Singular Value Decomposition Part 2: Theorem, Proof, Algorithm

I’m just going to jump right into the definitions and rigor, so if you haven’t read the previous post motivating the singular value decomposition, go back and do that first. This post will be theorem, proof, algorithm, data. The data set we test on is a thousand-story CNN news data set. All of the data, code, and examples used in this post is in a github repository, as usual.

We start with the best-approximating $ k$-dimensional linear subspace.

Definition: Let $ X = \{ x_1, \dots, x_m \}$ be a set of $ m$ points in $ \mathbb{R}^n$. The best approximating $ k$-dimensional linear subspace of $ X$ is the $ k$-dimensional linear subspace $ V \subset \mathbb{R}^n$ which minimizes the sum of the squared distances from the points in $ X$ to $ V$.

Let me clarify what I mean by minimizing the sum of squared distances. First we’ll start with the simple case: we have a vector $ x \in X$, and a candidate line $ L$ (a 1-dimensional subspace) that is the span of a unit vector $ v$. The squared distance from $ x$ to the line spanned by $ v$ is the squared length of $ x$ minus the squared length of the projection of $ x$ onto $ v$. Here’s a picture.


I’m saying that the pink vector $ z$ in the picture is the difference of the black and green vectors $ x-y$, and that the “distance” from $ x$ to $ v$ is the length of the pink vector. The reason is just the Pythagorean theorem: the vector $ x$ is the hypotenuse of a right triangle whose other two sides are the projected vector $ y$ and the difference vector $ z$.

Let’s throw down some notation. I’ll call $ \textup{proj}_v: \mathbb{R}^n \to \mathbb{R}^n$ the linear map that takes as input a vector $ x$ and produces as output the projection of $ x$ onto $ v$. In fact we have a brief formula for this when $ v$ is a unit vector. If we call $ x \cdot v$ the usual dot product, then $ \textup{proj}_v(x) = (x \cdot v)v$. That’s $ v$ scaled by the inner product of $ x$ and $ v$. In the picture above, since the line $ L$ is the span of the vector $ v$, that means that $ y = \textup{proj}_v(x)$ and $ z = x -\textup{proj}_v(x) = x-y$.

The dot-product formula is useful for us because it allows us to compute the squared length of the projection by taking a dot product $ |x \cdot v|^2$. So then a formula for the distance of $ x$ from the line spanned by the unit vector $ v$ is

$ \displaystyle (\textup{dist}_v(x))^2 = \left ( \sum_{i=1}^n x_i^2 \right ) – |x \cdot v|^2$

This formula is just a restatement of the Pythagorean theorem for perpendicular vectors.

$ \displaystyle \sum_{i} x_i^2 = (\textup{proj}_v(x))^2 + (\textup{dist}_v(x))^2$

In particular, the difference vector we originally called $ z$ has squared length $ \textup{dist}_v(x)^2$. The vector $ y$, which is perpendicular to $ z$ and is also the projection of $ x$ onto $ L$, it’s squared length is $ (\textup{proj}_v(x))^2$. And the Pythagorean theorem tells us that summing those two squared lengths gives you the squared length of the hypotenuse $ x$.

If we were trying to find the best approximating 1-dimensional subspace for a set of data points $ X$, then we’d want to minimize the sum of the squared distances for every point $ x \in X$. Namely, we want the $ v$ that solves $ \min_{|v|=1} \sum_{x \in X} (\textup{dist}_v(x))^2$.

With some slight algebra we can make our life easier. The short version: minimizing the sum of squared distances is the same thing as maximizing the sum of squared lengths of the projections. The longer version: let’s go back to a single point $ x$ and the line spanned by $ v$. The Pythagorean theorem told us that

$ \displaystyle \sum_{i} x_i^2 = (\textup{proj}_v(x))^2 + (\textup{dist}_v(x))^2$

The squared length of $ x$ is constant. It’s an input to the algorithm and it doesn’t change through a run of the algorithm. So we get the squared distance by subtracting $ (\textup{proj}_v(x))^2$ from a constant number,

$ \displaystyle \sum_{i} x_i^2 – (\textup{proj}_v(x))^2 = (\textup{dist}_v(x))^2$

which means if we want to minimize the squared distance, we can instead maximize the squared projection. Maximizing the subtracted thing minimizes the whole expression.

It works the same way if you’re summing over all the data points in $ X$. In fact, we can say it much more compactly this way. If the rows of $ A$ are your data points, then $ Av$ contains as each entry the (signed) dot products $ x_i \cdot v$. And the squared norm of this vector, $ |Av|^2$, is exactly the sum of the squared lengths of the projections of the data onto the line spanned by $ v$. The last thing is that maximizing a square is the same as maximizing its square root, so we can switch freely between saying our objective is to find the unit vector $ v$ that maximizes $ |Av|$ and that which maximizes $ |Av|^2$.

At this point you should be thinking,

Great, we have written down an optimization problem: $ \max_{v : |v|=1} |Av|$. If we could solve this, we’d have the best 1-dimensional linear approximation to the data contained in the rows of $ A$. But (1) how do we solve that problem? And (2) you promised a $ k$-dimensional approximating subspace. I feel betrayed! Swindled! Bamboozled!

Here’s the fantastic thing. We can solve the 1-dimensional optimization problem efficiently (we’ll do it later in this post), and (2) is answered by the following theorem.

The SVD Theorem: Computing the best $ k$-dimensional subspace reduces to $ k$ applications of the one-dimensional problem.

We will prove this after we introduce the terms “singular value” and “singular vector.”

Singular values and vectors

As I just said, we can get the best $ k$-dimensional approximating linear subspace by solving the one-dimensional maximization problem $ k$ times. The singular vectors of $ A$ are defined recursively as the solutions to these sub-problems. That is, I’ll call $ v_1$ the first singular vector of $ A$, and it is:

$ \displaystyle v_1 = \arg \max_{v, |v|=1} |Av|$

And the corresponding first singular value, denoted $ \sigma_1(A)$, is the maximal value of the optimization objective, i.e. $ |Av_1|$. (I will use this term frequently, that $ |Av|$ is the “objective” of the optimization problem.) Informally speaking, $ (\sigma_1(A))^2$ represents how much of the data was captured by the first singular vector. Meaning, how close the vectors are to lying on the line spanned by $ v_1$. Larger values imply the approximation is better. In fact, if all the data points lie on a line, then $ (\sigma_1(A))^2$ is the sum of the squared norms of the rows of $ A$.

Now here is where we see the reduction from the $ k$-dimensional case to the 1-dimensional case. To find the best 2-dimensional subspace, you first find the best one-dimensional subspace (spanned by $ v_1$), and then find the best 1-dimensional subspace, but only considering those subspaces that are the spans of unit vectors perpendicular to $ v_1$. The notation for “vectors $ v$ perpendicular to $ v_1$” is $ v \perp v_1$. Restating, the second singular vector $ v _2$ is defined as

$ \displaystyle v_2 = \arg \max_{v \perp v_1, |v| = 1} |Av|$

And the SVD theorem implies the subspace spanned by $ \{ v_1, v_2 \}$ is the best 2-dimensional linear approximation to the data. Likewise $ \sigma_2(A) = |Av_2|$ is the second singular value. Its squared magnitude tells us how much of the data that was not “captured” by $ v_1$ is captured by $ v_2$. Again, if the data lies in a 2-dimensional subspace, then the span of $ \{ v_1, v_2 \}$ will be that subspace.

We can continue this process. Recursively define $ v_k$, the $ k$-th singular vector, to be the vector which maximizes $ |Av|$, when $ v$ is considered only among the unit vectors which are perpendicular to $ \textup{span} \{ v_1, \dots, v_{k-1} \}$. The corresponding singular value $ \sigma_k(A)$ is the value of the optimization problem.

As a side note, because of the way we defined the singular values as the objective values of “nested” optimization problems, the singular values are decreasing, $ \sigma_1(A) \geq \sigma_2(A) \geq \dots \geq \sigma_n(A) \geq 0$. This is obvious: you only pick $ v_2$ in the second optimization problem because you already picked $ v_1$ which gave a bigger singular value, so $ v_2$’s objective can’t be bigger.

If you keep doing this, one of two things happen. Either you reach $ v_n$ and since the domain is $ n$-dimensional there are no remaining vectors to choose from, the $ v_i$ are an orthonormal basis of $ \mathbb{R}^n$. This means that the data in $ A$ contains a full-rank submatrix. The data does not lie in any smaller-dimensional subspace. This is what you’d expect from real data.

Alternatively, you could get to a stage $ v_k$ with $ k < n$ and when you try to solve the optimization problem you find that every perpendicular $ v$ has $ Av = 0$. In this case, the data actually does lie in a $ k$-dimensional subspace, and the first-through-$ k$-th singular vectors you computed span this subspace.

