02-25-14 - WiFi

So our WiFi stopped working recently, and I discovered a few things which I will now write down.

First of all, WiFi is fucked. 2.4 GHz is way overcrowded and just keeps getting more crowded. Lots of fucking routers now are offering increased bandwidth by using multiple channels simultaneously, etc. etc. It's one big interference fest.

The first issue I found was baby monitors. Baby monitors, like many wireless devices, are also in the 2.4 GHz band and just crap all over your wifi. Turning them off helped our signal a bit, but we were still getting constant problems.

Next issue is interference from neighbors'ses wifises. This is what inSSIDer looks like at my house :

We are *way* away from any neighbors, at least 50 feet in every direction, and we still get this amount of shit from them. Each of my cock-ass-fuck neighbors seems to have four or five wifi networks. Good job guys, way to fix your signal strength issues by just piling more shit in the spectrum.

What you can't see from the static image is that lots of the fucking neighbor wifis are not locked to a specific channel, many of them are constantly jumping around trying to find a clear channel, which just makes them crud them all up.

(I'd love to get some kind of super industrial strength wifi for my house and just crank it up to infinity and put it on every channel so that nobody for a mile around gets any wifi)

I've long had our WiFi on channel 8 because it looked like the most clear spot to be. Well, it turns out that was a classic newb mistake. Apparently it's worse to be slightly offset from a busy channel than it is to be right on it. When you're offset, you get signal leakage from the other channel that just looks like noise; being on the channel you're fighting with other people, but at least you are seing their data as real data that you can ignore. Anyway, switching our network to channel 11 fixed it.

It looks like in practice channel 6 and 11 are the only usable ones in noisy environments (eg. everywhere).

The new 802.11ac on 5 GHz should be a nice clean way to go for a few years until it too gets crudded up.

02-25-14 - ANS Applications

Some rambling on where I see ANS being useful.

In brief - anywhere you used Huffman in the past, you can use ANS instead.

ANS (or ABS) are not very useful for high end compressors. They are awkward for adaptive modeling. Even if all you care about is decode speed (so you don't mind the buffering up the models to make the encode work backwards) it's just not a big win over arithmetic coding. Things like PAQ/CM , LZMA, H264, all the high compression cases that use adaptive context models, there's no real win from ANS/ABS.

Some specific cases where I see ANS being a good win :

JPEG-ANS obviously. Won't be competitive with sophisticated coders like "packjpg" but will at least fix the cliff at low bit rate caused by Huffman coding.

JPEGNEXT-ANS. I've been thinking for a long time about writing a "JPEGNEXT". Back end coefficient classification; coefficients in each group sent by value with ANS. Front end 32x32 macroblocks with various DCT sizes. Try to keep it as simple as possible but be up to modern quality levels.

LZ-ANS. An "LZX class" (which means large window, "repeat match", entropy resets) LZ with ANS back end should be solid. Won't be quite LZMA compression levels, but way way faster.

Lossless image DPCM. ANS on prediction residual values is a clear killer. Should have a few statistics groups with block classification. No need to transmit the ANS counts if you use a parametric model ala TMW etc. Should be easy, fast, and good.

blocksort-ANS. Should replace bzip. Fast to decode.

MPEG-class video coders. Fast wavelet coders. Anywhere you are using minimal context modelling (only a few contexts) and are sending values by their magnitude (not bit plane coders).



02-18-14 - ans_fast implementation notes

Some notes about the ans_fast I posted earlier .

ans_fast contains a tANS (table-ANS) implementation and a rANS (range-ANS) implementation.

First, the benchmarking. You can compare to the more naive implementations I posted earlier . However, do not compare this tANS impl to Huffman or arithmetic and conclude "ANS is faster" because the tANS impl here has had rather more loving than those. Most of the tricks used on "ans_fast" can be equally used for other algorithms (though not all).

Here L=4096 to match the 12-bits used in the previous test. This is x64 on my lappy (1.7 Ghz Core i7 with turbo disabled). Compressed sizes do not include sending the counts. Time "withtable" includes all table construction but not histogramming or count normalization (that affects encoder only). ("fse" and "huf" on the previous page included table transmission and histogramming time)


tANS 768771 -> 435252.75

ticks to encode: 4.64 decode: 3.39
mbps encode: 372.92 decode: 509.63

withtable ticks to encode: 4.69 decode: 3.44
withtable mbps encode: 368.65 decode: 501.95

rANS 768771 -> 435980 bytes (v2)

ticks to encode: 6.97 decode: 5.06
mbps encode: 248.02 decode: 341.63

withtable ticks to encode: 6.97 decode: 5.07
withtable mbps encode: 247.92 decode: 341.27


tANS 513216 -> 78856.88

ticks to encode: 4.53 decode: 3.47
mbps encode: 382.02 decode: 497.75

withtable ticks to encode: 4.62 decode: 3.56
withtable mbps encode: 374.45 decode: 485.40

rANS 513216 -> 79480 bytes (v2)

ticks to encode: 5.62 decode: 3.53
mbps encode: 307.78 decode: 490.32

withtable ticks to encode: 5.63 decode: 3.54
withtable mbps encode: 307.26 decode: 488.88

First a note on file sizes : rANS file sizes are a few bytes larger than the "rans 12" posted last time. That's because that was a 32-bit impl. The rANS here is 64-bit and dual-state so I have to flush 16 bytes instead of 4. There are ways to recover some of those bytes.

The tANS file sizes here are smaller than comparable coders. The win comes from the improvements to normalizing counts and making the sort order. In fact, the +1 bias heuristic lets me beat "arith 12" and "rans 12" from the last post, which were coding nearly perfectly to the expected codelen of the normalized counts.

If you run "ans_learning" you will often see that the written bits are less than the predicted codelen :

H = 1.210176
CL = 1.238785
wrote 1.229845 bpb
this is because the +1 bias heuristic lets the codelens match the data better than the normalized counts do.

Okay, so on to the speed.

The biggest thing is that the above reported speeds are for 2x interleaved coders. That is, two independent states encoding the single buffer to a single compressed stream. I believe ryg will talk about this more soon. You can read his paper on arxiv now. Note that this is not just unrolling. Because the states are independent they allow independent execution chains to be in flight at the same time.

The speedup from interleaving is pretty huge (around 1.4X) :


rANS non-interleaved (v1)

ticks to encode: 26.84 decode: 7.33
mbps encode: 64.41 decode: 235.97

withtable ticks to encode: 26.85 decode: 7.38
withtable mbps encode: 64.41 decode: 234.19

rANS 2x interleaved (v1)

ticks to encode: 17.15 decode: 5.16
mbps encode: 100.84 decode: 334.95

withtable ticks to encode: 17.15 decode: 5.22
withtable mbps encode: 100.83 decode: 331.31

tANS non-interleaved

ticks to encode: 6.43 decode: 4.68
mbps encode: 269.10 decode: 369.44

withtable ticks to encode: 6.48 decode: 4.73
withtable mbps encode: 266.86 decode: 365.39

tANS 2x interleaved

ticks to encode: 4.64 decode: 3.39
mbps encode: 372.92 decode: 509.63

withtable ticks to encode: 4.69 decode: 3.44
withtable mbps encode: 368.65 decode: 501.95

But even non-interleaved it's fast. (note that interleaved tANS is using only a single shared bit buffer). The rest of the implementation discussion will use the non-interleaved versions for simplicity.

The tANS implementation is pretty straightforward.

Decoding one symbol is :

    struct decode_entry { uint16 next_state; uint8 num_bits; uint8 sym; };

    decode_entry * detable = table - L;

    #define DECODE_ONE() do { \
        de = detable + state; \
        nb = de->num_bits; \
        state = de->next_state; \
        BITIN_OR(bitin_bits,bitin_numbits,nb,state); \
        *toptr++ = (uint8) de->sym; \
    } while(0)

where BITIN_OR reads "nb" bits and ors them onto state.

With a 64-bit bit buffer, I can ensure >= 56 bits are in the buffer. That means with L up to 14 bits, I can do four decodes before checking for more bits needed. So the primary decode loop is :

        // I know >= 56 bits are available  
        // each decode consumes <= 14 bits

        // now >= 56 bits again

The fastest way I could find to do the bit IO was "big endian style". That's the next bits at the top of the word. Bits in the word are in order of bits in the file. This lets you unconditionally grab the next 8 bytes to refill, but requires a bswap (on little endian platforms). eg :

#define BITIN_REFILL(bitin_bits,bitin_numbits,bitin_ptr) do { \
        ASSERT( bitin_numbits > 0 && bitin_numbits <= 64 ); \
        int64 bytesToGet = (64 - bitin_numbits)>>3; \
        uint64 next8 = _byteswap_uint64( *( (uint64 *)bitin_ptr ) ); \
        bitin_ptr += bytesToGet; \
        bitin_bits |= (next8 >> 1) >> (bitin_numbits-1); \
        bitin_numbits += bytesToGet<<3; \
        ASSERT( bitin_numbits >= 56 && bitin_numbits <= 64 ); \
    } while(0)

The other nice thing about the bits-at-top style is that the encoder can put bits in the word without any masking. The encoder is :

    #define ENCODE_ONE() do { \
        sym = *bufptr--; ee = eetable+sym; \
        msnb = ee->max_state_numbits; \
        msnb += ( state >= ee->max_state_thresh ); \
        BITOUT_PUT(bout_bits,bout_numbits, state,msnb); \
        state = ee->packed_table_ptr[ state>>msnb ]; \
        } while(0)

    #define BITOUT_PUT(bout_bits,bout_numbits,val,nb) do { \
        ASSERT( (bout_numbits+nb) <= 64 ); \
        bout_bits >>= nb; \
        bout_bits |= ((uint64)val) << (64 - nb); \
        bout_numbits += nb; \
    } while(0)

the key interesting part being that the encoder just does BITOUT_PUT with "state", and by shifting it up to the top of the word for the bitio, it gets automatically masked. (and rotate-right is a way to make that even faster).

