algorithm, we’ll illustrate it running on a simple example. In doing this we’ll see what all the different parts of the tableau correspond to from the previous section in each step of the algorithm.

Spoiler alert: the optimum is $x_1 = 2, x_2 = 1$ and the value of the max is 8.

So let’s be more programmatically formal about this. The main routine is essentially pseudocode, and the difficulty is in implementing the helper functions

def simplex(c, A, b):
tableau = initialTableau(c, A, b)

while canImprove(tableau):
pivot = findPivotIndex(tableau)

return primalSolution(tableau), objectiveValue(tableau)


Let’s start with the initial tableau. We’ll assume the user’s inputs already include the slack variables. In particular, our example data before adding slack is

c = [3, 2]
A = [[1, 2], [1, -1]]
b = [4, 1]


c = [3, 2, 0, 0]
A = [[1,  2,  1,  0],
[1, -1,  0,  1]]
b = [4, 1]


Now to set up the initial tableau we need an initial feasible solution in mind. The reader is recommended to work this part out with a pencil, since it’s much easier to write down than it is to explain. Since we introduced slack variables, our initial feasible solution (basis) $B$ can just be $(0,0,1,1)$. And so $x_B$ is just the slack variables, $c_B$ is the zero vector, and $A_B$ is the 2×2 identity matrix. Now $A_B^{-1}A_{B'} = A_{B'}$, which is just the original two columns of $A$ we started with, and $A_B^{-1}b = b$. For the last row, $c_B$ is zero so the part under $A_B^{-1}A_B$ is the zero vector. The part under $A_B^{-1}A_{B'}$ is just $c_{B'} = (3,2)$.

Rather than move columns around every time the basis $B$ changes, we’ll keep the tableau columns in order of $(x_1, \dots, x_n, \xi_1, \dots, \xi_m)$. In other words, for our example the initial tableau should look like this.

[[ 1,  2,  1,  0,  4],
[ 1, -1,  0,  1,  1],
[ 3,  2,  0,  0,  0]]


So implementing initialTableau is just a matter of putting the data in the right place.

def initialTableau(c, A, b):
tableau = [row[:] + [x] for row, x in zip(A, b)]
tableau.append(c[:] + [0])
return tableau


As an aside: in the event that we don’t start with the trivial basic feasible solution of “trivially use the slack variables,” we’d have to do a lot more work in this function. Next, the primalSolution() and objectiveValue() functions are simple, because they just extract the encoded information out from the tableau (some helper functions are omitted for brevity).

def primalSolution(tableau):
# the pivot columns denote which variables are used
columns = transpose(tableau)
indices = [j for j, col in enumerate(columns[:-1]) if isPivotCol(col)]
return list(zip(indices, columns[-1]))

def objectiveValue(tableau):
return -(tableau[-1][-1])


Similarly, the canImprove() function just checks if there’s a nonnegative entry in the last row

def canImprove(tableau):
lastRow = tableau[-1]
return any(x &gt; 0 for x in lastRow[:-1])


Let’s run the first loop of our simplex algorithm. The first step is checking to see if anything can be improved (in our example it can). Then we have to find a pivot entry in the tableau. This part includes some edge-case checking, but if the edge cases aren’t a problem then the strategy is simple: find a positive entry corresponding to some entry $j$ of $B'$, and then pick an appropriate entry in that column to use as the pivot. Pivoting increases the value of $x_j$ (from zero) to whatever is the largest we can make it without making some other variables become negative. As we’ve said before, we’ll stop increasing $x_j$ when some other variable hits zero, and we can compute which will be the first to do so by looking at the current values of $x_B = A_B^{-1}b$ (in the last column of the tableau), and seeing how pivoting will affect them. If you stare at it for long enough, it becomes clear that the first variable to hit zero will be the entry $x_i$ of the basis for which $x_i / A_{i,j}$ is minimal (and $A_{i,j}$ has to be positve). This is because, in order to maintain the linear equalities, every entry of $x_B$ will be decreased by that value during a pivot, and we can’t let any of the variables become negative.