Let’s do a quick sanity check: how do we know that the singular vectors $ v_i$ form a basis? Well formally they only span a basis of the row space of $ A$, i.e. a basis of the subspace spanned by the data contained in the rows of $ A$. But either way the point is that each $ v_{i+1}$ spans a new dimension from the previous $ v_1, \dots, v_i$ because we’re choosing $ v_{i+1}$ to be orthogonal to all the previous $ v_i$. So the answer to our sanity check is “by construction.”

Back to the singular vectors, the discussion from the last post tells us intuitively that the data is probably never in a small subspace.  You never expect the process of finding singular vectors to stop before step $ n$, and if it does you take a step back and ask if something deeper is going on. Instead, in real life you specify how much of the data you want to capture, and you keep computing singular vectors until you’ve passed the threshold. Alternatively, you specify the amount of computing resources you’d like to spend by fixing the number of singular vectors you’ll compute ahead of time, and settle for however good the $ k$-dimensional approximation is.

Before we get into any code or solve the 1-dimensional optimization problem, let’s prove the SVD theorem.

Proof of SVD theorem.

Recall we’re trying to prove that the first $ k$ singular vectors provide a linear subspace $ W$ which maximizes the squared-sum of the projections of the data onto $ W$. For $ k=1$ this is trivial, because we defined $ v_1$ to be the solution to that optimization problem. The case of $ k=2$ contains all the important features of the general inductive step. Let $ W$ be any best-approximating 2-dimensional linear subspace for the rows of $ A$. We’ll show that the subspace spanned by the two singular vectors $ v_1, v_2$ is at least as good (and hence equally good).

Let $ w_1, w_2$ be any orthonormal basis for $ W$ and let $ |Aw_1|^2 + |Aw_2|^2$ be the quantity that we’re trying to maximize (and which $ W$ maximizes by assumption). Moreover, we can pick the basis vector $ w_2$ to be perpendicular to $ v_1$. To prove this we consider two cases: either $ v_1$ is already perpendicular to $ W$ in which case it’s trivial, or else $ v_1$ isn’t perpendicular to $ W$ and you can choose $ w_1$ to be $ \textup{proj}_W(v_1)$ and choose $ w_2$ to be any unit vector perpendicular to $ w_1$.

Now since $ v_1$ maximizes $ |Av|$, we have $ |Av_1|^2 \geq |Aw_1|^2$. Moreover, since $ w_2$ is perpendicular to $ v_1$, the way we chose $ v_2$ also makes $ |Av_2|^2 \geq |Aw_2|^2$. Hence the objective $ |Av_1|^2 + |Av_2|^2 \geq |Aw_1|^2 + |Aw_2|^2$, as desired.

For the general case of $ k$, the inductive hypothesis tells us that the first $ k$ terms of the objective for $ k+1$ singular vectors is maximized, and we just have to pick any vector $ w_{k+1}$ that is perpendicular to all $ v_1, v_2, \dots, v_k$, and the rest of the proof is just like the 2-dimensional case.

$ \square$

Now remember that in the last post we started with the definition of the SVD as a decomposition of a matrix $ A = U\Sigma V^T$? And then we said that this is a certain kind of change of basis? Well the singular vectors $ v_i$ together form the columns of the matrix $ V$ (the rows of $ V^T$), and the corresponding singular values $ \sigma_i(A)$ are the diagonal entries of $ \Sigma$. When $ A$ is understood we’ll abbreviate the singular value as $ \sigma_i$.

To reiterate with the thoughts from last post, the process of applying $ A$ is exactly recovered by the process of first projecting onto the (full-rank space of) singular vectors $ v_1, \dots, v_k$, scaling each coordinate of that projection according to the corresponding singular values, and then applying this $ U$ thing we haven’t talked about yet.

So let’s determine what $ U$ has to be. The way we picked $ v_i$ to make $ A$ diagonal gives us an immediate suggestion: use the $ Av_i$ as the columns of $ U$. Indeed, define $ u_i = Av_i$, the images of the singular vectors under $ A$. We can swiftly show the $ u_i$ form a basis of the image of $ A$. The reason is because if $ v = \sum_i c_i v_i$ (using all $ n$ of the singular vectors $ v_i$), then by linearity $ Av = \sum_{i} c_i Av_i = \sum_i c_i u_i$. It is also easy to see why the $ u_i$ are orthogonal (prove it as an exercise). Let’s further make sure the $ u_i$ are unit vectors and redefine them as $ u_i = \frac{1}{\sigma_i}Av_i$

If you put these thoughts together, you can say exactly what $ A$ does to any given vector $ x$. Since the $ v_i$ form an orthonormal basis, $ x = \sum_i (x \cdot v_i) v_i$, and then applying $ A$ gives

$ \displaystyle \begin{aligned}Ax &= A \left ( \sum_i (x \cdot v_i) v_i \right ) \\  &= \sum_i (x \cdot v_i) A_i v_i \\ &= \sum_i (x \cdot v_i) \sigma_i u_i \end{aligned}$

If you’ve been closely reading this blog in the last few months, you’ll recognize a very nice way to write the last line of the above equation. It’s an outer product. So depending on your favorite symbols, you’d write this as either $ A = \sum_{i} \sigma_i u_i \otimes v_i$ or $ A = \sum_i \sigma_i u_i v_i^T$. Or, if you like expressing things as matrix factorizations, as $ A = U\Sigma V^T$. All three are describing the same object.

Let’s move on to some code.

A black box example

Before we implement SVD from scratch (an urge that commands me from the depths of my soul!), let’s see a black-box example that uses existing tools. For this we’ll use the numpy library.

Recall our movie-rating matrix from the last post:


The code to compute the svd of this matrix is as simple as it gets:

from numpy.linalg import svd

movieRatings = [
    [2, 5, 3],
    [1, 2, 1],
    [4, 1, 1],
    [3, 5, 2],
    [5, 3, 1],
    [4, 5, 5],
    [2, 4, 2],
    [2, 2, 5],

U, singularValues, V = svd(movieRatings)

Printing these values out gives

[[-0.39458526  0.23923575 -0.35445911 -0.38062172 -0.29836818 -0.49464816 -0.30703202 -0.29763321]
 [-0.15830232  0.03054913 -0.15299759 -0.45334816  0.31122898  0.23892035 -0.37313346  0.67223457]
 [-0.22155201 -0.52086121  0.39334917 -0.14974792 -0.65963979  0.00488292 -0.00783684  0.25934607]
 [-0.39692635 -0.08649009 -0.41052882  0.74387448 -0.10629499  0.01372565 -0.17959298  0.26333462]
 [-0.34630257 -0.64128825  0.07382859 -0.04494155  0.58000668 -0.25806239  0.00211823 -0.24154726]
 [-0.53347449  0.19168874  0.19949342 -0.03942604  0.00424495  0.68715732 -0.06957561 -0.40033035]
 [-0.31660464  0.06109826 -0.30599517 -0.19611823 -0.01334272  0.01446975  0.85185852  0.19463493]
 [-0.32840223  0.45970413  0.62354764  0.1783041   0.17631186 -0.39879476  0.06065902  0.25771578]]
[ 15.09626916   4.30056855   3.40701739]
[[-0.54184808 -0.67070995 -0.50650649]
 [-0.75152295  0.11680911  0.64928336]
 [ 0.37631623 -0.73246419  0.56734672]]

Now this is a bit weird, because the matrices $ U, V$ are the wrong shape! Remember, there are only supposed to be three vectors since the input matrix has rank three. So what gives? This is a distinction that goes by the name “full” versus “reduced” SVD. The idea goes back to our original statement that $ U \Sigma V^T$ is a decomposition with $ U, V^T$ both orthogonal and square matrices. But in the derivation we did in the last section, the $ U$ and $ V$ were not square. The singular vectors $ v_i$ could potentially stop before even becoming full rank.

In order to get to square matrices, what people sometimes do is take the two bases $ v_1, \dots, v_k$ and $ u_1, \dots, u_k$ and arbitrarily choose ways to complete them to a full orthonormal basis of their respective vector spaces. In other words, they just make the matrix square by filling it with data for no reason other than that it’s sometimes nice to have a complete basis. We don’t care about this. To be honest, I think the only place this comes in useful is in the desire to be particularly tidy in a mathematical formulation of something.

We can still work with it programmatically. By fudging around a bit with numpy’s shapes to get a diagonal matrix, we can reconstruct the input rating matrix from the factors.

Sigma = np.vstack([
    np.zeros((5, 3)),

print(np.round(movieRatings -,, V)), decimals=10))

And the output is, as one expects, a matrix of all zeros. Meaning that we decomposed the movie rating matrix, and built it back up from the factors.

We can actually get the SVD as we defined it (with rectangular matrices) by passing a special flag to numpy’s svd.