Similarly to the decoder, the encoder can write 4 symbols before it has to check if the bit buffer needs any output.

The other crucial thing for fast tANS is the sort order construction. I do a real sort, using a radix sort. I do the first step of radix sorting (generating a histogram), and then I directly build the tables from that, reading out of the radix histogram. There's no need to explicitly generate the sorted symbol list as an intermediate step. I use only an 8-bit radix here (256 entries) but it's not significantly different (in speed or compression) than using a larger radix table.

The rANS implementation is pretty straightforward and I didn't spend much time on it, so it could probably be faster (particularly encoding which I didn't spend any time on (ADDENDUM : v2 rANS now sped up and encoder uses fast reciprocals)). I use a 64-bit state with 32-bit renormalization. The basic decode operation is :

        uint64 xm = x & mask;   
        const rans_decode_table::entry & e = table[xm];
        x = e.freq * (x >> cumprobtot_bits) + e.xm_minus_low;
        buffer[i] = (uint8) e.sym;
        if ( x < min_x )
            x <<= 32;
            x |= *((uint32 *)comp_ptr);
            comp_ptr += 4;

One thing I should note is that my rANS decode table is 2x bigger than the tANS decode table. I found it was fastest to use an 8-byte decode entry for rANS :

    // 8-byte decode entry
    struct entry { uint16 freq; uint16 xm_minus_low; uint8 sym; uint16 pad; };
obviously you can pack that a lot smaller (32 bits from 12+12+8) but it hurts speed.

For both tANS and rANS I make the encoder write backwards and the decoder read forwards to bias in favor of decoder speed. I make "L" a variable, not a constant, which hurts speed a little.

02-18-14 - Understanding ANS - Conclusion

I think we can finally say that we understand ANS pretty well, so this series will end. I may cover some more ANS topics but they won't be "Understanding ANS".

Here is the index of all posts on this topic :

cbloom rants 1-30-14 - Understanding ANS - 1
cbloom rants 1-31-14 - Understanding ANS - 2
cbloom rants 02-01-14 - Understanding ANS - 3
cbloom rants 02-02-14 - Understanding ANS - 4
cbloom rants 02-03-14 - Understanding ANS - 5
cbloom rants 02-04-14 - Understanding ANS - 6
cbloom rants 02-05-14 - Understanding ANS - 7
cbloom rants 02-06-14 - Understanding ANS - 8
cbloom rants 02-10-14 - Understanding ANS - 9
cbloom rants 02-11-14 - Understanding ANS - 10
cbloom rants 02-14-14 - Understanding ANS - 11
cbloom rants 02-18-14 - Understanding ANS - 12

And here is some source code for my ANS implementation : (v2 02/21/2014)

ans.zip - contains ans_fast and ans_learning
cblib.zip is required to build my code

My home code is MSVC 2005/2008. Port if you like. Email me if you need help.

NOTE : this release is not a library you just download and use. It is intended as documentation of research. If you want some ANS code you can just use off the shelf, go get FiniteStateEntropy . You may also want ryg_rans .

I think I'll do a followup post with the performance of ans_fast and some optimization notes so it doesn't crud up this index post. Please put implementation speed discussion in that followup post .

02-18-14 - Understanding ANS - 12

A little note about sorts and tables.


What's wrong with that sort?

(That's the naive rANS sort order; it's just a "cum2sym" table. It's each symbol Fs times in consecutive blocks. It has M=32 entries. M = sum{Fs} , L = coding precision)

(here I'm talking about a tANS implementation with L=M ; the larger (L/M) is, the more you preserve the information in the state x)

Think about what the state variable "x" does as you are coding. In the renormalization range it's in [32,63]. Its position in that range is a slider for the number of fraction bits it contains. At the bottom of the range, log2(x) is 5, at the top log2(x) is 6.

Any time you want to encode a "D" you must go back to a singleton precursor state, Is = [1]. That means you have to output all the bits in x, so all fractional bits are thrown away. All information about where you were in that I range is gone. Then from that singleton Is range you jump to the end of the I range.

(if Fs=2 , then you quantize the fractional bits up to 0.5 ; is Fs=3, you quantize to 1/3 of a bit, etc.)

Obviously the actual codelen for a "D" is longer than that for an "A". But so is the codelen for a "C", and the codelen for "A" is too short. Another way to think of it is that you're taking an initial state x that spans the whole interval [32,63] and thus has variable fractional bits, and you're mapping it into only a portion of the interval.

In order to preserve the fractional bit state size, you want to map from the whole interval back to the whole interval. In the most extreme case, something like :


(M=16) , when you encode an A you go from [16,31] to [8,15] and then back the A's in that string. The net result is that state just lost its bottom bit. That is, x &= ~1. You still have the full range of possible fractional bits from [0,1] , you just lost the bottom bit of precision.

I was thinking about this because I was making some weird alternative tANS tables. In fact I suppose not actually ANS tables, but more general coding tables.

For background, you can make one of the heuristic tANS tables thusly :

shuffle(s) = some permutation function
shuffle is one-to-one over the range [0,L-1]
such as Yann's stepping prime-mod-L
or bit reverse

    int next_state[256];    
    uint8 permutation[MAX_L];
    // make permutation :
    uint32 cumulative_count = 0;    
    for LOOP(sym,alphabet)
        uint32 count = normalized_counts[sym];
        if ( count == 0 ) continue;
        next_state[sym] = count;
        for LOOP(c,(int)count)
            uint32 index = shuffle(cumulative_count);
            permutation[index] = (uint8)sym;
    ASSERT( cumulative_count == (uint32)L );

    // permutation is now our "output string"   

    for LOOP(i,L) // iterating over destination state
        int sym = permutation[i];
        // step through states for this symbol
        int from_state = next_state[sym];
        next_state[sym] ++;
        int to_state = L + i;
        encode_packed_table_ptr[sym][from_state] = to_state;

which is all well and good. But I started thinking - can I eliminate the intermediate permutation[] table entirely? Well, yes. There are a few ways.

If you have a "cum2sym" table already handy, then you can just use shuffle() to look up directly into cum2sym[], and that is identical to the above. But you probably don't have cum2sym.

Well what if we just use shuffle() to make the destination state? Note that this is calling it in the opposite direction (from cum2sym index to to_state , rather than from to_state to cum2sym). If your shuffle is self-inverting like bit reversal is, then it's the same.

It gives you a very simple table construction :

    uint32 cumulative_count = 0;    
    for LOOP(sym,alphabet)
        uint32 count = normalized_counts[sym];
        if ( count == 0 ) continue;
        for LOOP(c,(int)count)
            uint32 index = shuffle(cumulative_count);

            uint32 to_state = index + L;
            int from_state = count + c; 

            encode_packed_table_ptr[sym][from_state] = to_state;
    ASSERT( cumulative_count == (uint32)L );

make_tans_shuffle_direct walks the Fs in a kind of cum2sym order and then scatters those symbols out to semi-random target locations using the shuffle() function.

It doesn't work. Or rather, it works, it encodes & decodes data correctly, but the total coded size is worse.

The problem is that the encode table is no longer monotonic. That is, as "from_state" increases, "to_state" does not necessarily increase. The Fs encode table entries for each symbol are not numerically in order.

In the images we've been picturing from earlier in the post we can see the problem. Some initial state x is renormalized down to the Is coding range. We then follow the state transition back to the I range - but we go somewhere random. We don't go to the same neighborhood where we started, so we randomly get more or less fractional bits.

You can fix it thusly :

    uint32 cumulative_count = 0;    
    for LOOP(sym,alphabet)
        uint32 count = normalized_counts[sym];
        if ( count == 0 ) continue;
        for LOOP(c,(int)count)
            uint32 index = shuffle(cumulative_count);

            uint32 to_state = index + L;
            int from_state = count + c; 

            encode_packed_table_ptr[sym][from_state] = to_state;

        // fix - to_states not monotonic
        // sort the destination states for this symbol :
        std::sort( encode_packed_table_ptr[sym]+count, encode_packed_table_ptr[sym]+2*count );
    ASSERT( cumulative_count == (uint32)L );

and then it is identical to "make_tans_shuffle" (identical if shuffle is self-inverting, and if not then it's different but equal, since shuffle is really just a random number generator so running it backwards doesn't hurt compression).

For the record the compression penalty for getting the state transition order wrong is 1-2% :

CCC total bytes out :

correct sort : 1788631
shuffle fixed: 1789655
shuffle bad  : 1813450


02-14-14 - Understanding ANS - 11

I want to do some hand waving about the different ways you can conceptually look at ANS.

Perhaps the best way to understand ANS mathematically is via the analogy with arithmetic coding . While ANS is not actually building up an arithmetic coder interval for the file, each step is very much like a LIFO arithmetic coder, and the (x/P) scaling is what makes x grow the right amount for each symbol. This is the most useful way to think about rANS or uANS, I find.

But there are other ways to think about it.

One is Duda's "asymmetric numeral system", which is how he starts the explanation in the paper, and really confused me to begin with. Now that we've come at ANS from the back way we can understand what he was on about.