All of this results in the following function, where we have left out the degeneracy/unboundedness checks.

def findPivotIndex(tableau):
# pick first nonzero index of the last row
column = [i for i,x in enumerate(tableau[-1][:-1]) if x &gt; 0][0]
quotients = [(i, r[-1] / r[column]) for i,r in enumerate(tableau[:-1]) if r[column] &gt; 0]

# pick row index minimizing the quotient
row = min(quotients, key=lambda x: x[1])[0]
return row, column


For our example, the minimizer is the $(1,0)$ entry (second row, first column). Pivoting is just doing the usual elementary row operations (we covered this in a primer a while back on row-reduction). The pivot function we use here is no different, and in particular mutates the list in place.

def pivotAbout(tableau, pivot):
i,j = pivot

pivotDenom = tableau[i][j]
tableau[i] = [x / pivotDenom for x in tableau[i]]

for k,row in enumerate(tableau):
if k != i:
pivotRowMultiple = [y * tableau[k][j] for y in tableau[i]]
tableau[k] = [x - y for x,y in zip(tableau[k], pivotRowMultiple)]


And in our example pivoting around the chosen entry gives the new tableau.

[[ 0.,  3.,  1., -1.,  3.],
[ 1., -1.,  0.,  1.,  1.],
[ 0.,  5.,  0., -3., -3.]]


In particular, $B$ is now $(1,0,1,0)$, since our pivot removed the second slack variable $\xi_2$ from the basis. Currently our solution has $x_1 = 1, \xi_1 = 3$. Notice how the identity submatrix is still sitting in there, the columns are just swapped around.

There’s still a positive entry in the bottom row, so let’s continue. The next pivot is (0,1), and pivoting around that entry gives the following tableau:

[[ 0.        ,  1.        ,  0.33333333, -0.33333333,  1.        ],
[ 1.        ,  0.        ,  0.33333333,  0.66666667,  2.        ],
[ 0.        ,  0.        , -1.66666667, -1.33333333, -8.        ]]


And because all of the entries in the bottom row are negative, we’re done. We read off the solution as we described, so that the first variable is 2 and the second is 1, and the objective value is the opposite of the bottom right entry, 8.

To see all of the source code, including the edge-case-checking we left out of this post, see the Github repository for this post.

An obvious question is: what is the runtime of the simplex algorithm? Is it polynomial in the size of the tableau? Is it even guaranteed to stop at some point? The surprising truth is that nobody knows the answer to all of these questions! Originally (in the 1940’s) the simplex algorithm actually had an exponential runtime in the worst case, though this was not known until 1972. And indeed, to this day while some variations are known to terminate, no variation is known to have polynomial runtime in the worst case. Some of the choices we made in our implementation (for example, picking the first column with a positive entry in the bottom row) have the potential to cycle, i.e., variables leave and enter the basis without changing the objective at all. Doing something like picking a random positive column, or picking the column which will increase the objective value by the largest amount are alternatives. Unfortunately, every single pivot-picking rule is known to give rise to exponential-time simplex algorithms in the worst case (in fact, this was discovered as recently as 2011!). So it remains open whether there is a variant of the simplex method that runs in guaranteed polynomial time.

But then, in a stunning turn of events, Leonid Khachiyan proved in the 70’s that in fact linear programs can always be solved in polynomial time, via a completely different algorithm called the ellipsoid method. Following that was a method called the interior point method, which is significantly more efficient. Both of these algorithms generalize to problems that are harder than linear programming as well, so we will probably cover them in the distant future of this blog.

Despite the celebratory nature of these two results, people still use the simplex algorithm for industrial applications of linear programming. The reason is that it’s much faster in practice, and much simpler to implement and experiment with.

The next obvious question has to do with the poignant observation that whole numbers are great. That is, you often want the solution to your problem to involve integers, and not real numbers. But adding the constraint that the variables in a linear program need to be integer valued (even just 0-1 valued!) is NP-complete. This problem is called integer linear programming, or just integer programming (IP). So we can’t hope to solve IP, and rightly so: the reader can verify easily that boolean satisfiability instances can be written as linear programs where each clause corresponds to a constraint.

