This is the post about l2w version 1.0, a Latex to WordPress converter painstakingly put together by me with big help from the LaTeX community. Click here to download it. Below is an example of what you can do, taken at random from my class notes which were compiled with this script. I also used this in conjunction with Lyx for several posts such as I believe P=NP, so you can also call this a Lyx to WordPress converter. I just export to latex and then run l2w.

This might work out of the box. More in detail, it needs tex4ht (which is included e.g. in MiKTeX distributions) and Perl (the script only uses minimalistic, shell perl commands). Simply unzip l2w.zip, which contains four files. The file post.tex is this document, which you can edit. To compile, run l2w.bat (which calls myConfig5.cfg). This will create the output post.html which you can copy and past in the wordpress HTML editor. I have tested it on an old Windows XP machine, and a more recent Windows 7 with MixTeX 2.9. I haven’t tested it on linux, which might require some simple changes to l2w.bat. For LyX I add certain commands in the preamble, and as an example the .lyx source of the post I believe P=NP is included in the zip archive.

The non-math source is compiled using full-fledged LaTeX, which means you can use your own macros and bibliography. The math source is not compiled, but more or less left as is for wordpress, which has its own LaTeX interpreter. This means that you can’t use your own macros in math mode. For the same reason, label and ref of equations are a problem. To make them work, the script fetches their values from the .aux file and then crudely applies them. This is a hack with a rather unreadable script; however, it works for me. One catch: your labels should start with eq:.

I hope this will spare you the enormous amount of time it took me to arrive to this solution. Let me know if you use it!

1 Example of what you can do

First, some of the problematic math references:

\begin{aligned} x = 2 ~~~~(1) \end{aligned}

Equation (1).

Next, some weird font stuff: \mathbb {A}, \mathrm {A}, \text {A}.

Lemma 1. Suppose that distributions A^0, A^1 over \{0,1\}^{n_A} are k_A-wise indistinguishable distributions; and distributions B^0, B^1 over \{0,1\}^{n_B} are k_B-wise indistinguishable distributions. Define C^0, C^1 over \{0,1\}^{n_A \cdot n_B} as follows:

C^b: draw a sample x \in \{0,1\}^{n_A} from A^b, and replace each bit x_i by a sample of B^{x_i} (independently).

Then C^0 and C^1 are k_A \cdot k_B-wise indistinguishable.

To finish the proof of the lower bound on the approximate degree of the AND-OR function, it remains to see that AND-OR can distinguish well the distributions C^0 and C^1. For this, we begin with observing that we can assume without loss of generality that the distributions have disjoint supports.

Claim 2. For any function f, and for any k-wise indistinguishable distributions A^0 and A^1, if f can distinguish A^0 and A^1 with probability \epsilon then there are distributions B^0 and B^1 with the same properties (k-wise indistinguishability yet distinguishable by f) and also with disjoint supports. (By disjoint support we mean for any x either \Pr [B^0 = x] = 0 or \Pr [B^1 = x] = 0.)

Proof. Let distribution C be the “common part” of A^0 and A^1. That is to say, we define C such that \Pr [C = x] := \min \{\Pr [A^0 = x], \Pr [A^1 = x]\} multiplied by some constant that normalize C into a distribution.

Then we can write A^0 and A^1 as

\begin{aligned} A^0 &= pC + (1-p) B^0 \,,\\ A^1 &= pC + (1-p) B^1 \,, \end{aligned}

where p \in [0,1], B^0 and B^1 are two distributions. Clearly B^0 and B^1 have disjoint supports.

Then we have

\begin{aligned} \mathbb {E}[f(A^0)] - \mathbb {E}[f(A^1)] =&~p \mathbb {E}[f(C)] + (1-p) \mathbb {E}[f(B^0)] \notag \\ &- p \mathbb {E}[f(C)] - (1-p) \mathbb {E}[f(B^1)] \\ =&~(1-p) \big ( \mathbb {E}[f(B^0)] - \mathbb {E}[f(B^1)] \big ) \\ \leq &~\mathbb {E}[f(B^0)] - \mathbb {E}[f(B^1)] \,. \end{aligned}

Therefore if f can distinguish A^0 and A^1 with probability \epsilon then it can also distinguish B^0 and B^1 with such probability.