The fundamental op in ANS is :

integer x contains some previous value

make x' = x scaled up in some way + new value 

with a normal "symmetric numeral system" , you would just do base-b math :

x' = x * b + v

which gives you an x' where the old value (x) is distributed evenly, and the v's just cycle :

b = 3 for example

x':  0  1  2  3  4  5  6  7  8 ... 
x :  0  0  0  1  1  1  2  2  2
v :  0  1  2  0  1  2  0  1  2

x' is a way of packing the old value x and the new value v together. This symmetric packing corresponds to the output string "012" in the parlance of this series. The growth factor (x'/x) determines the number of bits required to send our value, and it's uniform.

But it doesn't need to be uniform.

0102 :

x':  0  1  2  3  4  5  6  7  8 ... 
x :  0  0  1  0  2  1  3  1  4 
v :  0  1  0  2  0  1  0  2  0

Intuitively, the more often a symbol occurs in the output string, the more slots there are for the previous value (x) to get placed; that is, more bits of x can be sent in lower values of x' when the symbol occurs in many slots. Hence x' grows less. If you're thinking in terms of normalized x's, growing less means you have to output fewer bits to stay in the renormalization range.

You can draw these asymmetric numeral lines in different ways, which Duda does in the paper. For example :

input x as the axis line,
output x' in the table :

  0 1 2 3 4 5 6  x
A 0 2 4          x'
B 1 3 5


  0 1 2 3 4 5 6  x
A 0 1 3 4 6 7 9  x'
B 2 5 8 11

output x' as the axis line
input x in the table :

  0 1 2 3 4 5 6  x'
A 0   1   2   3  x
B   0   1   2

  0 1 2 3 4 5 6  x'
A 0 1   2 3   4  x
B     0     1

output x' line implicit
show x and output symbol :

0 0 1 1 2 2 3

0 1 0 2 3 1 4

That is, it's a funny way of just doing base-b math; we're shifting up the place value and adding our value in, but we're in an "asymmetric numeral system", so the base is nonuniform. I find this mental image not very useful when thinking about how the coding actually works.

There's another way to think about tANS in particular (tANS = table-based ANS), which is what Yann is talking about .

To get there mentally, we actually need to optimize our tANS code.

When I covered tANS encoding before , I described it something like this :

x is the current state
x starts in the range I = [L,2L-1]

to encode the next symbol s
we need to reach the precursor range Is = [Fs,2Fs-1]

to do that, output bits from x
b = x&1; x >>= 1;
until x is lowered to reach Is

then take the state transition C()
this takes x back into I

this should be familiar and straightforward.

To optimize, we know that x always starts in a single power of 2 interval [L,2L-1] , and it always lands in a power of 2 interval [Fs,2Fs-1]. That means the minimum number of bits we ever output is from L to 2Fs-1 , and the maximum number of bits is only 1 more than that. So the renormalization can be written as :

precompute :

max_state = 2Fs - 1;
min_num_bits = floor(log2(L/Fs));

to renormalize :

x in [L,2L-1]
output min_num_bits from x
x >>= min_num_bits

now ( x >= Fs && x < 2*max_state );

if ( x > max_state ) output 1 more bit; x>>= 1;

now x in [Fs,2Fs-1]

But you can move the check for the extra output bit earlier, before shifting x down :

precompute :

min_num_bits = log2(L) - log2ceil(Fs);  // if L is power of 2
threshold = (2*Fs)<<num_bits;

to renormalize :

x in [L,2L-1]
num_bits = min_num_bits;
if ( x >= threshold ) num_bits ++;
output num_bits from x
x >>= num_bits

x in [Fs,2Fs-1]

and then use C(x) since x is now in Is.

It's just straightforward optimization, but it actually allows us to picture the whole process in a different way. Let's write the same encoder, but just in terms of a table index :

let t = x - L
t in [0,L-1]

t is a table index.

to encode s :

num_bits = min_num_bits[s] + ( t >= threshold[s] );
bitoutput( t, num_bits );
t = encode_state_table[s][ (t+L)>>num_bits ];

That is, we're going from a table index to another table index. We're no longer thinking about going back to the [Fs,2Fs-1] precursor range at all.

Before we got our desired code len by the scaling of the intervals [L,2L)/[Fs,2Fs) , now the code len is the stored number of bits. We can see that we get fractional bits because sometimes we output one more.

Let's revisit an example that we went through previously , but with this new image.

L = 8
Fs = {3,3,2}
output string = ABCABCAB

We can see right away that our table index t is 3 bits. To encode a 'C' there will be only two slots on our numeral line that correspond to a lower digit of C, so we must output 2 bits and keep 1 bit of t. To encode an 'A' we can keep 3 values, so we can output 1 bit for t in [0,3] and 2 bits for t in [4,7] ; that will give us 2 retained values in the first region and 1 retained value in the second.

Explicitly :

t in [0,7]
I want to encode an A
so I want to reach {AxxAxxAx}

t in [0,3]
  output t&1
  index = (t+L)>>1 = 4 or 5
  take the last two A's {xxxAxxAx}
  so state -> 3 or 6

t in [4,7]
  output t&3
  index = (t+L)>>2 = 3
  take the first A {Axxxxxxx}
  state -> 0

Note that the way we're doing it, high states transition to low states, and vice versa. These comes up because of the +L sentry bit method to separate the subranges produced by the shift.

The tANS construction creates this encode table :

A : b=1+(t>=4) : {0,3,6}
B : b=1+(t>=4) : {1,4,7}
C : b=2+(t>=8) : {2,5}

It should be obvious that we can now drop all our mental ideas about "ANS" and just make these coding tables directly. All you need is an output string, and you think about doing these kinds of mapping :

t in [0,7]

I want to encode a B

[xxxxxxxx] -> [xBxxBxxB]

output bits to reduce the 3 values
and transition to one of the slots with a B

The decode table is trivial to make from the inverse :

 0: A -> 4 + getbits(2)
 1: B -> 4 + getbits(2)
 2: C -> 0 + getbits(2)
 3: A -> 0 + getbits(1)
 4: B -> 0 + getbits(1)
 5: C -> 4 + getbits(2)
 6: A -> 2 + getbits(1)
 7: B -> 2 + getbits(1)

Note that each symbol's decode covers the entire origin state range :

 0: A -> 4 + getbits(2)  from [4,7]
 3: A -> 0 + getbits(1)  from [0,1]
 6: A -> 2 + getbits(1)  from [2,3]

 1: B -> 4 + getbits(2)  from [4,7]
 4: B -> 0 + getbits(1)  from [0,1]
 7: B -> 2 + getbits(1)  from [2,3]

 2: C -> 0 + getbits(2)  from [0,3]
 5: C -> 4 + getbits(2)  from [4,7]

During decode we can think about our table index 't' as containing two pieces of information : one is the current symbol to output, but there's also some information about the range where t will be on the next step. That is, the current t contains some bits of the next t. The number of bits depends on where we are in the table. eg. in the example above; when t=4 we specify a B, but we also specify 2 bits worth of the next t.

Doing another example from that earlier post :



A : b=1+(t>=12) : {0,3,5,7,10,12,15}
B : b=1+(t>=8) : {1,4,6,9,11,14}
C : b=2+(t>=8) : {2,8,13}

 0: A -> 12 + getbits(2)
 1: B -> 8 + getbits(2)
 2: C -> 8 + getbits(3)
 3: A -> 0 + getbits(1)
 4: B -> 12 + getbits(2)
 5: A -> 2 + getbits(1)
 6: B -> 0 + getbits(1)
 7: A -> 4 + getbits(1)
 8: C -> 0 + getbits(2)
 9: B -> 2 + getbits(1)
10: A -> 6 + getbits(1)
11: B -> 4 + getbits(1)
12: A -> 8 + getbits(1)
13: C -> 4 + getbits(2)
14: B -> 6 + getbits(1)
15: A -> 10 + getbits(1)

and this concludes our conception of tANS in terms of just an [0,t-1] table.

I'm gonna be super redundant and repeat myself some more. I think it's intriguing that we went through all this ANS entropy coder idea, scaling values by (x/P) and so on, and from that we constructed tANS code. But you can get to the exact same tANS directly from the idea of the output string!

Let's invent tANS our new way, starting from scratch.

I'm given normalized frequencies {Fs}. Sum{Fs} = L. I want a state machine with L entries. Take each symbol and scatter it into our output string in some way.

To encode each symbol, I need to map the state machine index t in [0,L-1] to one of its occurances in the output string.

There are Fs occurances in the output string

I need to map an [0,L-1] value to an [0,Fs-1] value
by outputting either b or b+1 bits

now clearly if (L/Fs) is a power of 2, then the log2 of that is just b and we always output that many bits. (eg L=16, Fs=4, we just output 2 bits). In general if (L/Fs) is not a power of 2, then

b = floor(log2(L/Fs))
b+1 = ceil(log2(L/Fs))

so we just need two sub-ranges of L such that the total adds up to Fs :

threshold T
values < T output b bits
values >= T output b+1 bits

total of both ranges after output should equal Fs :

(T>>b) + (L-T)>>(b+1) = Fs

(2T + L-T)>>(b+1) = Fs

L+T = Fs<<(b+1)

T = (Fs<<(b+1)) - L

and that's it! We've just made a tANS encoder without talking about anything related to the ANS ideas at all.

The funny thing to me is that we got the exact same condition before from "b-uniqueness". That is, in order to be able to encode symbol s from any initial state, we worked out that the only valid precursor range was Is = [Fs,2*Fs-1] . That leads us to the renormalization loop :

while x > (2*Fs-1)
  output a bit from x; x>>= 1;

And from that we computed a minimum number of output bits, and a threshold state for one more. That threshold we computed was

(max_state + 1)<<min_num_bits

= (2*Fs-1 + 1)<<b
= Fs<<(b+1)

which is the same.


02-11-14 - Understanding ANS - 10 - Optimal normalized counts

Not really an ANS topic, but a piece you need for ANS so I've had a look at it.