This brings up a very interesting theoretical issue: if we take an integer program and just remove the integrality constraints, and solve the resulting linear program, how far away are the two solutions? If they’re close, then we can hope to give a good approximation to the integer program by solving the linear program and somehow turning the resulting solution back into an integer solution. In fact this is a very popular technique called LP-rounding. We’ll also likely cover that on this blog at some point.

Oh there’s so much to do and so little time! Until next time.

# Learning a single-variable polynomial, or the power of adaptive queries

Problem: Alice chooses a secret polynomial $p(x)$ with nonnegative integer coefficients. Bob wants to discover this polynomial by querying Alice for the value of $p(x)$ for some integer $x$ of Bob’s choice. What is the minimal number of queries Bob needs to determine $p(x)$ exactly?

Solution: Two queries. The first is $p(1)$, and if we call $N = p(1) + 1$, then the second query is $p(N)$.

To someone who is familiar with polynomials, this may seem shocking, and I’ll explain why it works in a second. After all, it’s very easy to prove that if Bob gives Alice all of his queries at the same time (if the queries are not adaptive), then it’s impossible to discover what $p(x)$ is using fewer than $\textup{deg}(p) + 1$ queries. This is due to a fact called polynomial interpolation, which we’ve seen on this blog before in the context of secret sharing. Specifically, there is a unique single-variable degree $d$ polynomial passing through $d+1$ points (with distinct $x$-values). So if you knew the degree of $p$, you could determine it easily. But Bob doesn’t know the degree of the polynomial, and there’s no way he can figure it out without adaptive queries! Indeed, if Bob tries and gives a set of $d$ queries, Alice could have easily picked a polynomial of degree $d+1$. So it’s literally impossible to solve this problem without adaptive queries.

The lovely fact is that once you allow adaptiveness, the number of queries you need doesn’t even depend on the degree of the secret polynomial!

Okay let’s get to the solution. It was crucial that our polynomial had nonnegative integer coefficients, because we’re going to do a tiny bit of number theory. Let $p(x) = a_0 + a_1 x + \dots + a_d x^d$. First, note that $p(1)$ is exactly the sum of the coefficients $\sum_i a_i$, and in particular $p(1) + 1$ is larger than any single coefficient. So call this $N$, and query $p(N)$. This gives us a number $y_0$ of the form

$\displaystyle y_0 = a_0 + a_1N + a_2N^2 + \dots + a_dN^d$

And because $N$ is so big, we can compute $a_0$ easily by computing $y_0 \mod N$. Now set $y_1 = (y_0 - a_0) / N$, and this has the form $a_1 + a_2N + \dots + a_dN^{d-1}$. We can compute modulus again to get $a_1$, and repeat until we have all the coefficients. We’ll stop once we get a $y_i$ that is zero.

As a small technical note, this is a polynomial-time algorithm in the number of bits needed to write down $p(x)$. So this demonstrates the power of adaptive queries: we get from something which is uncomputable with any number of queries to something which is efficiently computable with a constant number of queries.

The obvious follow-up question is: can you come up with an efficient algorithm if we allow the coefficients to be negative integers?

# On the Computational Complexity of MapReduce

I recently wrapped up a fun paper with my coauthors Ben Fish, Adam Lelkes, Lev Reyzin, and Gyorgy Turan in which we analyzed the computational complexity of a model of the popular MapReduce framework. Check out the preprint on the arXiv.

As usual I’ll give a less formal discussion of the research here, and because the paper is a bit more technically involved than my previous work I’ll be omitting some of the more pedantic details. Our project started after Ben Moseley gave an excellent talk at UI Chicago. He presented a theoretical model of MapReduce introduced by Howard Karloff et al. in 2010, and discussed his own results on solving graph problems in this model, such as graph connectivity. You can read Karloff’s original paper here, but we’ll outline his model below.

Basically, the vast majority of the work on MapReduce has been algorithmic. What I mean by that is researchers have been finding more and cleverer algorithms to solve problems in MapReduce. They have covered a huge amount of work, implementing machine learning algorithms, algorithms for graph problems, and many others. In Moseley’s talk, he posed a question that caught our eye:

Is there a constant-round MapReduce algorithm which determines whether a graph is connected?