U, singularValues, V = svd(movieRatings, full_matrices=False)

Sigma = np.diag(singularValues)
print(np.round(movieRatings -,, V)), decimals=10))

And the result

[[-0.39458526  0.23923575 -0.35445911]
 [-0.15830232  0.03054913 -0.15299759]
 [-0.22155201 -0.52086121  0.39334917]
 [-0.39692635 -0.08649009 -0.41052882]
 [-0.34630257 -0.64128825  0.07382859]
 [-0.53347449  0.19168874  0.19949342]
 [-0.31660464  0.06109826 -0.30599517]
 [-0.32840223  0.45970413  0.62354764]]
[ 15.09626916   4.30056855   3.40701739]
[[-0.54184808 -0.67070995 -0.50650649]
 [-0.75152295  0.11680911  0.64928336]
 [ 0.37631623 -0.73246419  0.56734672]]
[[-0. -0. -0.]
 [-0. -0.  0.]
 [ 0. -0.  0.]
 [-0. -0. -0.]
 [-0. -0. -0.]
 [-0. -0. -0.]
 [-0. -0. -0.]
 [ 0. -0. -0.]]

This makes the reconstruction less messy, since we can just multiply everything without having to add extra rows of zeros to $ \Sigma$.

What do the singular vectors and values tell us about the movie rating matrix? (Besides nothing, since it’s a contrived example) You’ll notice that the first singular vector $ \sigma_1 > 15$ while the other two singular values are around $ 4$. This tells us that the first singular vector covers a large part of the structure of the matrix. I.e., a rank-1 matrix would be a pretty good approximation to the whole thing. As an exercise to the reader, write a program that evaluates this claim (how good is “good”?).

The greedy optimization routine

Now we’re going to write SVD from scratch. We’ll first implement the greedy algorithm for the 1-d optimization problem, and then we’ll perform the inductive step to get a full algorithm. Then we’ll run it on the CNN data set.

The method we’ll use to solve the 1-dimensional problem isn’t necessarily industry strength (see this document for a hint of what industry strength looks like), but it is simple conceptually. It’s called the power method. Now that we have our decomposition of theorem, understanding how the power method works is quite easy.

Let’s work in the language of a matrix decomposition $ A = U \Sigma V^T$, more for practice with that language than anything else (using outer products would give us the same result with slightly different computations). Then let’s observe $ A^T A$, wherein we’ll use the fact that $ U$ is orthonormal and so $ U^TU$ is the identity matrix:

$ \displaystyle A^TA = (U \Sigma V^T)^T(U \Sigma V^T) = V \Sigma U^TU \Sigma V^T = V \Sigma^2 V^T$

So we can completely eliminate $ U$ from the discussion, and look at just $ V \Sigma^2 V^T$. And what’s nice about this matrix is that we can compute its eigenvectors, and eigenvectors turn out to be exactly the singular vectors. The corresponding eigenvalues are the squared singular values. This should be clear from the above derivation. If you apply $ (V \Sigma^2 V^T)$ to any $ v_i$, the only parts of the product that aren’t zero are the ones involving $ v_i$ with itself, and the scalar $ \sigma_i^2$ factors in smoothly. It’s dead simple to check.

Theorem: Let $ x$ be a random unit vector and let $ B = A^TA = V \Sigma^2 V^T$. Then with high probability, $ \lim_{s \to \infty} B^s x$ is in the span of the first singular vector $ v_1$. If we normalize $ B^s x$ to a unit vector at each $ s$, then furthermore the limit is $ v_1$.

Proof. Start with a random unit vector $ x$, and write it in terms of the singular vectors $ x = \sum_i c_i v_i$. That means $ Bx = \sum_i c_i \sigma_i^2 v_i$. If you recursively apply this logic, you get $ B^s x = \sum_i c_i \sigma_i^{2s} v_i$. In particular, the dot product of $ (B^s x)$ with any $ v_j$ is $ c_i \sigma_j^{2s}$.

What this means is that so long as the first singular value $ \sigma_1$ is sufficiently larger than the second one $ \sigma_2$, and in turn all the other singular values, the part of $ B^s x$  corresponding to $ v_1$ will be much larger than the rest. Recall that if you expand a vector in terms of an orthonormal basis, in this case $ B^s x$ expanded in the $ v_i$, the coefficient of $ B^s x$ on $ v_j$ is exactly the dot product. So to say that $ B^sx$ converges to being in the span of $ v_1$ is the same as saying that the ratio of these coefficients, $ |(B^s x \cdot v_1)| / |(B^s x \cdot v_j)| \to \infty$ for any $ j$. In other words, the coefficient corresponding to the first singular vector dominates all of the others. And so if we normalize, the coefficient of $ B^s x$ corresponding to $ v_1$ tends to 1, while the rest tend to zero.

Indeed, this ratio is just $ (\sigma_1 / \sigma_j)^{2s}$ and the base of this exponential is bigger than 1.

$ \square$

If you want to be a little more precise and find bounds on the number of iterations required to converge, you can. The worry is that your random starting vector is “too close” to one of the smaller singular vectors $ v_j$, so that if the ratio of $ \sigma_1 / \sigma_j$ is small, then the “pull” of $ v_1$ won’t outweigh the pull of $ v_j$ fast enough. Choosing a random unit vector allows you to ensure with high probability that this doesn’t happen. And conditioned on it not happening (or measuring “how far the event is from happening” precisely), you can compute a precise number of iterations required to converge. The last two pages of these lecture notes have all the details.

We won’t compute a precise number of iterations. Instead we’ll just compute until the angle between $ B^{s+1}x$ and $ B^s x$ is very small. Here’s the algorithm

import numpy as np
from numpy.linalg import norm

from random import normalvariate
from math import sqrt

def randomUnitVector(n):
    unnormalized = [normalvariate(0, 1) for _ in range(n)]
    theNorm = sqrt(sum(x * x for x in unnormalized))
    return [x / theNorm for x in unnormalized]

def svd_1d(A, epsilon=1e-10):
    ''' The one-dimensional SVD '''

    n, m = A.shape
    x = randomUnitVector(m)
    lastV = None
    currentV = x
    B =, A)

    iterations = 0
    while True:
        iterations += 1
        lastV = currentV
        currentV =, lastV)
        currentV = currentV / norm(currentV)

        if abs(, lastV)) &gt; 1 - epsilon:
            print(&quot;converged in {} iterations!&quot;.format(iterations))
            return currentV

We start with a random unit vector $ x$, and then loop computing $ x_{t+1} = Bx_t$, renormalizing at each step. The condition for stopping is that the magnitude of the dot product between $ x_t$ and $ x_{t+1}$ (since they’re unit vectors, this is the cosine of the angle between them) is very close to 1.

And using it on our movie ratings example:

if __name__ == &quot;__main__&quot;:
    movieRatings = np.array([
        [2, 5, 3],
        [1, 2, 1],
        [4, 1, 1],
        [3, 5, 2],
        [5, 3, 1],
        [4, 5, 5],
        [2, 4, 2],
        [2, 2, 5],
    ], dtype='float64')


With the result

converged in 6 iterations!
[-0.54184805 -0.67070993 -0.50650655]

Note that the sign of the vector may be different from numpy’s output because we start with a random vector to begin with.

The recursive step, getting from $ v_1$ to the entire SVD, is equally straightforward. Say you start with the matrix $ A$ and you compute $ v_1$. You can use $ v_1$ to compute $ u_1$ and $ \sigma_1(A)$. Then you want to ensure you’re ignoring all vectors in the span of $ v_1$ for your next greedy optimization, and to do this you can simply subtract the rank 1 component of $ A$ corresponding to $ v_1$. I.e., set $ A’ = A – \sigma_1(A) u_1 v_1^T$. Then it’s easy to see that $ \sigma_1(A’) = \sigma_2(A)$ and basically all the singular vectors shift indices by 1 when going from $ A$ to $ A’$. Then you repeat.

If that’s not clear enough, here’s the code.

def svd(A, epsilon=1e-10):
    n, m = A.shape
    svdSoFar = []

    for i in range(m):
        matrixFor1D = A.copy()

        for singularValue, u, v in svdSoFar[:i]:
            matrixFor1D -= singularValue * np.outer(u, v)

        v = svd_1d(matrixFor1D, epsilon=epsilon)  # next singular vector
        u_unnormalized =, v)
        sigma = norm(u_unnormalized)  # next singular value
        u = u_unnormalized / sigma

        svdSoFar.append((sigma, u, v))

    # transform it into matrices of the right shape
    singularValues, us, vs = [np.array(x) for x in zip(*svdSoFar)]

    return singularValues, us.T, vs

And we can run this on our movie rating matrix to get the following

&gt;&gt;&gt; theSVD = svd(movieRatings)
&gt;&gt;&gt; theSVD[0]
array([ 15.09626916,   4.30056855,   3.40701739])
&gt;&gt;&gt; theSVD[1]
array([[ 0.39458528, -0.23923093,  0.35446407],
       [ 0.15830233, -0.03054705,  0.15299815],
       [ 0.221552  ,  0.52085578, -0.39336072],
       [ 0.39692636,  0.08649568,  0.41052666],
       [ 0.34630257,  0.64128719, -0.07384286],
       [ 0.53347448, -0.19169154, -0.19948959],
       [ 0.31660465, -0.0610941 ,  0.30599629],
       [ 0.32840221, -0.45971273, -0.62353781]])
&gt;&gt;&gt; theSVD[2]
array([[ 0.54184805,  0.67071006,  0.50650638],
       [ 0.75151641, -0.11679644, -0.64929321],
       [-0.37632934,  0.73246611, -0.56733554]])

Checking this against our numpy output shows it’s within a reasonable level of precision (considering the power method took on the order of ten iterations!)