For ANS and many other statistical coders (eg. arithmetic coding) you need to create scaled frequencies (the Fs in ANS terminology) from the true counts.

But how do you do that? I've seen many heuristics over the years that are more or less good, but I've never actually seen the right answer. How do you scale to minimize total code len? Well let's do it.

Let's state the problem :

You are given some true counts Cs

Sum{Cs} = T  the total of true counts

the true probabilities then are

Ps = Cs/T

and the ideal code lens are log2(1/Ps)

You need to create scaled frequencies Fs
such that

Sum{Fs} = M

for some given M.

and our goal is to minimize the total code len under the counts Fs.

The ideal entropy of the given counts is :

H = Sum{ Ps * log2(1/Ps) }

The code len under the counts Fs is :

L = Sum{ Ps * log2(M/Fs) }

The code len is strictly worse than the entropy

L >= H

We must also meet the constraint

if ( Cs != 0 ) then Fs > 0

That is, all symbols that exist in the set must be codeable. (note that this is not actually optimal; it's usually better to replace all rare symbols with a single escape symbol, but we won't do that here).

The naive solution is :

Fs = round( M * Ps )

if ( Cs > 0 ) Fs = MAX(Fs,1);

which is just scaling up the Ps by M. This has two problems - one is that Sum{Fs} is not actually M. The other is that just rounding the float does not actually distribute the integer counts to minimize codelen.

The usual heuristic is to do something like the above, and then apply some fix to make the sum right.

So first let's address how to fix the sum. We will always have issues with the sum being off M because of integer rounding.

What you will have is some correction :

correction = M - Sum{Fs}

that can be positive or negative. This is a count that needs to be added onto some symbols. We want to add it to the symbols that give us the most benefit to L, the total code len. Well that's simple, we just measure the affect of changing each Fs :

correction_sign = correction > 0 ? 1 : -1;

Ls_before = Ps * log2(M/Fs)
Ls_after = Ps * log2(M/(Fs + correction_sign))

Ls_delta = Ls_after - Ls_before
Ls_delta = Ps * ( log2(M/(Fs + correction_sign)) - log2(M/Fs) )
Ls_delta = Ps * log2(Fs/(Fs + correction_sign))

so we need to just find the symbol that gives us the lowest Ls_delta. This is either an improvement to total L, or the least increase in L.

We need to apply multiple corrections. We don't want a solution thats O(alphabet*correction) , since that can be 256*256 in bad cases. (correction is <= alphabet and typically in the 1-50 range for a typical 256-symbol file). The obvious solution is a heap. In pseudocode :

For all s
    push_heap( Ls_delta , s )

For correction
    s = pop_heap
    adjust Fs
    compute new Ls_delta for s
    push_heap( Ls_delta , s )

note that after we adjust the count we need to recompute Ls_delta and repush that symbol, because we might want to choose the same symbol again later.

In STL+cblib this is :

to[] = Fs
from[] = original counts

struct sort_sym
    int sym;
    float rank;
    sort_sym() { }
    sort_sym( int s, float r ) : sym(s) , rank(r) { }
    bool operator < (const sort_sym & rhs) const { return rank < rhs.rank; }


    if ( correction != 0 )
        int32 correction_sign = (correction > 0) ? 1 : -1;

        vector<sort_sym> heap;

        for LOOP(i,alphabet)
            if ( from[i] == 0 ) continue;
            ASSERT( to[i] != 0 );
            if ( to[i] > 1 || correction_sign == 1 )
                double change = log( (double) to[i] / (to[i] + correction_sign) ) * from[i];
                heap.push_back( sort_sym(i,change) );
        while( correction != 0 )
            ASSERT_RELEASE( ! heap.empty() );
            sort_sym ss = heap.back();
            int i = ss.sym;
            ASSERT( from[i] != 0 );
            to[i] += correction_sign;
            correction -= correction_sign;
            ASSERT( to[i] != 0 );
            if ( to[i] > 1 || correction_sign == 1 )
                double change = log( (double) to[i] / (to[i] + correction_sign) ) * from[i];
                heap.push_back( sort_sym(i,change) );
        ASSERT( cb::sum(to,to+alphabet) == (uint32)to_sum_desired );

You may have noted that the above code is using natural log instead of log2. The difference is only a constant scaling factor, so it doesn't affect the heap order; you may use whatever log base is fastest.

Errkay. So our first attempt is to just use the naive scaling Fs = round( M * Ps ) and then fix the sum using the heap correction algorithm above.

Doing round+correct gets you 99% of the way there. I measured the difference between the total code len made that way and the optimal, and they are less than 0.001 bpb different on every file I tried. But it's still not quite right, so what is the right way?

To guide my search I had a look at the cases where round+correct was not optimal. When it's not optimal it means there is some symbol a and some symbol b such that { Fa+1 , Fb-1 } gives a better total code len than {Fa,Fb}. An example of that is :

count to inc : (1/1024) was (1866/1286152 = 0.0015)
count to dec : (380/1024) was (482110/1286152 = 0.3748)
to inc; cl before : 10.00 cl after : 9.00 , true cl : 9.43
to dec; cl before : 1.43 cl after : 1.43 , true cl : 1.42

The key point is on the 1 count :

count to inc : (1/1024) was (1866/1286152 = 0.0015)
to inc; cl before : 10.00 cl after : 9.00 , true cl : 9.43

1024*1866/1286152 = 1.485660
round(1.485660) = 1

so Fs = 1 , which is a codelen of 10

but Fs = 2 gives a codelen (9) closer to the true codelen (9.43)

And this provided the key observation : rather than rounding the scaled count, what we should be doing is either floor or ceil of the fraction, whichever gives a codelen closer to the true codelen.

BTW before you go off hacking a special case just for Fs==1, it also happens with higher counts :

count to inc : (2/1024) was (439/180084) scaled = 2.4963
to inc; cl before : 9.00 cl after : 8.42 , true cl : 8.68

count to inc : (4/1024) was (644/146557) scaled = 4.4997
to inc; cl before : 8.00 cl after : 7.68 , true cl : 7.83

though obviously the higher Fs, the less likely it is because the rounding gets closer to being perfect.

So it's easy enough just to solve exactly, simply pick the floor or ceil of the ratio depending on which makes the closer codelen :

Ps = Cs/T from the true counts

down = floor( M * Ps )
down = MAX( down,1)

Fs = either down or (down+1)

true_codelen = -log2( Ps )
down_codelen = -log2( down/M )
  up_codelen = -log2( (down+1)/M )

if ( |down_codelen - true_codelen| < |up_codelen - true_codelen| )
  Fs = down
  Fs = down+1

And since all we care about is the inequality, we can do some maths and simplify the expressions. I won't write out all the algebra to do the simplification because it's straightforward, but there are a few key steps :

| log(x) | = log( MAX(x,1/x) )

log(x) >= log(y)  is the same as x >= y

down <= M*Ps
down+1 >= M*Ps

the result of the simplification in code is :

from[] = original counts (Cs) , sum to T
to[] = normalized counts (Fs) , will sum to M

    double from_scaled = from[i] * M/T;

    uint32 down = (uint32)( from_scaled );
    to[i] = ( from_scaled*from_scaled <= down*(down+1) ) ? down : down+1;

Note that there's no special casing needed to ensure that (from_scaled < 1) gives you to[i] = 1 , we get that for free with this expression.

I was delighted when I got to this extremely simple final form.

And that is the conclusion. Use that to find the initial scaled counts. There will still be some correction that needs to be applied to reach the target sum exactly, so use the heap correction algorithm above.

As a final note, if we look at the final expression :

to[i] = ( from_scaled*from_scaled < down*(down+1) ) ? down : down+1;

to[i] = ( test < 0 ) ? down : down+1;

test = from_scaled*from_scaled - down*(down+1); 

from_scaled = down + frac

test = (down + frac)^2 - down*(down+1);

solve for frac where test = 0

frac = sqrt( down^2 + down ) - down

That gives you the fractional part of the scaled count where you should round up or down. It varies with floor(from_scaled). The actual values are :

1 : 0.414214
2 : 0.449490
3 : 0.464102
4 : 0.472136
5 : 0.477226
6 : 0.480741
7 : 0.483315
8 : 0.485281
9 : 0.486833
10 : 0.488088
11 : 0.489125
12 : 0.489996
13 : 0.490738
14 : 0.491377
15 : 0.491933
16 : 0.492423
17 : 0.492856
18 : 0.493242
19 : 0.493589

You can see as Fs gets larger, it goes to 0.5 , so just using rounding is close to correct. It's really in the very low values where it's quite far from 0.5 that errors are most likely to occur.


02-10-14 - Understanding ANS - 9

If you just want to understand the basics of how ANS works, you may skip this post. I'm going to explore some unsolved issues about the sort order.

Some issues about constructing the ANS sort order are still mysterious to me. I'm going to try to attack a few points.

One thing I said wrote last time needs some clarification - "Every slot has an equal probability of 1/M."

What is true is that every character of the output string is equiprobable (assuming again that the Fs are the true probabilities). That is, if you have the string S[] with L symbols, each symbol s occurs Fs times, then you can generate symbols with the correct probability by just drawing S[i] with random i.

The output string S[] also corresponds to the destination state of the encoder in the renormalization range I = [L,2L-1]. What is not true is that all states in I are equally probable.

To explore this I did 10,000 random runs of encoding 10,000 symbols each time. I used L=1024 each time, and gathered stats from all the runs.

This is the actual frequency of the state x having each value in [1024,2047] (scaled so that the average is 1000) :

The lowest most probable states (x=1024) have roughly 2X the frequency of the high least probable states (x=2047).