Similarly, for all S \neq \varnothing such that |S| \leq k, we have

\begin{aligned} 0 = \mathbb {E}[\chi _S(A^0)] - \mathbb {E}[\chi _S(A^1)] = (1-p) \big ( \mathbb {E}[\chi _S(B^0)] - \mathbb {E}[\chi _S(B^1)] \big ) = 0 \,. \end{aligned}

Hence, B^0 and B^1 are k-wise indistinguishable. \square

Equipped with the above lemma and claim, we can finally prove the following lower bound on the approximate degree of AND-OR.

Theorem 3. d_{1/3}(AND-OR) = \Omega (\sqrt {RN}).

Proof. Let A^0, A^1 be \Omega (\sqrt {R})-wise indistinguishable distributions for AND with advantage 0.99, i.e. \Pr [\mathrm {AND}(A^1) = 1] > \Pr [\mathrm {AND}(A^0) = 1] + 0.99. Let B^0, B^1 be \Omega (\sqrt {N})-wise indistinguishable distributions for OR with advantage 0.99. By the above claim, we can assume that A^0, A^1 have disjoint supports, and the same for B^0, B^1. Compose them by the lemma, getting \Omega (\sqrt {RN})-wise indistinguishable distributions C^0,C^1. We now show that AND-OR can distinguish C^0, C^1:

  • C_0: First sample A^0. As there exists a unique x = 1^R such that \mathrm {AND}(x)= 1, \Pr [A^1 = 1^R] >0. Thus by disjointness of support \Pr [A^0 = 1^R] = 0. Therefore when sampling A^0 we always get a string with at least one “0”. But then “0” is replaced with sample from B^0. We have \Pr [B^0 = 0^N] \geq 0.99, and when B^0 = 0^N, AND-OR=0.
  • C_1: First sample A^1, and we know that A^1 = 1^R with probability at least 0.99. Each bit “1” is replaced by a sample from B^1, and we know that \Pr [B^1 = 0^N] = 0 by disjointness of support. Then AND-OR=1.

Therefore we have d_{1/3}(AND-OR)= \Omega (\sqrt {RN}). \square

1.1 Lower Bound of d_{1/3}(SURJ)

In this subsection we discuss the approximate degree of the surjectivity function. This function is defined as follows.

Definition 4. The surjectivity function SURJ\colon \left (\{0,1\}^{\log R}\right )^N \to \{0,1\}, which takes input (x_1, \dots , x_N) where x_i \in [R] for all i, has value 1 if and only if \forall j \in [R], \exists i\colon x_i = j.

First, some history. Aaronson first proved that the approximate degree of SURJ and other functions on n bits including “the collision problem” is n^{\Omega (1)}. This was motivated by an application in quantum computing. Before this result, even a lower bound of \omega (1) had not been known. Later Shi improved the lower bound to n^{2/3}, see [AS04]. The instructor believes that the quantum framework may have blocked some people from studying this problem, though it may have very well attracted others. Recently Bun and Thaler [BT17] reproved the n^{2/3} lower bound, but in a quantum-free paper, and introducing some different intuition. Soon after, together with Kothari, they proved [BKT17] that the approximate degree of SURJ is \Theta (n^{3/4}).

We shall now prove the \Omega (n^{3/4}) lower bound, though one piece is only sketched. Again we present some things in a different way from the papers.

For the proof, we consider the AND-OR function under the promise that the Hamming weight of the RN input bits is at most N. Call the approximate degree of AND-OR under this promise d_{1/3}^{\leq N}(AND-OR). Then we can prove the following theorems.

Theorem 5. d_{1/3}(SURJ) \geq d_{1/3}^{\leq N}(AND-OR).

Theorem 6. d_{1/3}^{\leq N}(AND-OR) \geq \Omega (N^{3/4}) for some suitable R = \Theta (N).