&gt;&gt;&gt; np.round(np.abs(npSVD[0]) - np.abs(theSVD[1]), decimals=5)
array([[ -0.00000000e+00,  -0.00000000e+00,   0.00000000e+00],
       [  0.00000000e+00,  -0.00000000e+00,   0.00000000e+00],
       [  0.00000000e+00,  -1.00000000e-05,   1.00000000e-05],
       [  0.00000000e+00,   0.00000000e+00,  -0.00000000e+00],
       [  0.00000000e+00,  -0.00000000e+00,   1.00000000e-05],
       [ -0.00000000e+00,   0.00000000e+00,  -0.00000000e+00],
       [  0.00000000e+00,  -0.00000000e+00,   0.00000000e+00],
       [ -0.00000000e+00,   1.00000000e-05,  -1.00000000e-05]])
&gt;&gt;&gt; np.round(np.abs(npSVD[2]) - np.abs(theSVD[2]), decimals=5)
array([[  0.00000000e+00,   0.00000000e+00,  -0.00000000e+00],
       [ -1.00000000e-05,  -1.00000000e-05,   1.00000000e-05],
       [  1.00000000e-05,   0.00000000e+00,  -1.00000000e-05]])
&gt;&gt;&gt; np.round(np.abs(npSVD[1]) - np.abs(theSVD[0]), decimals=5)
array([ 0.,  0., -0.])

So there we have it. We added an extra little bit to the svd function, an argument $ k$ which stops computing the svd after it reaches rank $ k$.

CNN stories

One interesting use of the SVD is in topic modeling. Topic modeling is the process of taking a bunch of documents (news stories, or emails, or movie scripts, whatever) and grouping them by topic, where the algorithm gets to choose what counts as a “topic.” Topic modeling is just the name that natural language processing folks use instead of clustering.

The SVD can help one model topics as follows. First you construct a matrix $ A$ called a document-term matrix whose rows correspond to words in some fixed dictionary and whose columns correspond to documents. The $ (i,j)$ entry of $ A$ contains the number of times word $ i$ shows up in document $ j$. Or, more precisely, some quantity derived from that count, like a normalized count. See this table on wikipedia for a list of options related to that. We’ll just pick one arbitrarily for use in this post.

The point isn’t how we normalize the data, but what the SVD of $ A = U \Sigma V^T$ means in this context. Recall that the domain of $ A$, as a linear map, is a vector space whose dimension is the number of stories. We think of the vectors in this space as documents, or rather as an “embedding” of the abstract concept of a document using the counts of how often each word shows up in a document as a proxy for the semantic meaning of the document. Likewise, the codomain is the space of all words, and each word is embedded by which documents it occurs in. If we compare this to the movie rating example, it’s the same thing: a movie is the vector of ratings it receives from people, and a person is the vector of ratings of various movies.

Say you take a rank 3 approximation to $ A$. Then you get three singular vectors $ v_1, v_2, v_3$ which form a basis for a subspace of words, i.e., the “idealized” words. These idealized words are your topics, and you can compute where a “new word” falls by looking at which documents it appears in (writing it as a vector in the domain) and saying its “topic” is the closest of the $ v_1, v_2, v_3$. The same process applies to new documents. You can use this to cluster existing documents as well.

The dataset we’ll use for this post is a relatively small corpus of a thousand CNN stories picked from 2012. Here’s an excerpt from one of them

$ cat data/cnn-stories/story479.txt
3 things to watch on Super Tuesday
Here are three things to watch for: Romney's big day. He's been the off-and-on frontrunner throughout the race, but a big Super Tuesday could begin an end game toward a sometimes hesitant base coalescing behind former Massachusetts Gov. Mitt Romney. Romney should win his home state of Massachusetts, neighboring Vermont and Virginia, ...

So let’s first build this document-term matrix, with the normalized values, and then we’ll compute it’s SVD and see what the topics look like.

Step 1 is cleaning the data. We used a bunch of routines from the nltk library that boils down to this loop:

    for filename, documentText in documentDict.items():
        tokens = tokenize(documentText)
        tagged_tokens = pos_tag(tokens)
        wnl = WordNetLemmatizer()
        stemmedTokens = [wnl.lemmatize(word, wordnetPos(tag)).lower()
                         for word, tag in tagged_tokens]

This turns the Super Tuesday story into a list of words (with repetition):

[&quot;thing&quot;, &quot;watch&quot;, &quot;three&quot;, &quot;thing&quot;, &quot;watch&quot;, &quot;big&quot;, ... ]

If you’ll notice the name Romney doesn’t show up in the list of words. I’m only keeping the words that show up in the top 100,000 most common English words, and then lemmatizing all of the words to their roots. It’s not a perfect data cleaning job, but it’s simple and good enough for our purposes.

Now we can create the document term matrix.

def makeDocumentTermMatrix(data):
    words = allWords(data)  # get the set of all unique words

    wordToIndex = dict((word, i) for i, word in enumerate(words))
    indexToWord = dict(enumerate(words))
    indexToDocument = dict(enumerate(data))

    matrix = np.zeros((len(words), len(data)))
    for docID, document in enumerate(data):
        docWords = Counter(document['words'])
        for word, count in docWords.items():
            matrix[wordToIndex[word], docID] = count

    return matrix, (indexToWord, indexToDocument)

This creates a matrix with the raw integer counts. But what we need is a normalized count. The idea is that a common word like “thing” shows up disproportionately more often than “election,” and we don’t want raw magnitude of a word count to outweigh its semantic contribution to the classification. This is the applied math part of the algorithm design. So what we’ll do (and this technique together with SVD is called latent semantic indexing) is normalize each entry so that it measures both the frequency of a term in a document and the relative frequency of a term compared to the global frequency of that term. There are many ways to do this, and we’ll just pick one. See the github repository if you’re interested.

So now lets compute a rank 10 decomposition and see how to cluster the results.

    data = load()
    matrix, (indexToWord, indexToDocument) = makeDocumentTermMatrix(data)
    matrix = normalize(matrix)
    sigma, U, V = svd(matrix, k=10)

This uses our svd, not numpy’s. Though numpy’s routine is much faster, it’s fun to see things work with code written from scratch. The result is too large to display here, but I can report the singular values.

&gt;&gt;&gt; sigma
array([ 42.85249098,  21.85641975,  19.15989197,  16.2403354 ,
        15.40456779,  14.3172779 ,  13.47860033,  13.23795002,
        12.98866537,  12.51307445])

Now we take our original inputs and project them onto the subspace spanned by the singular vectors. This is the part that represents each word (resp., document) in terms of the idealized words (resp., documents), the singular vectors. Then we can apply a simple k-means clustering algorithm to the result, and observe the resulting clusters as documents.

    projectedDocuments =, U)
    projectedWords =, V.T)

    documentCenters, documentClustering = cluster(projectedDocuments)
    wordCenters, wordClustering = cluster(projectedWords)

    wordClusters = [
        [indexToWord[i] for (i, x) in enumerate(wordClustering) if x == j]
        for j in range(len(set(wordClustering)))

    documentClusters = [
         for (i, x) in enumerate(documentClustering) if x == j]
        for j in range(len(set(documentClustering)))

And now we can inspect individual clusters. Right off the bat we can tell the clusters aren’t quite right simply by looking at the sizes of each cluster.

&gt;&gt;&gt; Counter(wordClustering)
Counter({1: 9689, 2: 1051, 8: 680, 5: 557, 3: 321, 7: 225, 4: 174, 6: 124, 9: 123})
&gt;&gt;&gt; Counter(documentClustering)
Counter({7: 407, 6: 109, 0: 102, 5: 87, 9: 85, 2: 65, 8: 55, 4: 47, 3: 23, 1: 15})

What looks wrong to me is the size of the largest word cluster. If we could group words by topic, then this is saying there’s a topic with over nine thousand words associated with it! Inspecting it even closer, it includes words like “vegan,” “skunk,” and “pope.” On the other hand, some word clusters are spot on. Examine, for example, the fifth cluster which includes words very clearly associated with crime stories.