After we describe the model below it’ll be clear what we mean by “solve” and what we mean by “constant-round,” but the conjecture is that this is impossible, particularly for the case of sparse graphs. We know we can solve it in a logarithmic number of rounds, but anything better is open.

In any case, we started thinking about this problem and didn’t make much progress. To the best of my knowledge it’s still wide open. But along the way we got into a whole nest of more general questions about the power of MapReduce. Specifically, Karloff proved a theorem relating MapReduce to a very particular class of circuits. What I mean is he proved a theorem that says “anything that can be solved in MapReduce with so many rounds and so much space can be solved by circuits that are yae big and yae complicated, and vice versa.

But this question is so specific! We wanted to know: is MapReduce as powerful as polynomial time, our classical notion of efficiency (does it equal P)? Can it capture all computations requiring logarithmic space (does it contain L)? MapReduce seems to be somewhere in between, but it’s exact relationship to these classes is unknown. And as we’ll see in a moment the theoretical model uses a novel communication model, and processors that never get to see the entire input. So this led us to a host of natural complexity questions:

1. What computations are possible in a model of parallel computation where no processor has enough space to store even one thousandth of the input?
2. What computations are possible in a model of parallel computation where processor’s can’t request or send specific information from/to other processors?
3. How the hell do you prove that something can’t be done under constraints of this kind?
4. How do you measure the increase of power provided by giving MapReduce additional rounds or additional time?

These questions are in the domain of complexity theory, and so it makes sense to try to apply the standard tools of complexity theory to answer them. Our paper does this, laying some brick for future efforts to study MapReduce from a complexity perspective.

In particular, one of our theorems is the following:

Theorem: Any problem requiring sublogarithmic space $o(\log n)$ can be solved in MapReduce in two rounds.

This theorem is nice for two reasons. First is it’s constructive. If you give me a problem and I know it classically takes less than logarithmic space, then this gives an automatic algorithm to implement it in MapReduce, and if you were so inclined you could even automate the process of translating a classical algorithm to a MapReduce algorithm (it’s not pretty, but it’s possible).

Our other results are a bit more esoteric. We relate the questions about MapReduce to old, deep questions about computations that have simultaneous space and time bounds. In the end we give a (conditional, nonconstructive) answer to question 4 above, which I’ll sketch without getting too deep in the details.

So let’s start with the model of Karloff et al., which they named “MRC” for MapReduce Class.

## The MRC Model of Karloff et al.

MapReduce is a programming paradigm which is intended to make algorithm design for distributed computing easier. Specifically, when you’re writing massively distributed algorithms by hand, you spend a lot of time and effort dealing with really low-level questions. Like, what do I do if a processor craps out in the middle of my computation? Or, what’s the most efficient way to broadcast a message to all the processors, considering the specific topology of my network layout means the message will have to be forwarded? Note these are questions that have nothing to do with the actual problem you’re trying to solve.

So the designers of MapReduce took a step back, analyzed the class of problems they were most often trying to solve (turns out, searching, sorting, and counting), and designed their framework to handle all of the low-level stuff automatically. The catch is that the algorithm designer now has to fit their program into the MapReduce paradigm, which might be hard.

In designing a MapReduce algorithm, one has to implement two functions: a mapper and a reducer. The mapper takes a list of key-value pairs, and applies some operation to each element independently (just like the map function in most functional programming languages). The reducer takes a single key and a list of values for that key, and outputs a new list of values with the same key. And then the MapReduce protocol iteratively applies mappers and reducers in rounds to the input data. There are a lot of pictures of this protocol on the internet. Here’s one

Image source: ibm.com

An interesting part of this protocol is that the reducers never get to communicate with each other, except indirectly through the mappers in the next round. MapReduce handles the implied grouping and message passing, as well as engineering issues like fault-tolerance and load balancing. In this sense, the mappers and reducers are ignorant cogs in the machine, so it’s interesting to see how powerful MapReduce is.