Note : this data was generated using Duda's "precise initialization" (my "sort by sorting" with 0.5 bias). Different table constructions will create different utilization graphs. In particular the various heuristics will have some weird bumps. And we'll see what different bias does later on.

This is the same data with 1/X through it :

This probability distribution (1/X) can be reproduced just from doing this :

            x = x*b + irandmod(b); // for any base b
            while( x >= 2*K ) x >>= 1;
            stats_count[x-K] ++;            

though I'd still love to see an analytic proof and understand that better.

So, the first thing I should correct is : final states (the x' in I) are not equally likely.

How that should be considered in sort construction, I do not know.

The other thing I've been thinking about was why did I find that the + 1.0 bias is better in practice than the + 0.5 bias that Duda suggests ("precise initialization") ?

What the +1 bias does is push low probability symbols further towards the end of the sort order. I've been contemplating why that might be good. The answer is not that the end of the sort order makes longer codelens, because that kind of issue has already been accounted for.

My suspicion was that the +1 bias was beating the +0.5 bias because of the difference between normalized counts and unnormalized original counts.

Recall that to construct the table we had to make normalized frequences Fs that sum to L. These, however, are not the true symbol frequencies (except in synthetic tests). The true symbol frequencies had to be scaled to sum to L to make the Fs.

The largest coding error from frequency scaling is on the least probable symbols. In fact the very worst case is symbols that occur only once in a very large file. eg. in a 1 MB file a symbol occurs once; its true probability is 2^-20 and it should be coded in 20 bits. But we scale the frequencies to sum to 1024 (for example), it still must get a count of 1, so it's coded in 10 bits.

What the +1 bias does is take the least probable symbols and push them to the end of the table, which maximizes the number of bits they take to code. If the {Fs} were the true frequencies, this would be bad, and the + 0.5 bias would be better. But the {Fs} are not the true frequencies.

This raises the question - could we make the sort order from the true frequencies instead of the scaled ones? Yes, but you would then have to either transmit the true frequencies to the decoder, or transmit the sort order. Either way takes many more bits than transmitting the scaled frequencies. (in fact in the real world you may wish to transmit even approximations of the scaled frequencies). You must ensure the encoder and decoder use the same frequencies so they build the same sort order.

Anyway, I tested this hypothesis by making buffers synthetically by drawing symbols from the {Fs} random distribution. I took my large testset, for each file I counted the real histogram, made the scaled frequencies {Fs}, then regenerated the buffer from the frequencies {Fs} so that the statistics match the data exactly. I then ran tANS on the synthetic buffers and on the original file data :

synthetic data :

total bytes out : 146068969.00  bias=0.5
total bytes out : 146117818.63  bias=1.0

real data :

total bytes out : 144672103.38  bias=0.5
total bytes out : 144524757.63  bias=1.0

On the synthetic data, bias=0.5 is in fact slightly better. On the real data, bias=1.0 is slightly better. This confirms that the difference between the normalized counts & unnormalized counts is in fact the origin of 1.0's win in my previous tests, but doesn't necessarily confirm my guess for why.

An idea for an alternative to the bias=1 heuristic is you could use bias=0.5 , but instead of using the Fs for the sort order, use the estimated original count before normalization. That is, for each Fs you can have a probability model of what the original count was, and select the maximum-likelihood count from that. This is exactly analoguous to restoring to expectation rather than restoring to middle in a quantizer.

Using bias=1.0 and measuring state occurance counts, we get this :

Which mostly has the same 1/x curve, but with a funny tail at the end. Note that these graphs are generated on synthetic data.

I'm now convinced that the 0.5 bias is "right". It minimizes measured output len on synthetic data where the Fs are the true frequencies. It centers each symbol's occurances in the output string. It reproduces the 1/x distribution of state frequencies. However there is still the missing piece of how to derive it from first principles.


While I was at it, I gathered the average number of bits output when coding from each state. If you're following along with Yann's blog he's been explaining FSE in terms of this. tANS outputs bits to get the state x down into the coding range Is for the next symbol. The Is are always lower than I (L), so you have to output some bits to scale down x to reach the Is. x starts in [L,2L) and we have to output bits to reach [Fs,2Fs) ; the average number of bits required is like log2(L/Fs) which is log2(1/P) which is the code length we want. Because our range is [L,2L) we know the average output bit count from each state must differ by 1 from the top of the range to the bottom. In fact it looks like this :

Another way to think about it is that at state=L , the state is empty. As state increases, it is holding some fractional bits of information in the state variable. That number of fraction bits goes from 0 at L up to 1 at 2L.

Ryg just pointed me at a proof of the 1/x distribution in Moffat's "Arithmetic Coding Revisited" (DCC98).

The "x" in ANS has the same properties as the range ("R") in an arithmetic coder.

The bits of information in x is I ~= log( x )

I is in [0,1] and is a uniform random value, Pr(I) ~= 1

if log(x) has Pr ~= 1 , then Pr(x) must be ~= 1/x

The fact that I is uniform is maybe not entirely obvious; Moffat just hand-waves about it. Basically you're accumulating a random variable into I ( -log2(P_sym) ) and then dropping the integer part; the result is a fractional part that's random and uniform.


02-06-14 - Understanding ANS - 8

Time to address an issue that we've skirted for some time - how do you make the output string sort order?

Recall : The output string contains Fs occurances of each symbol. For naive rANS the output string is just in alphabetical order (eg. "AAABBBCCC"). With tANS we can use any permutation of that string.

So what permutation should we use? Well, the output string determines the C() and D() encode and decode tables. It is in fact the only degree of freedom in table construction (assuming the same constraints as last time, b=2 and L=M). So we should choose the output string to minimize code length.

The guiding principle will be (x/P). That is, we achieve minimum total length when we make each code length as close to log2(1/P) as possible. We do that by making the input state to output state ratio (x'/x) as close to (1/P) as possible.

(note for the record : if you try to really solve to minimize the error, it should not just be a distance between (x'/x) and (1/P) , it needs to be log-scaled to make it a "minimum rate" solution). (open question : is there an exact solution for table building that finds the minimum rate table that isn't NP (eg. not just trying all permutations)).

Now we know that the source state always come from the precursor ranges Is, and we know that

destination range :
I = [ M , 2*M - 1]

source range :
Is = [ Fs, 2*Fs - 1 ] for each symbol s

and Ps = Fs/M

so the ideal target for the symbols in each source range is :

target in I = (1/Ps) * (Is) = (M/Fs) * [ Fs, 2*Fs - 1 ] = 

and taking off the +M bias to make it a string index in the range [0,M-1] :

Ts = target in string = target in I - M

Ts = { 0 , M * 1/Fs , M * 2/Fs) , ... }

Essentially, we just need to take each symbol and spread its Fs occurances evenly over the output string.

Now there's a step that I don't know how to justify without waving my hands a bit. It works slightly better if we imagine that the source x was not just an integer, but rather a bucket that covers the unit range of that integer. That is, rather that starting exactly at the value "x = Fs" you start in the range [Fs,Fs+1]. So instead of just mapping up that integer by 1/P we map up the range, and we can assign a target anywhere in that range. In the paper Duda uses a bias of 0.5 for "precise initialization" , which corresponds to assuming the x's start in the middle of their integer buckets. That is :

Ts = { M * (b/Fs), M* (1+b)/Fs, M * (2+b)/Fs , ... }

with b = 0.5 for Duda's "precise initialization". Obviously b = 0.5 makes T centered on the range [0,M] , but I see no reason why that should be preferred.

Now assuming we have these known target locations, you can't just put all the symbols into the target slots that they want, because lots of symbols want the same spot.

For example :


T_A = { 8 * 0.5/3 , 8 * 1.5 / 3 , 8 * 2.5 / 3 } = { 1 1/3 , 4 , 6 2/3 }
T_B = T_A
T_C = { 8 * 0.5/2 , 8 * 1.5/2 } = { 2 , 6 }

One way to solve this problem is to start assigning slots, and when you see that one is full you just look in the neighbor slot, etc. So you might do something like :

initial string is empty :

string = "        "

put A's in 1,4,6

string = " A  A A "

put B's in 1,4,6 ; oops they're taken, shift up one to find empty slots :

string = " AB ABAB"

put C's in 2,6 ; oops they're taken, hunt around to find empty slots :

string = "CABCABAB"

now obviously you could try to improve this kind of algorithm, but there's no point. It's greedy so it makes mistakes in the overall optimization problem (it's highly dependant on order). It can also be slow because it spends a lot of time hunting for empty slots; you'd have to write a fast slot allocator to avoid degenerate bad cases. There are other ways.

Another thing I should note is that when doing these target slot assignments, there's no reason to prefer the most probable symbol first, or the least probable or whatever. The reason is every symbol occurance is equally probable. That is, symbol s has frequency Fs, but there are Fs slots for symbol s, so each slot has a frequency of 1. Every slot has an equal probability of 1/M.

An alternative algorithm that I have found to work well is to sort the targets. That is :

make a sorting_array of size M

add { Ts, s } to sorting_array for each symbol  (that's Fs items added)

sort sorting_array by target location

the symbols in sorting_array are in output string order

I believe that this is identical to Duda's "precise initialization" which he describes using push/pop operations on a heap; the result is the same - assigning slots in the order of desired target location.

Using the sort like this is a little weird. We are no longer explicitly trying to put the symbols in their target slots. But the targets (Ts) span the range [0, M] and the sort is an array of size M, so they wind up distributed over that range. In practice it works well, and it's fast because sorting is fast.