&gt;&gt;&gt; wordClusters[4]
['account', 'accuse', 'act', 'affiliate', 'allegation', 'allege', 'altercation', 'anything', 'apartment', 'arrest', 'arrive', 'assault', 'attorney', 'authority', 'bag', 'black', 'blood', 'boy', 'brother', 'bullet', 'candy', 'car', 'carry', 'case', 'charge', 'chief', 'child', 'claim', 'client', 'commit', 'community', 'contact', 'convenience', 'court', 'crime', 'criminal', 'cry', 'dead', 'deadly', 'death', 'defense', 'department', 'describe', 'detail', 'determine', 'dispatcher', 'district', 'document', 'enforcement', 'evidence', 'extremely', 'family', 'father', 'fear', 'fiancee', 'file', 'five', 'foot', 'friend', 'front', 'gate', 'girl', 'girlfriend', 'grand', 'ground', 'guilty', 'gun', 'gunman', 'gunshot', 'hand', 'happen', 'harm', 'head', 'hear', 'heard', 'hoodie', 'hour', 'house', 'identify', 'immediately', 'incident', 'information', 'injury', 'investigate', 'investigation', 'investigator', 'involve', 'judge', 'jury', 'justice', 'kid', 'killing', 'lawyer', 'legal', 'letter', 'life', 'local', 'man', 'men', 'mile', 'morning', 'mother', 'murder', 'near', 'nearby', 'neighbor', 'newspaper', 'night', 'nothing', 'office', 'officer', 'online', 'outside', 'parent', 'person', 'phone', 'police', 'post', 'prison', 'profile', 'prosecute', 'prosecution', 'prosecutor', 'pull', 'racial', 'racist', 'release', 'responsible', 'return', 'review', 'role', 'saw', 'scene', 'school', 'scream', 'search', 'sentence', 'serve', 'several', 'shoot', 'shooter', 'shooting', 'shot', 'slur', 'someone', 'son', 'sound', 'spark', 'speak', 'staff', 'stand', 'store', 'story', 'student', 'surveillance', 'suspect', 'suspicious', 'tape', 'teacher', 'teen', 'teenager', 'told', 'tragedy', 'trial', 'vehicle', 'victim', 'video', 'walk', 'watch', 'wear', 'whether', 'white', 'witness', 'young']

As sad as it makes me to see that ‘black’ and ‘slur’ and ‘racial’ appear in this category, it’s a reminder that naively using the output of a machine learning algorithm can perpetuate racism.

Here’s another interesting cluster corresponding to economic words:

&gt;&gt;&gt; wordClusters[6]
['agreement', 'aide', 'analyst', 'approval', 'approve', 'austerity', 'average', 'bailout', 'beneficiary', 'benefit', 'bill', 'billion', 'break', 'broadband', 'budget', 'class', 'combine', 'committee', 'compromise', 'conference', 'congressional', 'contribution', 'core', 'cost', 'currently', 'cut', 'deal', 'debt', 'defender', 'deficit', 'doc', 'drop', 'economic', 'economy', 'employee', 'employer', 'erode', 'eurozone', 'expire', 'extend', 'extension', 'fee', 'finance', 'fiscal', 'fix', 'fully', 'fund', 'funding', 'game', 'generally', 'gleefully', 'growth', 'hamper', 'highlight', 'hike', 'hire', 'holiday', 'increase', 'indifferent', 'insistence', 'insurance', 'job', 'juncture', 'latter', 'legislation', 'loser', 'low', 'lower', 'majority', 'maximum', 'measure', 'middle', 'negotiation', 'offset', 'oppose', 'package', 'pass', 'patient', 'pay', 'payment', 'payroll', 'pension', 'plight', 'portray', 'priority', 'proposal', 'provision', 'rate', 'recession', 'recovery', 'reduce', 'reduction', 'reluctance', 'repercussion', 'rest', 'revenue', 'rich', 'roughly', 'sale', 'saving', 'scientist', 'separate', 'sharp', 'showdown', 'sign', 'specialist', 'spectrum', 'spending', 'strength', 'tax', 'tea', 'tentative', 'term', 'test', 'top', 'trillion', 'turnaround', 'unemployed', 'unemployment', 'union', 'wage', 'welfare', 'worker', 'worth']

One can also inspect the stories, though the clusters are harder to print out here. Interestingly the first cluster of documents are stories exclusively about Trayvon Martin. The second cluster is mostly international military conflicts. The third cluster also appears to be about international conflict, but what distinguishes it from the first cluster is that every story in the second cluster discusses Syria.

&gt;&gt;&gt; len([x for x in documentClusters[1] if 'Syria' in x]) / len(documentClusters[1])
&gt;&gt;&gt; len([x for x in documentClusters[2] if 'Syria' in x]) / len(documentClusters[2])

Anyway, you can explore the data more at your leisure (and tinker with the parameters to improve it!).

Issues with the power method

Though I mentioned that the power method isn’t an industry strength algorithm I didn’t say why. Let’s revisit that before we finish. The problem is that the convergence rate of even the 1-dimensional problem depends on the ratio of the first and second singular values, $ \sigma_1 / \sigma_2$. If that ratio is very close to 1, then the convergence will take a long time and need many many matrix-vector multiplications.

One way to alleviate that is to do the trick where, to compute a large power of a matrix, you iteratively square $ B$. But that requires computing a matrix square (instead of a bunch of matrix-vector products), and that requires a lot of time and memory if the matrix isn’t sparse. When the matrix is sparse, you can actually do the power method quite quickly, from what I’ve heard and read.

But nevertheless, the industry standard methods involve computing a particular matrix decomposition that is not only faster than the power method, but also numerically stable. That means that the algorithm’s runtime and accuracy doesn’t depend on slight changes in the entries of the input matrix. Indeed, you can have two matrices where $ \sigma_1 / \sigma_2$ is very close to 1, but changing a single entry will make that ratio much larger. The power method depends on this, so it’s not numerically stable. But the industry standard technique is not. This technique involves something called Householder reflections. So while the power method was great for a proof of concept, there’s much more work to do if you want true SVD power.

Until next time!

When Greedy Algorithms are Good Enough: Submodularity and the (1 – 1/e)-Approximation

Greedy algorithms are among the simplest and most intuitive algorithms known to humans. Their name essentially gives their description: do the thing that looks best right now, and repeat until nothing looks good anymore or you’re forced to stop. Some of the best situations in computer science are also when greedy algorithms are optimal or near-optimal. There is a beautiful theory of this situation, known as the theory of matroids. We haven’t covered matroids on this blog (edit: we did), but in this post we will focus on the next best thing: when the greedy algorithm guarantees a reasonably good approximation to the optimal solution.

This situation isn’t hard to formalize, and we’ll make it as abstract as possible. Say you have a set of objects $ X$, and you’re looking to find the “best” subset $ S \subset X$. Here “best” is just measured by a fixed (known, efficiently computable) objective function $ f : 2^X \to \mathbb{R}$. That is, $ f$ accepts as input subsets of $ X$ and outputs numbers so that better subsets have larger numbers. Then the goal is to find a subset maximizing $ f$.

In this generality the problem is clearly impossible. You’d have to check all subsets to be sure you didn’t miss the best one. So what conditions do we need on either $ X$ or $ f$ or both that makes this problem tractable? There are plenty you could try, but one very rich property is submodularity.

The Submodularity Condition

I think the simplest way to explain submodularity is in terms of coverage. Say you’re starting a new radio show and you have to choose which radio stations to broadcast from to reach the largest number of listeners. For simplicity say each radio station has one tower it broadcasts from, and you have a good estimate of the number of listeners you would reach if you broadcast from a given tower. For more simplicity, say it costs the same to broadcast from each tower, and your budget restricts you to a maximum of ten stations to broadcast from. So the question is: how do you pick towers to maximize your overall reach?

The hidden condition here is that some towers overlap in which listeners they reach. So if you broadcast from two towers in the same city, a listener who has access to both will just pick one or the other. In other words, there’s a diminished benefit to picking two overlapping towers if you already have chosen one.

In our version of the problem, picking both of these towers has some small amount of "overkill."

In our version of the problem, picking both of these towers has some small amount of “overkill.”

This “diminishing returns” condition is a general idea you can impose on any function that takes in subsets of a given set and produces numbers. If $ X$ is a set, then for a strange reason we denote by $ 2^X$ the set of all subsets of $ X$. So we can state this condition more formally,

Definition: Let $ X$ be a finite set. A function $ f: 2^X \to \mathbb{R}$ is called submodular if for all subsets $ S \subset T \subset X$ and all $ x \in X \setminus T$,

$ \displaystyle f(S \cup \{ x \}) – f(S) \geq f(T \cup \{ x \}) – f(T)$

In other words, if $ f$ measures “benefit,” then the marginal benefit of adding $ x$ to $ S$ is at least as high as the marginal benefit of adding it to $ T$. Since $ S \subset T$ and $ x$ are all arbitrary, this is as general as one could possibly make it: adding $ x$ to a bigger set can’t be better than adding it to a smaller set.