The model of Karloff, Suri, and Vassilvitskii is a relatively straightforward translation of this protocol to mathematics. They make a few minor qualifications, though, the most important of which is that they encode a bound on the total space used. In particular, if the input has size $n$, they impose that there is some $\varepsilon > 0$ for which every reducer uses space $O(n^{1 - \varepsilon})$. Moreover, the number of reducers in any round is also bounded by $O(n^{1 - \varepsilon})$.

We can formalize all of this with a few easy definitions.

Definition: An input to a MapReduce algorithm is a list of key-value pairs $\langle k_i,v_i \rangle_{i=1}^N$ of total size $n = \sum_{i=1}^N |k_i| + |v_i|$.

For binary languages (e.g., you give me a binary string $x$ and you want to know if there are an odd number of 1’s), we encode a string $x \in \{ 0,1 \}^m$ as the list $\langle x \rangle := \langle i, x_i \rangle_{i=1}^n$. This won’t affect any of our space bounds, because the total blowup from $m = |x|$ to $n = |\langle x \rangle |$ is logarithmic.

Definition: A mapper $\mu$ is a Turing machine which accepts as input a single key-value pair $\langle k, v \rangle$ and produces as output a list of key-value pairs $\langle k_1', v'_1 \rangle, \dots, \langle k'_s, v'_s \rangle$.

Definition: reducer $\rho$ is a Turing machine which accepts as input a key $k$ and a list of values $v_1, \dots, v_m$ and produces as output the same key and a new list of values $v'_1, \dots, v'_M$.

Now we can describe the MapReduce protocol, i.e. what a program is and how to run it. I copied this from our paper because the notation is the same so far.

The MRC protocol

All we’ve done here is just describe the MapReduce protocol in mathematics. It’s messier than it is complicated. The last thing we need is the space bounds.

Definition: A language $L$ is in $\textup{MRC}[f(n), g(n)]$ if there is a constant $0 < c < 1$ and a sequence of mappers and reducers $\mu_1, \rho_1, \mu_2, \rho_2, \dots$ such that for all $x \in \{ 0,1 \}^n$ the following is satisfied:

1. Letting $R = O(f(n))$ and $M = (\mu_1, \rho_1, \dots, \mu_R, \rho_R)$, $M$ accepts $x$ if and only if $x \in L$.
2. For all $1 \leq r \leq R$, $\mu_r, \rho_r$ use $O(n^c)$ space and $O(g(n))$ time.
3. Each $\mu_r$ outputs $O(n^c)$ keys in round $r$.

To be clear, the first argument to $\textup{MRC}[f(n), g(n)]$ is the round bound, and the second argument is the time bound. The last point implies the bound on the number of processors. Since we are often interested in a logarithmic number of rounds, we can define

$\displaystyle \textup{MRC}^i = \textup{MRC}[\log^i(n), \textup{poly}(n)]$

So we can restate the question from the beginning of the post as,

Is graph connectivity in $\textup{MRC}^0$?

Here there’s a bit of an issue with representation. You have to show that if you’re just given a binary string representing a graph, that you can translate that into a reasonable key-value description of a graph in a constant number of rounds. This is possible, and gives evidence that the key-value representation is without loss of generality for all intents and purposes.

## A Pedantic Aside

If our goal is to compare MRC with classes like polynomial time (P) and logarithmic space (L), then the definition above has a problem. Specifically the definition above allows one to have an infinite list of reducers, where each one is potentially different, and the actual machine that is used depends on the input size. In formal terminology, MRC as defined above is a nonuniform model of computation.

Indeed, we give a simple proof that this is a problem by showing any unary language is in $\textup{MRC}^1$ (which is where many of the MRC algorithms in recent years have been). For those who don’t know, this is a problem because you can encode versions of the Halting problem as a unary language, and the Halting problem is undecidable.

While this level of pedantry might induce some eye-rolling, it is necessary in order to make statements like “MRC is contained in P.” It’s simply not true for the nonuniform model above. To fix this problem we refined the MRC model to a uniform version. The details are not trivial, but also not that interesting. Check out the paper if you want to know exactly how we do it. It doesn’t have that much of an effect on the resulting analysis. The only important detail is that now we’re representing the entire MRC machine as a single Turing machine. So now, unlike before, any MRC machine can be written down as a finite string, and there are infinitely many strings representing a given MRC machine. We call our model MRC, and Karloff’s model nonuniform MRC.