A few small notes :

You want to use a "stable" sort, or bias the target by some small number based on the symbol. The reason is you will have lots of ties, and you want the ties broken consistently. eg. for "AABBCC" you want "ABCABC" or "CBACBA" but not "ABCCAB". One way to get a stable sort is to make the sorting_array work on U32's, and pack the sort rank into the top 24 bits and the symbol id into the bottom 8 bits.

The bias = 0.5 that Duda uses is not strongly justified, so I tried some other numbers. bias = 0 is much worse. It turns out that bias = 1.0 is better. I tried a bunch of values on a large test set and found that bias = 1 is consistently good.

One very simple way to get a decent sort is to bit-reverse the rANS indexes. That is, start from a rANS/alphabetical order string ("AABB..") and take the index of each element, bit-reverse that index (so 0001 -> 1000) , and put the symbol in the bit reversed slot. While this is not competitive with the proper sort, it is simple and one pass.

Another possible heuristic is to just scatter the symbols by doing steps that are prime with M. This is what Yann does in fse.c

All the files in Calgary Corpus :
(compression per file; sum of output sizes)

M = 1024

rANS/alpahabetical : 1824053.75

bit reverse : 1805230.75

greedy search for empty slots : 1801351

Yann's heuristic in fse.c : 1805503.13

sort , bias = 0.0 : 1817269.88

sort , bias = 0.5 : 1803676.38  (Duda "precise")

sort , bias = 1.0 : 1798930.75

Before anyone throws a fit - yes, I tested on my very large test set, not just calgary. The results were consistent on all the test sets I tried. I also tested with larger M (4096) and the results were again the same, though the differences are smaller the larger you make M.

For completeness, here is what the sorts actually do :


bit reverse :   ABABABACABACABBC

greedy search : CABABACABABACABB

greedy search, LPS first :  ABCABAACBABACBAB

Yann fse :          AAABBCAABBCAABBC

sort , bias = 0.0 : ABCABABCABABCABA

sort , bias = 0.5 : ABCABABACBABACBA

sort , bias = 1.0 : ABABCABABCABAABC

but I caution against judging sorts by whether they "look good" since that criteria does not seem to match coding performance.

Finally for clarity, here's the code for the simpler sorts :

void make_sort(int * sorted_syms, int sort_count, const uint32 * normalized_counts, int alphabet)
    ASSERT( (int) cb::sum(normalized_counts,normalized_counts+alphabet) == sort_count );
    const int fse_step = (sort_count>>1) + (sort_count>>3) + 1;
    int fse_pos = 0;
    int s = 0;
    for LOOP(a,alphabet)
        int count = normalized_counts[a];

        for LOOP(c,count)
            // choose one :

            // rANS :
            sorted_syms[s] = a;

            // fse :
            sorted_syms[fse_pos] = a;
            fse_pos = (fse_pos + step) % sort_count;

            // bitreverse :
            sorted_syms[ bitreverse(s, numbits(sort_count)) ] = a;


and the code for the actual sorting sort (recommended) :

struct sort_sym
    int sym;
    float rank;
    bool operator < (const sort_sym & rhs) const
        return rank < rhs.rank;

void make_sort(int * sorted_syms, int sort_count, const uint32 * normalized_counts, int alphabet)
    ASSERT( (int) cb::sum(normalized_counts,normalized_counts+alphabet) == sort_count );

    vector<sort_sym> sort_syms;

    int s = 0;

    for LOOP(sym,alphabet)
        uint32 count = normalized_counts[sym];
        if ( count == 0 ) continue;
        float invp = 1.f / count;
        float base =  1.f * invp; // 0.5f for Duda precise

        for LOOP(c,(int)count)
            sort_syms[s].sym = sym;
            sort_syms[s].rank = base + c * invp;
    ASSERT_RELEASE( s == sort_count );
    for LOOP(s,sort_count)
        sorted_syms[s] = sort_syms[s].sym;

and for the greedy search :

void make_sort(int * sorted_syms, int sort_count, const uint32 * normalized_counts, int alphabet)
    ASSERT( (int) cb::sum(normalized_counts,normalized_counts+alphabet) == sort_count );

    // make all slots empty :
    for LOOP(s,sort_count)
        sorted_syms[s] = -1;
    for LOOP(a,alphabet)
        uint32 count = normalized_counts[a];
        if ( count == 0 ) continue;
        uint32 step = (sort_count + (count/2) ) / count;
        uint32 first = step/2;
        for LOOP(c,(int)count)
            uint32 slot = first + step * c;
            // find an empty slot :
                if ( sorted_syms[slot] == -1 )
                    sorted_syms[slot] = a;
                slot = (slot + 1)%sort_count;

small note : the reported results use a greedy search that searches away from slot using +1,-1,+2,-2 , instead of the simpler +1,+2 in this code snippet. This simpler version is very slightly worse.


02-05-14 - Understanding ANS - 7

And we're ready to cover table-based ANS (or "tANS") now.

I'm going to be quite concrete and consider a specific choice of implementation, rather than leaving everything variable. But extrapolation to the general solution is straightforward.

You have integer symbol frequences Fs. They sum to M. The cumulative frequencies are Bs.

I will stream the state x in bits. I will use the smallest possible renormalization range for this example , I = [ M , 2*M - 1]. You can always use any integer multiple of M that you want (k*M, any k), which will give you more coding resolution (closer to entropy). This is equivalent to scaling up all the F's by a constant factor, so it doesn't change the construction here.

Okay. We will encode/decode symbols using this procedure :

ENCODE                      DECODE

|                           ^
V                           |

stream out                  stream in

|                           ^
V                           |

C(s,x) coding function      D(x) decoding function

|                           ^
V                           |

x'                          x'

We need tables for C() and D(). The constraints are :

D(x') = { x , s }  outputs a state and a symbol

D(x) must be given for x in I = [ M , 2*M - 1 ]

D(x) in I must output each symbol s Fs times

that is, D(x in I) must be an output string made from a permutation of "AA..BB.." , each symbol Fs times

D( C( s, x ) ) = { x , s }  decode must invert coding

C(s,x) = x'  outputs the following state

C(s,x) must be given for x' in I
 that's x in Is

The precursor ranges Is = { x : C(s,x) is in I }
must exist and be of the form Is = [ k , 2k-1 ] for some k

Now, if we combine the precursor range requirement and the invertability we can see :

D(x in I) outputs each s Fs times

C(s,x) with x' in I must input each s Fs times

the size of Is must be Fs

the precursor ranges must be Is = [ Fs, 2*Fs - 1 ]

C(s,x) must given in M slots

And I believe that's it; those are the necessary and sufficient conditions to make a valid tANS system. I'll go over some more points and fill in some details.

Here's an example of the constraint for an alphabet of "ABC" and M = 8 :

Now, what do you put in the shaded area? You just fill in the output states from 8-15. The order you fill them in corresponds to the output string. In this case the output string must be some permutation of "AAABBBCC".

Here's one way : (and in true Duda style I have confusingly used different notation in the image, since I drew this a long time ago before I started this blog series. yay!)

In the image above I have also given the corresponding output string and the decode table. If you're following along in Duda's paper arxiv 1311.2540v2 this is figure 9 on page 18. What you see in figure 9 is a decode table. The "step" part of figure 9 is showing one method of making the sort string. The shaded bars on the right are showing various permutations of an output string, with a shading for each symbol.

Before I understood ANS I was trying tables like this :

Fs = {7,6,3}

 S |  0|  1|  2
  1|  2|  3|  4
  2|  5|  6| 10
  3|  7|  8| 15
  4|  9| 11| 20
  5| 12| 13| 26
  6| 14| 16| 31
  7| 17| 19|   
  8| 18| 22|   
  9| 21| 24|   
 10| 23| 27|   
 11| 25| 29|   
 12| 28|   |   
 13| 30|   |   

This table does not work. If you're in state x = 7 and you want to encode symbol 2, you need to stream out bits to get into the precursor range I2. So you stream out from x=7 and get to x=3. Now you look in the table and you are going to state 15 - that's not in the range I=[16,31]. No good!

A correct table for those frequencies is :

 S |  0|  1|  2
  3|   |   | 18
  4|   |   | 24
  5|   |   | 29
  6|   | 17|   
  7| 16| 20|   
  8| 19| 22|   
  9| 21| 25|   
 10| 23| 27|   
 11| 26| 31|   
 12| 28|   |   
 13| 30|   |   

Building the decode table from the encode table is trivial.

Note that the decode table D(x) only has to be defined for x in I - that's M entries.

C(x,s) also only has M entries. If you made it naively as a 2d array, it would be |alphabet|*M . eg. something like (256*4096) slots, but most of it would be empty. Of course you don't want to do that.

The key observation is that C(x,s) is only defined over consecutive ranges of x for each s. In fact it's defined over [Fs, 2*Fs-1]. So, we can just pack these ranges together. The starting point in the packed array is just Bs - the cumulative frequency of each symbol. That is :

PC = packed coding table
PC has M entries

C(x,s) = PC[ Bs + (x - Fs) ]

eg. for the {3,3,2} table shown in the image above :

PC = { 8,11,14, 9,12,15, 10,13 }

this allows you to store the coding table also in an array of size M.

There are a few topics on tANS left to cover but I'll leave them for the next post.


02-04-14 - Understanding ANS - 6

Okay, let's get to streaming.

For illustration let's go back to the simple example of packing arbitrary base numbers into an integer :

// encode : put val into state
void encode(int & state, int val, int mod)
    ASSERT( val >= 0 && val < mod );
    state = state*mod + val;

// decode : remove a value from state and return it
int decode(int & state, int mod )
    int val = state % mod;
    state /= mod;
    return val;

as you encode, state grows, and eventually gets too big to fit in an integer. So we need to flush out some bits (or bytes).