In our settings, we consider R = \Theta (N). Theorem 5 shows surprisingly that we can somehow “shrink” \Theta (N^2) bits of input into N\log N bits while maintaining the approximate degree of the function, under some promise. Without this promise, we just showed in the last subsection that the approximate degree of AND-OR is \Omega (N) instead of \Omega (N^{3/4}) as in Theorem 6.

Proof of Theorem 5. Define an N \times R matrix Y s.t. the 0/1 variable y_{ij} is the entry in the i-th row j-th column, and y_{ij} = 1 iff x_i = j. We can prove this theorem in following steps:

  1. d_{1/3}(SURJ(\overline {x})) \geq d_{1/3}(AND-OR(\overline {y})) under the promise that each row has weight 1;
  2. let z_j be the sum of the j-th column, then d_{1/3}(AND-OR(\overline {y})) under the promise that each row has weight 1, is at least d_{1/3}(AND-OR(\overline {z})) under the promise that \sum _j z_j = N;
  3. d_{1/3}(AND-OR(\overline {z})) under the promise that \sum _j z_j = N, is at least d_{1/3}^{=N}(AND-OR(\overline {y}));
  4. we can change “=N” into “\leq N”.

Now we prove this theorem step by step.

  1. Let P(x_1, \dots , x_N) be a polynomial for SURJ, where x_i = (x_i)_1, \dots , (x_i)_{\log R}. Then we have
    \begin{aligned} (x_i)_k = \sum _{j: k\text {-th bit of }j \text { is } 1} y_{ij}. \end{aligned}

    Then the polynomial P'(\overline {y}) for AND-OR(\overline {y}) is the polynomial P(\overline {x}) with (x_i)_k replaced as above, thus the degree won’t increase. Correctness follows by the promise.

  2. This is the most extraordinary step, due to Ambainis [Amb05]. In this notation, AND-OR becomes the indicator function of \forall j, z_j \neq 0. Define
    \begin{aligned} Q(z_1, \dots , z_R) := \mathop {\mathbb {E}}_{\substack {\overline {y}: \text { his rows have weight } 1\\ \text {and is consistent with }\overline {z}}} P(\overline {y}). \end{aligned}

    Clearly it is a good approximation of AND-OR(\overline {z}). It remains to show that it’s a polynomial of degree k in z’s if P is a polynomial of degree k in y’s.

    Let’s look at one monomial of degree k in P: y_{i_1j_1}y_{i_2j_2}\cdots y_{i_kj_k}. Observe that all i_\ell ’s are distinct by the promise, and by u^2 = u over \{0,1\}. By chain rule we have

    \begin{aligned} \mathbb {E}[y_{i_1j_1}\cdots y_{i_kj_k}] = \mathbb {E}[y_{i_1j_1}]\mathbb {E}[y_{i_2j_2}|y_{i_1j_1} = 1] \cdots \mathbb {E}[y_{i_kj_k}|y_{i_1j_1}=\cdots =y_{i_{k-1}j_{k-1}} = 1]. \end{aligned}

    By symmetry we have \mathbb {E}[y_{i_1j_1}] = \frac {z_{j_1}}{N}, which is linear in z’s. To get \mathbb {E}[y_{i_2j_2}|y_{i_1j_1} = 1], we know that every other entry in row i_1 is 0, so we give away row i_1, average over y’s such that \left \{\begin {array}{ll} y_{i_1j_1} = 1 &\\ y_{ij} = 0 & j\neq j_1 \end {array}\right . under the promise and consistent with z’s. Therefore

    \begin{aligned} \mathbb {E}[y_{i_2j_2}|y_{i_1j_1} = 1] = \left \{ \begin {array}{ll} \frac {z_{j_2}}{N-1} & j_1 \neq j_2,\\ \frac {z_{j_2}-1}{N-1} & j_1 = j_2. \end {array}\right . \end{aligned}

    In general we have

    \begin{aligned} \mathbb {E}[y_{i_kj_k}|y_{i_1j_1}=\cdots =y_{i_{k-1}j_{k-1}} = 1] = \frac {z_{j_k} - \#\ell < k \colon j_\ell = j_k}{N-k + 1}, \end{aligned}

    which has degree 1 in z’s. Therefore the degree of Q is not larger than that of P.