Before we start doing things with submodular functions, let’s explore some basic properties. The first is an equivalent definition of submodularity

Proposition: $ f$ is submodular if and only if for all $ A, B \subset X$, it holds that

$ \displaystyle f(A \cap B) + f(A \cup B) \leq f(A) + f(B)$.

Proof. If we assume $ f$ has the condition from this proposition, then we can set $ A=T, B=S \cup \{ x \}$, and the formula just works out. Conversely, if we have the condition from the definition, then using the fact that $ A \cap B \subset B$ we can inductively apply the inequality to each element of $ A \setminus B$ to get

$ \displaystyle f(A \cup B) – f(B) \leq f(A) – f(A \cap B)$

$ \square$

Next, we can tweak and combine submodular functions to get more submodular functions. In particular, non-negative linear combinations of sub-modular functions are submodular. In other words, if $ f_1, \dots, f_k$ are submodular on the same set $ X$, and $ \alpha_1, \dots, \alpha_k$ are all non-negative reals, then $ \alpha_1 f_1 + \dots + \alpha_k f_k$ is also a submodular function on $ X$. It’s an easy exercise in applying the definition to see why this is true. This is important because when we’re designing objectives to maximize, we can design them by making some simple submodular pieces, and then picking an appropriate combination of those pieces.

The second property we need to impose on a submodular function is monotonicity. That is, as your sets get more elements added to them, their value under $ f$ only goes up. In other words, $ f$ is monotone when $ S \subset T$ then $ f(S) \leq f(T)$. An interesting property of functions that are both submodular and monotone is that the truncation of such a function is also submodular and monotone. In other words, $ \textup{min}(f(S), c)$ is still submodular when $ f$ is monotone submodular and $ c$ is a constant.

Submodularity and Monotonicity Give 1 – 1/e

The wonderful thing about submodular functions is that we have a lot of great algorithmic guarantees for working with them. We’ll prove right now that the coverage problem (while it might be hard to solve in general) can be approximated pretty well by the greedy algorithm.

Here’s the algorithmic setup. I give you a finite set $ X$ and an efficient black-box to evaluate $ f(S)$ for any subset $ S \subset X$ you want. I promise you that $ f$ is monotone and submodular. Now I give you an integer $ k$ between 1 and the size of $ X$, and your task is to quickly find a set $ S$ of size $ k$ for which $ f(S)$ is maximal among all subsets of size $ k$. That is, you design an algorithm that will work for any $ k, X, f$ and runs in polynomial time in the sizes of $ X, k$.

In general this problem is NP-hard, meaning you’re not going to find a solution that works in the worst case (if you do, don’t call me; just claim your million dollar prize). So how well can we approximate the optimal value for $ f(S)$ by a different set of size $ k$? The beauty is that, if your function is monotone and submodular, you can guarantee to get within 63% of the optimum. The hope (and reality) is that in practice it will often perform much better, but still this is pretty good! More formally,

Theorem: Let $ f$ be a monotone, submodular, non-negative function on $ X$. The greedy algorithm, which starts with $ S$ as the empty set and at every step picks an element $ x$ which maximizes the marginal benefit $ f(S \cup \{ x \}) – f(S)$, provides a set $ S$ that achieves a $ (1- 1/e)$-approximation of the optimum.

We’ll prove this in just a little bit more generality, and the generality is quite useful. If we call $ S_1, S_2, \dots, S_l$ the sets chosen by the greedy algorithm (where now we might run the greedy algorithm for $ l > k$ steps), then for all $ l, k$, we have

$ \displaystyle f(S_l) \geq \left ( 1 – e^{-l/k} \right ) \max_{T: |T| \leq k} f(T)$

This allows us to run the algorithm for more than $ k$ steps to get a better approximation by sets of larger size, and quantify how much better the guarantee on that approximation would be. It’s like an algorithmic way of hedging your risk. So let’s prove it.

Proof. Let’s set up some notation first. Fix your $ l$ and $ k$, call $ S_i$ the set chosen by the greedy algorithm at step $ i$, and call $ S^*$ the optimal subset of size $ k$. Further call $ \textup{OPT}$ the value of the best set $ f(S^*)$. Call $ x_1^*, \dots, x_k^*$ the elements of $ S^*$ (the order is irrelevant). Now for every $ i < l$ monotonicity gives us $ f(S^*) \leq f(S^* \cup S_i)$. We can unravel this into a sum of marginal gains of adding single elements. The first step is

$ \displaystyle f(S^* \cup S_i) = f(S^* \cup S_i) – f(\{ x_1^*, \dots, x_{k-1}^* \} \cup S_i) + f(\{ x_1^*, \dots, x_{k-1}^* \} \cup S_i)$

The second step removes $ x_{k-1}^*$, from the last term, the third removes $ x_{k-2}^*$, and so on until we have removed all of $ S^*$ and get this sum

$ \displaystyle f(S^* \cup S_i) = f(S_i) + \sum_{j=1}^k \left ( f(S_i \cup \{ x_1^*, \dots, x_j^* \}) – f(S_i \cup \{ x_1^*, \dots, x_{j-1}^* \} ) \right )$

Now, applying submodularity, we can change all of these marginal benefits of “adding one more $ S^*$ element to $ S_i$ already with some $ S^*$ stuff” to “adding one more $ S^*$ element to just $ S_i$.” In symbols, the equation above is at most

$ \displaystyle f(S_i) + \sum_{x \in S^*} f(S_i \cup \{ x \}) – f(S_i)$

and because $ S_{i+1}$ is greedily chosen to maximize the benefit of adding a single element, so the above is at most

$ \displaystyle f(S_i) + \sum_{x \in S^*} f(S_{i+1}) – f(S_i) = f(S_i) + k(f(S_{i+1}) – f(S_i))$

Chaining all of these together, we have $ f(S^*) – f(S_i) \leq k(f(S_{i+1}) – f(S_i))$. If we call $ a_{i} = f(S^*) – f(S_i)$, then this inequality can be rewritten as $ a_{i+1} \leq (1 – 1/k) a_{i}$. Now by induction we can relate $ a_l \leq (1 – 1/k)^l a_0$. Now use the fact that $ a_0 \leq f(S^*)$ and the common inequality $ 1-x \leq e^{-x}$ to get

$ \displaystyle a_l = f(S^*) – f(S_l) \leq e^{-l/k} f(S^*)$

And rearranging gives $ f(S_l) \geq (1 – e^{-l/k}) f(S^*)$.

$ \square$

Setting $ l=k$ gives the approximation bound we promised. But note that allowing the greedy algorithm to run longer can give much stronger guarantees, though it requires you to sacrifice the cardinality constraint. $ 1 – 1/e$ is about 63%, but doubling the size of $ S$ gives about an 86% approximation guarantee. This is great for people in the real world, because you can quantify the gains you’d get by relaxing the constraints imposed on you (which are rarely set in stone).

So this is really great! We have quantifiable guarantees on a stupidly simple algorithm, and the setting is super general. And so if you have your problem and you manage to prove your function is submodular (this is often the hardest part), then you are likely to get this nice guarantee.

Extensions and Variations

This result on monotone submodular functions is just one part of a vast literature on finding approximation algorithms for submodular functions in various settings. In closing this post we’ll survey some of the highlights and provide references.

What we did in this post was maximize a monotone submodular function subject to a cardinality constraint $ |S| \leq k$. There are three basic variations we could do: we could drop constraints and see whether we can still get guarantees, we could look at minimization instead of maximization, and we could modify the kinds of constraints we impose on the solution.

There are a ton of different kinds of constraints, and we’ll discuss two. The first is where you need to get a certain value $ f(S) \geq q$, and you want to find the smallest set that achieves this value. Laurence Wolsey (who proved a lot of these theorems) showed in 1982 that a slight variant of the greedy algorithm can achieve a set whose size is a multiplicative factor of $ 1 + \log (\max_x f(\{ x \}))$ worse than the optimum.

The second kind of constraint is a generalization of a cardinality constraint called a knapsack constraint. This means that each item $ x \in X$ has a cost, and you have a finite budget with which to spend on elements you add to $ S$. One might expect this natural extension of the greedy algorithm to work: pick the element which maximizes the ratio of increasing the value of $ f$ to the cost (within your available budget). Unfortunately this algorithm can perform arbitrarily poorly, but there are two fun caveats. The first is that if you do both this augmented greedy algorithm and the greedy algorithm that ignores costs, then at least one of these can’t do too poorly. Specifically, one of them has to get at least a 30% approximation. This was shown by Leskovec et al in 2007. The second is that if you’re willing to spend more time in your greedy step by choosing the best subset of size 3, then you can get back to the $ 1-1/e$ approximation. This was shown by Sviridenko in 2004.