You can also make randomized versions of MRC, but we’ll stick to the deterministic version here.

## Sublogarithmic Space

Now we can sketch the proof that sublogarithmic space is in $\textup{MRC}^0$. In fact, the proof is simpler for regular languages (equivalent to constant space algorithms) being in $\textup{MRC}^0$. The idea is as follows:

A regular language is one that can be decided by a DFA, say $G$ (a graph representing state transitions as you read a stream of bits). And the DFA is independent of the input size, so every mapper and reducer can have it encoded into them. Now what we’ll do is split the input string up in contiguous chunks of size $O(\sqrt{n})$ among the processors (mappers can do this just fine). Now comes the trick. We have each reducer compute, for each possible state $s$ in $G$, what the ending state would be if the DFA had started in state $s$ after processing their chunk of the input. So the output of reducer $j$ would be an encoding of a table of the form:

$\displaystyle s_1 \to T_j(s_1) \\ s_2 \to T_j(s_2) \\ \vdots \\ s_{|S|} \to T_j(s_{|S|})$

And the result would be a lookup table of intermediate computation results. So each reducer outputs their part of the table (which has constant size). Since there are only $O(\sqrt{n})$ reducers, all of the table pieces can fit on a single reducer in the second round, and this reducer can just chain the functions together from the starting state and output the answer.

The proof for sublogarithmic space has the same structure, but is a bit more complicated because one has to account for the tape of a Turing machine which has size $o(\log n)$. In this case, you split up the tape of the Turing machine among the processors as well. And then you have to keep track of a lot more information. In particular, each entry of your function has to look like

“if my current state is A and my tape contents are B and the tape head starts on side C of my chunk of the input”

then

“the next time the tape head would leave my chunk, it will do so on side C’ and my state will be A’ and my tape contents will be B’.”

As complicated as it sounds, the size of one of these tables for one reducer is still less than $n^c$ for some $c < 1/2$. And so we can do the same trick of chaining the functions together to get the final answer. Note that this time the chaining will be adaptive, because whether the tape head leaves the left or right side will determine which part of the table you look at next. And moreover, we know the chaining process will stop in polynomial time because we can always pick a Turing machine to begin with that will halt in polynomial time (i.e., we know that L is contained in P).

The size calculations are also just large enough to stop us from doing the same trick with logarithmic space. So that gives the obvious question: is $\textup{L} \subset \textup{MRC}^0$? The second part of our paper shows that even weaker containments are probably very hard to prove, and they relate questions about MRC to the problem of L vs P.

## Hierarchy Theorems

This part of the paper is where we go much deeper into complexity theory than I imagine most of my readers are comfortable with. Our main theorem looks like this:

Theorem: Assume some conjecture is true. Then for every $a, b > 0$, there is are bigger $c > a, d > b$ such that the following hold:

$\displaystyle \textup{MRC}[n^a, n^b] \subsetneq \textup{MRC}[n^c, n^b],$
$\displaystyle \textup{MRC}[n^a, n^b] \subsetneq \textup{MRC}[n^a, n^d].$

This is a kind of hierarchy theorem that one often likes to prove in complexity theory. It says that if you give MRC enough extra rounds or time, then you’ll get strictly more computational power.

The conjecture we depend on is called the exponential time hypothesis (ETH), and it says that the 3-SAT problem can’t be solved in $2^{cn}$ time for any $0 < c < 1$. Actually, our theorem depends on a weaker conjecture (implied by ETH), but it’s easier to understand our theorems in the context of the ETH. Because 3-SAT is this interesting problem: we believe it’s time-intensive, and yet it’s space efficient because we can solve it in linear space given $2^n$ time. This time/space tradeoff is one of the oldest questions in all of computer science, but we still don’t have a lot of answers.