  3. Note that \forall j, z_j = \sum _i y_{ij}. Hence by replacing z’s by y’s, the degree won’t increase.
  4. We can add a “slack” variable z_0, or equivalently y_{01}, \dots , y_{0N}; then the condition \sum _{j=0}^R z_j = N actually means \sum _{j=1}^R z_j \leq N.


Proof idea for Theorem 6. First, by the duality argument we can verify that d_{1/3}^{\leq N}(f) \geq d if and only if there exists d-wise indistinguishable distributions A, B such that:

  • f can distinguish A, B;
  • A and B are supported on strings of weight \leq N.

Claim 7. d_{1/3}^{\leq \sqrt {N}}(OR_N) = \Omega (N^{1/4}).

The proof needs a little more information about the weight distribution of the indistinguishable distributions corresponding to this claim. Basically, their expected weight is very small.

Now we combine these distributions with the usual ones for And using the lemma mentioned at the beginning.

What remains to show is that the final distribution is supported on Hamming weight \le N. Because by construction the R copies of the distributions for Or are sampled independently, we can use concentration of measure to prove a tail bound. This gives that all but an exponentially small measure of the distribution is supported on strings of weight \le N. The final step of the proof consists of slightly tweaking the distributions to make that measure 0. \square

1.2 Groups

Groups have many applications in theoretical computer science. Barrington [Bar89] used the permutation group S_5 to prove a very surprising result, which states that the majority function can be computed efficiently using only constant bits of memory (something which was conjectured to be false). More recently, catalytic computation [BCK^{+}14] shows that if we have a lot of memory, but it’s full with junk that cannot be erased, we can still compute more than if we had little memory. We will see some interesting properties of groups in the following.

Some famous groups used in computer science are:

  • \{0,1\}^n with bit-wise addition;
  • \mathbb {Z}_m with addition mod m ;
  • S_n, which are permutations of n elements;
  • Wreath product G:= (\mathbb {Z}_m \times \mathbb {Z}_m) \wr \mathbb {Z}_2\,, whose elements are of the form (a,b)z where z is a “flip bit”, with the following multiplication rules:
    • (a, b) 1 = 1 (b, a) ;
    • z\cdot z' := z+z' in \mathbb {Z}_2 ;
    • (a,b) \cdot (a',b') := (a+a', b+b') is the \mathbb {Z}_m\times \mathbb {Z}_m operation;

    An example is (5,7)1 \cdot (2,1) 1 = (5,7) 1 \cdot 1 (1, 2) = (6,9)0 . Generally we have

    \begin{aligned} (a, b) z \cdot (a', b') z' = \left \{ \begin {array}{ll} (a + a', b+b') z+z' & z = 1\,,\\ (a+b', b + a') z+z' & z = 0\,; \end {array}\right . \end{aligned}

  • SL_2(q) := \{2\times 2 matrices over \mathbb {F}_q with determinant 1\}, in other words, group of matrices \begin {pmatrix} a & b\\ c & d \end {pmatrix} such that ad - bc = 1.

The group SL_2(q) was invented by Galois. (If you haven’t, read his biography on wikipedia.)

Quiz. Among these groups, which is the “least abelian”? The latter can be defined in several ways. We focus on this: If we have two high-entropy distributions X, Y over G, does X \cdot Y has more entropy? For example, if X and Y are uniform over some \Omega (|G|) elements, is X\cdot Y close to uniform over G? By “close to” we mean that the statistical distance is less that a small constant from the uniform distribution. For G=(\{0,1\}^n, +), if Y=X uniform over \{0\}\times \{0,1\}^{n-1}, then X\cdot Y is the same, so there is not entropy increase even though X and Y are uniform on half the elements.

Definition 8.[Measure of Entropy] For \lVert A\rVert _2 = \left (\sum _xA(x)^2\right )^{\frac {1}{2}}, we think of \lVert A\rVert ^2_2 = 100 \frac {1}{|G|} for “high entropy”.