Now we could try dropping the monotonicity constraint. In this setting cardinality constraints are also superfluous, because it could be that the very large sets have low values. Now it turns out that if $ f$ has no other restrictions (in particular, if it’s allowed to be negative), then even telling whether there’s a set $ S$ with $ f(S) > 0$ is NP-hard, but the optimum could be arbitrarily large and positive when it exists. But if you require that $ f$ is non-negative, then you can get a 1/3-approximation, if you’re willing to add randomness you can get 2/5 in expectation, and with more subtle constraints you can get up to a 1/2 approximation. Anything better is NP-hard. Fiege, Mirrokni, and Vondrak have a nice FOCS paper on this.

Next, we could remove the monotonicity property and try to minimize the value of $ f(S)$. It turns out that this problem always has an efficient solution, but the only algorithm I have heard of to solve it involves a very sophisticated technique called the ellipsoid algorithm. This is heavily related to linear programming and convex optimization, something which I hope to cover in more detail on this blog.

Finally, there are many interesting variations in the algorithmic procedure. For example, one could require that the elements are provided in some order (the streaming setting), and you have to pick at each step whether to put the element in your set or not. Alternatively, the objective functions might not be known ahead of time and you have to try to pick elements to jointly maximize them as they are revealed. These two settings have connections to bandit learning problems, which we’ve covered before on this blog. See this survey of Krause and Golovin for more on the connections, which also contains the main proof used in this post.

Indeed, despite the fact that many of the big results were proved in the 80’s, the analysis of submodular functions is still a big research topic. There was even a paper posted just the other day on the arXiv about its relation to ad serving! And wouldn’t you know, they proved a $ (1-1/e)$-approximation for their setting. There’s just something about $ 1-1/e$.

Until next time!

On Coloring Resilient Graphs

I’m pleased to announce that another paper of mine is finished. This one just got accepted to MFCS 2014, which is being held in Budapest this year (this whole research thing is exciting!). This is joint work with my advisor, Lev Reyzin. As with my first paper, I’d like to explain things here on my blog a bit more informally than a scholarly article allows.

A Recent History of Graph Coloring

One of the first important things you learn when you study graphs is that coloring graphs is hard. Remember that coloring a graph with $ k$ colors means that you assign each vertex a color (a number in $ \left \{ 1, 2, \dots, k \right \}$) so that no vertex is adjacent to a vertex of the same color (no edge is monochromatic). In fact, even deciding whether a graph can be colored with just $ 3$ colors (not to mention finding such a coloring) has no known polynomial time algorithm. It’s what’s called NP-hard, which means that almost everyone believes it’s hopeless to solve efficiently in the worst case.

One might think that there’s some sort of gradient to this problem, that as the graphs get more “complicated” it becomes algorithmically harder to figure out how colorable they are. There are some notions of “simplicity” and “complexity” for graphs, but they hardly fall on a gradient. Just to give the reader an idea, here are some ways to make graph coloring easy:

  • Make sure your graph is planar. Then deciding 4-colorability is easy because the answer is always yes.
  • Make sure your graph is triangle-free and planar. Then finding a 3-coloring is easy.
  • Make sure your graph is perfect (which again requires knowledge about how colorable it is).
  • Make sure your graph has tree-width or clique-width bounded by a constant.
  • Make sure your graph doesn’t have a certain kind of induced subgraph (such as having no induced paths of length 4 or 5).

Let me emphasize that these results are very difficult and tricky to compare. The properties are inherently discrete (either perfect or imperfect, planar or not planar). The fact that the world has not yet agreed upon a universal measure of complexity for graphs (or at least one that makes graph coloring easy to understand) is not a criticism of the chef but a testament to the challenge and intrigue of the dish.

Coloring general graphs is much bleaker, where the focus has turned to approximations. You can’t “approximate” the answer to whether a graph is colorable, so now the key here is that we are actually trying to find an approximate coloring. In particular, if you’re given some graph $ G$ and you don’t know the minimum number of colors needed to color it (say it’s $ \chi(G)$, this is called the chromatic number), can you easily color it with what turns out to be, say, $ 2 \chi(G)$ colors?

Garey and Johnson (the gods of NP-hardness) proved this problem is again hard. In fact, they proved that you can’t do better than twice the number of colors. This might not seem so bad in practice, but the story gets worse. This lower bound was improved by Zuckerman, building on the work of Håstad, to depend on the size of the graph! That is, unless $ P=NP$, all efficient algorithms will use asymptotically more than $ \chi(G) n^{1 – \varepsilon}$ colors for any $ \varepsilon > 0$ in the worst case, where $ n$ is the number of vertices of $ G$. So the best you can hope for is being off by something like a multiplicative factor of $ n / \log n$. You can actually achieve this (it’s nontrivial and takes a lot of work), but it carries that aura of pity for the hopeful graph colorer.

The next avenue is to assume you know the chromatic number of your graph, and see how well you can do then. For example: if you are given the promise that a graph $ G$ is 3-colorable, can you efficiently find a coloring with 8 colors? The best would be if you could find a coloring with 4 colors, but this is already known to be NP-hard.

The best upper bounds, algorithms to find approximate colorings of 3-colorable graphs, also pitifully depend on the size of the graph. Remember I say pitiful not to insult the researchers! This decades-long line of work was extremely difficult and deserves the highest praise. It’s just frustrating that the best known algorithm to color a 3-colorable graph requires as many as $ n^{0.2}$ colors. At least it bypasses the barrier of $ n^{1 – \varepsilon}$ mentioned above, so we know that knowing the chromatic number actually does help.

The lower bounds are a bit more hopeful; it’s known to be NP-hard to color a $ k$-colorable graph using $ 2^{\sqrt[3]{k}}$ colors if $ k$ is sufficiently large. There are a handful of other linear lower bounds that work for all $ k \geq 3$, but to my knowledge this is the best asymptotic result. The big open problem (which I doubt many people have their eye on considering how hard it seems) is to find an upper bound depending only on $ k$. I wonder offhand whether a ridiculous bound like $ k^{k^k}$ colors would be considered progress, and I bet it would.

Our Idea: Resilience

So without big breakthroughs on the front of approximate graph coloring, we propose a new front for investigation. The idea is that we consider graphs which are not only colorable, but remain colorable under the adversarial operation of adding a few new edges. More formally,

Definition: A graph $ G = (V,E)$ is called $ r$-resiliently $ k$-colorable if two properties hold

  1. $ G$ is $ k$-colorable.
  2. For any set $ E’$ of $ r$ edges disjoint from $ E$, the graph $ G’ = (V, E \cup E’)$ is $ k$-colorable.

The simplest nontrivial example of this is 1-resiliently 3-colorable graphs. That is a graph that is 3-colorable and remains 3-colorable no matter which new edge you add. And the question we ask of this example: is there a polynomial time algorithm to 3-color a 1-resiliently 3-colorable graph? We prove in our paper that this is actually NP-hard, but it’s not a trivial thing to see.

The chief benefit of thinking about resiliently colorable graphs is that it provides a clear gradient of complexity from general graphs (zero-resilient) to the empty graph (which is $ (\binom{k+1}{2} – 1)$-resiliently $ k$-colorable). We know that the most complex case is NP-hard, and maximally resilient graphs are trivially colorable. So finding the boundary where resilience makes things easy can shed new light on graph coloring.

Indeed, we argue in the paper that lots of important graphs have stronger resilience properties than one might expect. For example, here are the resilience properties of some famous graphs.

From left to right: the Petersen graph, 2-resiliently 3-colorable; the Dürer graph, 4-resiliently 4-colorable; the Grötzsch graph, 4-resiliently 4-colorable; and the Chvátal graph, 3-resiliently 4-colorable. These are all maximally resilient (no graph is more resilient than stated) and chromatic (no graph is colorable with fewer colors)

From left to right: the Petersen graph, 2-resiliently 3-colorable; the Dürer graph, 4-resiliently 4-colorable; the Grötzsch graph, 4-resiliently 4-colorable; and the Chvátal graph, 3-resiliently 4-colorable. These are all maximally resilient (no graph is more resilient than stated) and chromatic (no graph is colorable with fewer colors)

If I were of a mind to do applied graph theory, I would love to know about the resilience properties of graphs that occur in the wild. For example, the reader probably knows the problem of register allocation is a natural graph coloring problem. I would love to know the resilience properties of such graphs, with the dream that they might be resilient enough on average to admit efficient coloring algorithms.