But we can't just stream out bits. The problem is that the decoder does a modulo to get the next value. If we stream in and out high bits, that's equivalent to doing something like +65536 on the value. When you do a mod-3 (or whatever) on that, you have changed what you decode.

If you only ever did mod-pow2's, you could stream bits out of the top at any time, because the decoding of the low bits is not affected by the high bits. This is how the Huffman special case of ANS works. With Huffman coding you can stream in and out any bits that are above the current symbol, because they don't affect the mask at the bottom.

In general we want to stream bits (base 2) or bytes (base 256). To do ANS in general we need to mod and multiply by arbitrary values that are not factors of 2 or 256.

To ensure that we get decodability, we have to stream such that the decoder sees the exact value that the encoder made. That is :

ENCODE                      DECODE

|                           ^
V                           |

stream out                  stream in

|                           ^
V                           |

C(s,x) coding function      D(x) decoding function

|                           ^
V                           |

x'                          x'

The key thing is that the value of x' that C(s,x) makes is exactly the same that goes into D(x).

This is different from Huffman, as noted above. It's also different than arithmetic coding, which can have an encoder and decoder that are out of sync. An arithmetic decoder only uses the top bits, so you can have more or less of the rest of the stream in the low bits. While the basic ANS step (x/P + C) is a kind of arithmetic coding step, the funny trick we did to take some bits of x and mod it back down to the low bits (see earlier posts) means that ANS is *not* making a continuous arithmetic stream for the whole message that you can jump into anywhere.

Now it's possible there are multiple streaming methods that work. For example with M = a power of 2 in rANS you might be able to stream high bytes. I'm not sure, and I'm not going to talk about that in general. I'm just going to talk about one method of streaming that does work, which Duda describes.

To ensure that our encode & decode streaming produce the same value of x', we need a range to keep it in. If you're streaming in base b, this range is of the form [L, b*L-1] . So, I'll use Duda's terminology and call "I" the range we want x' to be in for decoding, that is

I = [L, b*L-1]

Decoder streams into x :

x <- x*b + get_base_b();

until x is in I

but the encoder must do something a bit funny :

stream out from x

x' = C(s,x)  , coding function

x' now in I

that is, the stream out must be done before the coding function, and you must wind up in the streaming range after the coding function. x' in the range I ensures that the encoder and decoder see exactly the same value (because any more streaming ops would take it out of I).

To do this, we must know the "precursor" ranges for C(). That is :

Is = { x : C(s,x) is in I }

that is, the values of x such that after coding with x' = C(s,x), x' is in I

these precursor ranges depend on s. So the encoder streaming is :

I'm about to encode symbol s

stream out bits from x :

put_base_b( x % b )
x <- x/b

until x is in Is

so we get into the precursor range, and then after the coding step we are in I.

Now this is actually a constraint on the coding function C (because it determines what the Is are). You must be able to encode any symbol from any state. That means you must be able to reach the Is precursor range for any symbol from any x in the output range I. For that to be true, the Is must span a power of b, just like "I" does. That is,

all Is must be of the form

Is = [ K, b*K - 1 ]

for some K

eg. to be concrete, if b = 2, we're streaming out bits, then Is = { 3,4,5 } is okay, you will be able to get there from any larger x by streaming out bits, but Is = {4,5,6} is not okay.

I = [8, 15]

Is = {4,5,6}

x = 14

x is out of Is, so stream out a bit ; 14 -> 7

x is out of Is, so stream out a bit ; 7 -> 3

x is below Is!  crap!

this constraint will be our primary guide in building the table-based version of ANS.


02-03-14 - Understanding ANS - 5

First in case you aren't following them already, you should follow along with ryg and Yann as we all go through this :
RealTime Data Compression A comparison of Arithmetic Encoding with FSE
rANS notes The ryg blog

Getting back to my slow exposition track.

We talked before about how strings like "ABC" specify an ANS state machine. The string is the symbols that should be output by the decoder in each state, and there's an implicit cyclic repeat, so "ABC" means "ABCABC..". The cyclic repeat corresponds to only using some modulo of state in the decoder output.

Simple enumerations of the alphabet (like "ABC") are just flat codes. We saw before that power of two binary-distributed strings like "ABAC" are Huffman codes.

What about something like "AAB" ? State 0 and 1 both output an A. State 2 outputs a B. That means A should have twice the probability of B.

How do we encode a state like that? Putting in a B is obvious, we need to make the bottom of state be a 2 (mod 3) :

x' = x*3 + 2

but to encode an A, if we just did the naive op :

x' = x*3 + 0
x' = x*3 + 1

we're wasting a value. Either a 0 or 1 at the bottom would produce an A, so we have a free bit there. We need to make use of that bit or we are wasting code space. So we need to find a random bit to transmit to make use of that freedom. Fortunately, we have a value sitting around that needs to be transmitted that we can pack into that bit - x!

take a bit off x :
b = x&1
x >>= 1

x' = x*3 + b

or :

x' = (x/2)*3 + (x%2)

more generally if the output string is of length M and symbol s occurs Fs times, you would do :

x' = (x/Fs)*M + (x%Fs) + Bs

which is the formula for rANS.

Now, note that rANS always makes output strings where the symbols are not interleaved. That is, it can make "AAB" but it can't make "ABA". The states that output the same symbol are in consecutive blocks of length Fs.

This is actually not what we want, it's an approximation in rANS.

For example, consider a 3-letter alphabet and M=6. rANS corresponds to an output string of "AABBCC". We'd prefer "ABCABC". To see why, recall the arithmetic coding formula that we wanted to use :

x' = (x/P) + C

the important part being the (x/P). We want x to grow by that amount, because that's what gives us compression to the entropy. If x grows by too much or too little, we aren't coding with codelens that match the probabilities, so there will be some loss.

P = F/M , and we will assume for now that the probabilities are such that the rational expression F/M is the true probability. What we want is to do :

x' = (x*M)/F + C

to get a more accurate scaling of x. But we can't do that because in general that division will cause x' to not fall in the bucket [Cs,Cs+1) , which would make decoding impossible.

So instead, in rANS we did :

x' = floor(x/F)*M + C + (x%F)

the key part here being that we had to do floor(x/F) instead of (x/P), which means the bottom bits of x are not contributing to the 1/P scaling the way we want them to.


x = 7
F = 2
M = 6
P = 1/3

should scale like

x -> x/P = 21

instead scales like

x -> (x/F)*M + (x%F) = (7/2)*6 + (7%2) = 3*6 + 1 = 19

too low
because we lost the bottom bit of 7 when we did (7/2)

In practice this does in fact make a difference when the state value (x) is small. When x is generally large (vs. M), then (x/F) is a good approximation of the correct scaling. The closer x is to M, the worse the approximation.

In practice with rANS, you should use something like x in 32-bits and M < 16-bits, so you have decent resolution. For tANS we will be using much lower state values, so getting this right is important.

As a concrete example :

alphabet = 3
1000 random symbols coded
H = 1.585
K = 6
6 <= x < 12

output string "ABCABC"
wrote 1608 bits

output string "AABBCC"
wrote 1690 bits

And a drawing of what's happening :

I like the way Jarek Duda called rANS the "range coder" variant of ANS. While the analogy is not perfect, there is a similarity in the way it approximates the ideal scaling and gains speed.

The crucial difference between a "range coder" and prior (more accurate) arithmetic coders is that the range coder does a floor divide :

range coder :

symhigh * (range / symtot)

CACM (and such) :

(symhigh * range) / symtot

this approximation is good as long as range is large and symtot is small, just like with rANS.


02-02-14 - Understanding ANS - 4

Another detour from the slow exposition to mention something that's on my mind.

Let's talk about arithmetic coding.

Everyone is familiar with the standard simplified idea of arithmetic coding. Each symbol has a probability P(s). The sum of all preceding probability is the cumulative probability, C(s).

You start with an interval in [0,1]. Each symbol is specified by a range equal to P(s) located at C(s). You reduce your interval to the range of the current symbol, then put the next symbol within that range, etc. Like this :

As you go, you are making a large number < 1, with more and more bits of precision being added at the bottom. In the end you have to send enough bits so that the stream you wanted is specified. You get compression because more likely streams have larger intervals, and thus can be specified with fewer bits.

In the end we just made a single number that we had to send :

x = C0 + P0 * ( C1 + P1 * (C2 ...

in order to make that value in a FIFO stream, we would have to use the normal arithmetic coding style of tracking a current low and range :

currenty at [low,range]
add in Cn,Pn

low <- low + range * Cn
range <- range * Pn

and of course for the moment we're assuming we get to use infinite precision numbers.

But you can make the same final value x another way. Start at the end of the stream and work backwards, LIFO :


x contains all following symbols already encoded
x in [0,1]

x' <- Cn + Pn * x

there's no need to track two variables [low,range], you work from back to front and then send x to specify the whole stream. (This in fact is an ancient arithmetic coder. I think it was originally described by Elias even before the Pasco work. I mention this to emphasize that single variable LIFO coding is nothing new, though the details of ANS are in fact quite new. Like "range coding" vs prior arithmetic coders , it can be the tiny details that make all the difference.) (umm, holy crap, I just noticed the ps. at the bottom of that post ("ps. the other new thing is the Jarek Duda lattice coder thing which I have yet to understand")).

    ADD 12/9 : citation :
    This is the clearest one I've found : ibmrd2302G.pdf "Arithmetic Coding" , Rissanen & Langdon, IBM J Res. D, March 1979 The "dual" Elias code uses a single state variable and is the same as the LIFO arithmetic coder I describe. The non-dual is the forward two-variable version. In equation 4, P is a cumulative probability (my C), and l (lower case) is the length of a codeword. Let l(k)=-log2(p(k)) then C_bar <- P(k) + p(k) * C_bar is the single-variable LIFO Elias coder.