Note that \lVert A\rVert ^2_2 is exactly the “collision probability”, i.e. \Pr [A = A']. We will consider the entropy of the uniform distribution U as very small, i.e. \lVert U\rVert ^2_2 = \frac {1}{|G|} \approx \lVert \overline {0}\rVert ^2_2. Then we have

\begin{aligned} \lVert A - U \rVert ^2_2 &= \sum _x \left (A(x) - \frac {1}{|G|}\right )^2\\ &= \sum _x A(x)^2 - 2A(x) \frac {1}{|G|} + \frac {1}{|G|^2} \\ &= \lVert A \rVert ^2_2 - \frac {1}{|G|} \\ &= \lVert A \rVert ^2_2 - \lVert U \rVert ^2_2\\ &\approx \lVert A \rVert ^2_2\,. \end{aligned}

Theorem 9.[[Gow08], [BNP08]] If X, Y are independent over G, then

\begin{aligned} \lVert X\cdot Y - U \rVert _2 \leq \lVert X \rVert _2 \lVert Y \rVert _2 \sqrt {\frac {|G|}{d}}, \end{aligned}

where d is the minimum dimension of irreducible representation of G.

By this theorem, for high entropy distributions X and Y, we get \lVert X\cdot Y - U \rVert _2 \leq \frac {O(1)}{\sqrt {|G|d}}, thus we have

\begin{aligned} ~~~~(2) \lVert X\cdot Y - U \rVert _1 \leq \sqrt {|G|} \lVert X\cdot Y - U \rVert _2 \leq \frac {O(1)}{\sqrt {d}}. \end{aligned}

If d is large, then X \cdot Y is very close to uniform. The following table shows the d’s for the groups we’ve introduced.

G \{0,1\}^n \mathbb {Z}_m (\mathbb {Z}_m \times \mathbb {Z}_m) \wr \mathbb {Z}_2 A_n SL_2(q)

d 1 1 should be very small \frac {\log |G|}{\log \log |G|} |G|^{1/3}

Here A_n is the alternating group of even permutations. We can see that for the first groups, Equation ((2)) doesn’t give non-trivial bounds.

But for A_n we get a non-trivial bound, and for SL_2(q) we get a strong bound: we have \lVert X\cdot Y - U \rVert _2 \leq \frac {1}{|G|^{\Omega (1)}}.


[Amb05]    Andris Ambainis. Polynomial degree and lower bounds in quantum complexity: Collision and element distinctness with small range. Theory of Computing, 1(1):37–46, 2005.

[AS04]    Scott Aaronson and Yaoyun Shi. Quantum lower bounds for the collision and the element distinctness problems. J. of the ACM, 51(4):595–605, 2004.

[Bar89]    David A. Mix Barrington. Bounded-width polynomial-size branching programs recognize exactly those languages in NC^1. J. of Computer and System Sciences, 38(1):150–164, 1989.

[BCK^{+}14]    Harry Buhrman, Richard Cleve, Michal Koucký, Bruno Loff, and Florian Speelman. Computing with a full memory: catalytic space. In ACM Symp. on the Theory of Computing (STOC), pages 857–866, 2014.

[BKT17]    Mark Bun, Robin Kothari, and Justin Thaler. The polynomial method strikes back: Tight quantum query bounds via dual polynomials. CoRR, arXiv:1710.09079, 2017.

[BNP08]    László Babai, Nikolay Nikolov, and László Pyber. Product growth and mixing in finite groups. In ACM-SIAM Symp. on Discrete Algorithms (SODA), pages 248–257, 2008.

[BT17]    Mark Bun and Justin Thaler. A nearly optimal lower bound on the approximate degree of AC0. CoRR, abs/1703.05784, 2017.

[Gow08]    W. T. Gowers. Quasirandom groups. Combinatorics, Probability & Computing, 17(3):363–387, 2008.


3 thoughts on “l2w

  1. Very neat, Emanuele, great work! What about using KaTeX or Mathjax to display the equations? In my browser, for instance, your math bits appear quite smaller than the body of the text. Cheers!

      1. Goodness, really? I did not know that, sorry. Adding KaTeX is literally just one more line in the page, that’s a shame. Cheers!

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