Unfortunately the only way that I know how to compute resilience properties is via brute-force search, and of course this only works for small graphs and small $ k$. If readers are interested I could post such a program (I wrote it in vanilla python), but for now I’ll just post a table I computed on the proportion of small graphs that have various levels of resilience (note this includes graphs that vacuously satisfy the definition).

Percentage of k-colorable graphs on 6 vertices which are n-resilient
k\n       1       2       3       4
3       58.0    22.7     5.9     1.7
4       93.3    79.3    58.0    35.3
5       99.4    98.1    94.8    89.0
6      100.0   100.0   100.0   100.0

Percentage of k-colorable graphs on 7 vertices which are n-resilient
k\n       1       2       3       4
3       38.1     8.2     1.2     0.3
4       86.7    62.6    35.0    14.9
5       98.7    95.6    88.5    76.2
6       99.9    99.7    99.2    98.3

Percentage of k-colorable graphs on 8 vertices which are n-resilient
k\n       1       2       3       4
3       21.3     2.1     0.2     0.0
4       77.6    44.2    17.0     4.5

The idea is this: if this trend continues, that only some small fraction of all 3-colorable graphs are, say, 2-resiliently 3-colorable graphs, then it should be easy to color them. Why? Because resilience imposes structure on the graphs, and that structure can hopefully be realized in a way that allows us to color easily. We don’t know how to characterize that structure yet, but we can give some structural implications for sufficiently resilient graphs.

For example, a 7-resiliently 5-colorable graph can’t have any subgraphs on 6 vertices with $ \binom{6}{2} – 7$ edges, or else we can add enough edges to get a 6-clique which isn’t 5-colorable. This gives an obvious general property about the sizes of subgraphs in resilient graphs, but as a more concrete instance let’s think about 2-resilient 3-colorable graphs $ G$. This property says that no set of 4 vertices may have more than $ 4 = \binom{4}{2} – 2$ edges in $ G$. This rules out 4-cycles and non-isolated triangles, but is it enough to make 3-coloring easy? We can say that $ G$ is a triangle-free graph and a bunch of disjoint triangles, but it’s known 3-colorable non-planar triangle-free graphs can have arbitrarily large chromatic number, and so the coloring problem is hard. Moreover, 2-resilience isn’t enough to make $ G$ planar. It’s not hard to construct a non-planar counterexample, but proving it’s 2-resilient is a tedious task I relegated to my computer.

Speaking of which, the problem of how to determine whether a $ k$-colorable graph is $ r$-resiliently $ k$-colorable is open. Is this problem even in NP? It certainly seems not to be, but if it had a nice characterization or even stronger necessary conditions than above, we might be able to use them to find efficient coloring algorithms.

In our paper we begin to fill in a table whose completion would characterize the NP-hardness of coloring resilient graphs


The known complexity of k-coloring r-resiliently k-colorable graphs

Ignoring the technical notion of 2-to-1 hardness, the paper accomplishes this as follows. First, we prove some relationships between cells. In particular, if a cell is NP-hard then so are all the cells to the left and below it. So our Theorem 1, that 3-coloring 1-resiliently 3-colorable graphs is NP-hard, gives us the entire black region, though more trivial arguments give all except the (3,1) cell. Also, if a cell is in P (it’s easy to $ k$-color graphs with that resilience), then so are all cells above and to its right. We prove that $ k$-coloring $ \binom{k}{2}$-resiliently $ k$-colorable graphs is easy. This is trivial: no vertex may have degree greater than $ k-1$, and the greedy algorithm can color such graphs with $ k$ colors. So that gives us the entire light gray region.

There is one additional lower bound comes from the fact that it’s NP-hard to $ 2^{\sqrt[3]{k}}$-color a $ k$-colorable graph. In particular, we prove that if you have any function $ f(k)$ that makes it NP-hard to $ f(k)$-color a $ k$-colorable graph, then it is NP-hard to $ f(k)$-color an $ (f(k) – k)$-resiliently $ f(k)$-colorable graph. The exponential lower bound hence gives us a nice linear lower bound, and so we have the following “sufficiently zoomed out” picture of the table


The zoomed out version of the classification table above.

The paper contains the details of how these observations are proved, in addition to the NP-hardness proof for 1-resiliently 3-colorable graphs. This leaves the following open problems:

  • Get an unconditional, concrete linear resilience lower bound for hardness.
  • Find an algorithm that colors graphs that are less resilient than $ O(k^2)$. Even determining specific cells like (4,5) or (5,9) would likely give enough insight for this.
  • Classify the tantalizing (3,2) cell (determine if it’s hard or easy to 3-color a 2-resiliently 3-colorable graph) or even better the (4,2) cell.
  • Find a way to relate resilient coloring back to general coloring. For example, if such and such cell is hard, then you can’t approximate k-coloring to within so many colors.

But Wait, There’s More!

Though this paper focuses on graph coloring, our idea of resilience doesn’t stop there (and this is one reason I like it so much!). One can imagine a notion of resilience for almost any combinatorial problem. If you’re trying to satisfy boolean formulas, you can define resilience to mean that you fix the truth value of some variable (we do this in the paper to build up to our main NP-hardness result of 3-coloring 1-resiliently 3-colorable graphs). You can define resilient set cover to allow the removal of some sets. And any other sort of graph-based problem (Traveling salesman, max cut, etc) can be resiliencified by adding or removing edges, whichever makes the problem more constrained.

So this resilience notion is quite general, though it’s hard to define precisely in a general fashion. There is a general framework called Constraint Satisfaction Problems (CSPs), but resilience here seem too general. A CSP is literally just a bunch of objects which can be assigned some set of values, and a set of constraints (k-ary 0-1-valued functions) that need to all be true for the problem to succeed. If we were to define resilience by “adding any constraint” to a given CSP, then there’s nothing to stop us from adding the negation of an existing constraint (or even the tautologically unsatisfiable constraint!). This kind of resilience would be a vacuous definition, and even if we try to rule out these edge cases, I can imagine plenty of weird things that might happen in their stead. That doesn’t mean there isn’t a nice way to generalize resilience to CSPs, but it would probably involve some sort of “constraint class” of acceptable constraints, and I don’t know a reasonable property to impose on the constraint class to make things work.

So there’s lots of room for future work here. It’s exciting to think where it will take me.

Until then!

Complete Sequences and Magic Tricks

Numberphile posted a video today describing a neat trick based on complete sequences:

The mathematics here is pretty simple, but I noticed at the end of the video that Dr. Grime was constructing the cards by hand, when really this is a job for a computer program. I thought it would be a nice warmup exercise (and a treat to all of the Numberphile viewers) to write a program to construct the cards for any complete sequence.

For the sake of bringing some definitions into it, let’s review the idea of the card trick.

Definition: A sequence of integers $ a_n$ is complete if every natural number $ k$ can be represented as a sum of numbers in $ a_n$.

The examples used in the video are the binary numbers and the fibonacci numbers. The latter is a famous theorem Zeckendorf’s Theorem. Other famous complete sequences include the prime numbers (if you add 1) and the lazy caterer’s sequence.

Now the question is, given a complete sequence, how do we generate the cards from the video? One might recall our post on dynamic programming, in which we wrote a program to compute the optimal coin change for a given amount of money. Indeed, this method works, but we have some added structure in this problem: we know that any number can only be used once. In that case, it is easy to see that there is no need for dynamic programming. Instead, we just use a greedy algorithm.

That is, we can determine which card contains a number $ k$ (which number in the sequence occurs in the representation of $ k$) as follows. Start with the largest card smaller than $ k$, call it $ n$, and we know that $ n$ shows up in the representation of $ k$. To find the next number in the representation, simply repeat the process with the difference $ k-n$. The next number to appear in the representation of $ k$ is hence the largest card less than $ k-n$. We can repeat this process until we run out of cards or the difference becomes zero.

One interesting fact about this algorithm is that it will always produce the smallest representation. That is, in performing the card trick, one is guaranteed the least amount of work in remembering which cards contained the hidden number, and the least amount of work in adding those numbers up.

To implement this algorithm in code, we used Javascript. We chose Javascript so that the reader can run the program and view its source. But we extract the most important bit of the code here:

// compute each card
var j;
for (i = 1; i <= usersMax; i++) {
   var remainder = i;
   for (j = numbers.length-1; j >= 0; j--) {
      if (numbers[j] <= remainder) {
         remainder -= numbers[j];

In words, the “numbers” array contains the complete sequence in question, the cards array is an array of arrays, where each element in the array represents one card. We then loop over all numbers (the $ i$ variable), and for each number we check the sequence in decreasing order to look for membership of $ i$ on a card.

For the sake of brevity, we omit all of the code to deal with input and output, which comprises all of the remaining code in this program.

So to use the program, for example with prime numbers as the complete sequence, one would type “1,2,3,5,7,13,17” into the first text box, and whatever the maximum allowable guess is in the second box, and click the “Generate cards” button.