For instance, define $\textup{TISP}(f(n), g(n))$ to the the class of languages that can be decided by Turing machines that have simultaneous bounds of $O(f(n))$ time and $O(g(n))$ space. For example, we just said that $\textup{3-SAT} \in \textup{TISP}(2^n, n)$, and there is a famous theorem that says that SAT is not in $\textup{TISP}(n^{1.1} n^{0.1})$; apparently this is the best we know. So clearly there are some very basic unsolved problems about TISP. An important one that we address in our paper is whether there are hierarchies in TISP for a fixed amount of space. This is the key ingredient in proving our hierarchy theorem for MRC. In particular here is an open problem:

Conjecture: For every two integers $0 < a < b$, the classes $\textup{TISP}(n^a, n)$ and $\textup{TISP}(n^b, n)$ are different.

We know this is true of time and space separately, i.e., that $\textup{TIME}(n^a) \subsetneq \textup{TIME}(n^b)$ and $\textup{SPACE}(n^a) \subsetneq \textup{SPACE}(n^b)$. but for TISP we only know that you get more power if you increase both parameters.

So we prove a theorem that address this, but falls short in two respects: it depends on a conjecture like ETH, and it’s not for every $a, b$.

Theorem: Suppose ETH holds, then for every $a > 0$ there is a $b > a$ for which $\textup{TIME}(n^a) \not \subset \textup{TISP}(n^b, n)$.

In words, this suggests that there are problems that need both $n^2$ time and $n^2$ space, but can be solved with linear space if you allow enough extra time. Since $\textup{TISP}(n^a, n) \subset \textup{TIME}(n^a)$, this gives us the hierarchy we wanted.

To prove this we take 3-SAT and give it exponential padding so that it becomes easy enough to do in polynomial TISP (and it still takes linear space, in fact sublinear), but not so easy that you can do it in $n^a$ time. It takes some more work to get from this TISP hierarchy to our MRC hierarchy, but the details are a bit much for this blog. One thing I’d like to point out is that we prove that statements that are just about MRC have implications beyond MapReduce. In particular, the last corollary of our paper is the following.

Corollary: If $\textup{MRC}[\textup{poly}(n), 1] \subsetneq \textup{MRC}[\textup{poly}(n), \textup{poly}(n)]$, then L is different from P.

In other words, if you’re afforded a polynomial number of rounds in MRC, then showing that constant time per round is weaker than polynomial time is equivalently hard to separating L from P. The theorem is true because, as it turns out, $\textup{L} \subset \textup{MRC}[textup{poly}(n), 1]$, by simulating one step of a TM across polynomially many rounds. The proof is actually not that complicated (and doesn’t depend on ETH at all), and it’s a nice demonstration that studying MRC can have implications beyond parallel computing.

The other side of the coin is also interesting. Our first theorem implies the natural question of whether $\textup{L} \subset \textup{MRC}^0$. We’d like to say that this would imply the separation of L from P, but we don’t quite get that. In particular, we know that

$\displaystyle \textup{MRC}[1, n] \subsetneq \textup{MRC}[n, n] \subset \textup{MRC}[1, n^2] \subsetneq \textup{MRC}[n^2, n^2] \subset \dots$

But at the same time we could live in a world where

$\displaystyle \textup{MRC}[1, \textup{poly}(n)] = \textup{MRC}[\textup{poly}(n), \textup{poly}(n)]$

It seems highly unlikely, but to the best of our knowledge none of our techniques prove this is not the case. If we could rule this out, then we could say that $\textup{L} \subset \textup{MRC}^0$ implies the separation of L and P. And note this would not depend on any conjectures.

## Open Problems

Our paper has a list of open problems at the end. My favorite is essentially: how do we prove better lower bounds in MRC? In particular, it would be great if we could show that some problems need a lot of rounds simply because the communication model is too restrictive, and nobody has true random access to the entire input. For example, this is why we think graph connectivity needs a logarithmic number of rounds. But as of now nobody really knows how to prove it, and it seems like we’ll need some new and novel techniques in order to do it. I only have the wisps of ideas in that regard, and it will be fun to see which ones pan out.

Until next time!