You can decode an individual step thusly :

x in [0,1]
find s such that C(s) <= x < C(s+1)

x' <- (x - Cs)/Ps

Now let's start thinking about doing this in finite precision, or at least fixed point.

If we think of our original arithmetic coding image, growing "x" up from 0, instead of keeping x in [0,1] the whole time, let's keep the active interval in [0,1]. That is, as we go we rescale so that the bottom range is [0,1] :

That is, instead of keeping the decimal at the far left and making a fraction, we keep the decimal at the far right of x and we grow the number upwards. Each coding step is :

x' <- x/Ps + Cs

in the end we get a large number that we have to send using log2(x) bits. We get compression because highly probable symbols will make x grow less than improbable symbols, so more compressable streams will make smaller values of x.

(x = the value before the current step, x' = the value after the current step)

We can decode each step simply if the (x/Ps) are integers, then the Cs is the fractional part, so we just do :

f = frac(x)
find s such that C(s) <= f < C(s+1)

x' <- (x - Cs)*Ps
that is, we think of our big number x as having a decimal point with the fractional part Cs on the right, and the rest of the stream is in a big integer on the left. That is :

[ (x/Ps) ] . [ Cs ]

Of course (x/Ps) is not necessarily an integer, and we don't get to do infinite precision math, so let's fix that.

Let :

Ps = Fs / M

P is in [0,1] , a symbol probability
F is an integer frequency
M is the frequency denominator

Sum{F} = M

Cs = Bs / M

B is the cumulative frequency

now we're going to keep x an integer, and our "decimal" that separates the current symbol from the rest of the stream is a fixed point in the integer. (eg. if M = 2^12 then we have a 12 bit fixed point x).

Our coding step is now :

x' = x*M/Fs + Bs

and we can imagine this as a fixed point :

[ (x*M/Fs) ] . [ Bs ]

in particular the bottom M-ary fraction specifies the current symbol :

( x' mod M ) = Bs

the crucial thing for decoding is that the first part, the (x/P) part which is now (x*M/F) shouldn't mess up the bottom M-ary fraction.

But now that we have it written like this, it should be obvious how to do that, if we just write :

x*M/F -> floor(x/F)*M

then the (mod M) operator on that gives you 0, because it has an explicit *M

so we've made the right (x/P) scaling, and made something that doesn't mess up our bottom mod M for decodability.

But we've lost some crucial bits from x, which contains the rest of the stream. When we did floor(x/F) we threw out some bottom bits of x that we can't get rid of. So we need that (x mod F) back.

Fortunately we have the perfect place to put it. We can specify the current symbol not just with Bs, but with anything in the interval [Bs , Bs + Fs) ! So we can do :

x' = M*floor(x/Fs) + Bs + (x mod Fs)

which is :

x' = [ floor(x/Fs) ] . [ Bs + (x mod Fs) ]

with the integer part growing on the left and the base-M fractional part on the right specifying the current symbol s

and this is exactly the rANS encoding step !

As we encode x grows by (1/P) with each step. We wind up sending x with log2(x) bits, which means the code length of the stream is log2(1/P0*P1...) which is what we want.

For completeness, decoding is straightforwardly undoing the encode step :

f = M-ary fraction of x  (x mod M)
find s such that Bs <= f < Bs+1

x' = Fs * (x/M) + (x mod M) - Bs


x' = Fs * intM(x) + fracM(x) - Bs

and we know

(fracM(x) - Bs) is in [0, Fs)

which is the same as the old arithmetic decode step : x' = (x - Cs)/Ps

Of course we still have to deal with the issue of keeping x in fixed width integers and streaming, which we'll come back to.


02-01-14 - Understanding ANS - 3

I'm gonna take an aside from the slow exposition and jump way ahead to some results. Skip to the bottom for summary.

There have been some unfortunate claims made about ANS being "faster than Huffman". That's simply not true. And in fact it should be obvious that it's impossible for ANS to be faster than Huffman, since ANS is a strict superset of Huffman. You can always implement your Huffman coder by putting the Huffman code tree into your ANS coder, therefore the speed of Huffman is strictly >= ANS.

In practice, the table-based ANS decoder is so extremely similar to a table-based Huffman decoder that they are nearly identical in speed, and all the variation comes from minor implementation details (such as how you do your bit IO).

The "tANS" (table ANS, aka FSE) decode is :

  int sym = decodeTable[state].symbol;
  *op++ = sym;
  int nBits = decodeTable[state].nbBits;
  state = decodeTable[state].newState + getbits(nBits);

while a standard table-based Huffman decode is :

  int sym = decodeTable[state].symbol;
  *op++ = sym;
  int nBits = codelen[sym];
  state = ((state<<nBits)&STATE_MASK) + getbits(nBits);  

where for similarly I'm using a Huffman code with the first bits at the bottom. In the Huffman case, "state" is just a portion of the bit stream that you keep in a variable. In the ANS case, "state" is a position in the decoder state machine that has memory; this allows it to carry fractional bits forward.

If you so chose, you could of course put the Huffman codelen and next state into decodeTable[] just like for ANS and they would be identical.

So, let's see some concrete results comparing some decent real world implementations.

I'm going to compare four compressors :

huf = order-0 static Huffman

fse = Yann's implementation of tANS

rans = ryg's implementation of rANS

arith = arithmetic coder with static power of 2 cumulative frequency total and decode table

For fse, rans, and arith I use a 12-bit table (the default in fse.c)
huf uses a 10-bit table and does not limit code length

Runs are on x64 code, but the implementations are 32 bit. (no 64-bit math used)

Some notes on the four implementations will follow. First the raw results :

inName : book1
H = 4.527

arith 12:   768,771 ->   435,378 =  4.531 bpb =  1.766 to 1 
arith encode     : 0.006 seconds, 69.44 b/kc, rate= 120.08 mb/s
arith decode     : 0.011 seconds, 40.35 b/kc, rate= 69.77 mb/s

"rans 12:   768,771 ->   435,378 =  4.531 bpb =  1.766 to 1 
rans encode      : 0.010 seconds, 44.04 b/kc, rate= 76.15 mb/s
rans decode      : 0.006 seconds, 80.59 b/kc, rate= 139.36 mb/s

fse :   768,771 ->   435,981 =  4.537 bpb =  1.763 to 1 
fse encode       : 0.005 seconds, 94.51 b/kc, rate= 163.44 mb/s
fse decode       : 0.003 seconds, 166.95 b/kc, rate= 288.67 mb/s

huf :   768,771 ->   438,437 =  4.562 bpb =  1.753 to 1 
huf encode       : 0.003 seconds, 147.09 b/kc, rate= 254.34 mb/s
huf decode       : 0.003 seconds, 163.54 b/kc, rate= 282.82 mb/s
huf decode       : 0.003 seconds, 175.21 b/kc, rate= 302.96 mb/s (*1)

inName : pic
H = 1.210

arith 12:   513,216 ->    79,473 =  1.239 bpb =  6.458 to 1 
arith encode     : 0.003 seconds, 91.91 b/kc, rate= 158.90 mb/s
arith decode     : 0.007 seconds, 45.07 b/kc, rate= 77.93 mb/s

rans 12:   513,216 ->    79,474 =  1.239 bpb =  6.458 to 1 
rans encode      : 0.007 seconds, 45.52 b/kc, rate= 78.72 mb/s
rans decode      : 0.003 seconds, 96.50 b/kc, rate= 166.85 mb/s

fse :   513,216 ->    80,112 =  1.249 bpb =  6.406 to 1 
fse encode       : 0.003 seconds, 93.86 b/kc, rate= 162.29 mb/s
fse decode       : 0.002 seconds, 164.42 b/kc, rate= 284.33 mb/s

huf :   513,216 ->   106,691 =  1.663 bpb =  4.810 to 1 
huf encode       : 0.002 seconds, 162.57 b/kc, rate= 281.10 mb/s
huf decode       : 0.002 seconds, 189.66 b/kc, rate= 328.02 mb/s

And some conclusions :

1. "tANS" (fse) is almost the same speed to decode as huffman, but provides fractional bits. Obviously this is a huge win on skewed files like "pic". But even on more balanced distributions, it's a decent little compression win for no decode speed hit, so probably worth doing everywhere.

2. Huffman encode is still significantly faster than tANS encode.

3. "rANS" and "arith" almost have their encode and decode speeds swapped. Round trip time is nearly identical. They use identical tables for encode and decode. In fact they are deeply related, which is something we will explore more in the future.

4. "tANS" is about twice as fast as "rANS". (at this point)

And some implementation notes for the record :

"fse" and "rans" encode the array by walking backwards.  The "fse" encoder output bits forwards and
consume them backwards, while the "rans" encoder writes bits backwards and consumes them forwards.

"huf" and "fse" are transmitting their code tables.  "arith" and "rans" are not.  
They should add about 256 bytes of header to be fair.

"arith" is a standard Schindler range coder, with byte-at-a-time renormalization

"arith" and "rans" here are nearly identical, both byte-at-a-time, and use the exact same tables
for encode and decode.

All times include their table-build times, and the time to histogram and normalize counts.
If you didn't include those times, the encodes would appear faster.

"huf" here is not length-limited.  A huf decoder with a 12-bit table and 12-bit length limitted
codes (like "fse" uses) should be faster.
(*1 = I did a length-limited version with a non-overflow handling decoder)

"huf" here is was implemented with PowerPC and SPU in mind.  A more x86/x64 oriented version should be
a little faster.  ("fse" is pretty x86/x64 oriented).

and todo : compare binary rANS with a comparable binary arithmetic coder.

old rants