# Considering the order of results when computing Cartesian product [short]

Sometimes in programming you need to do a pairwise comparison of some elements coming from two collections, for example, checking possible collisions between particles (which may be embedded inside a quadtree representation for efficiency). A handy operation is then the Cartesian product of the two sets of elements, to get the set of all pairs, which can then be traversed.

Whenever I need a Cartesian product of two lists in Haskell, I whip out the list monad to generate the Cartesian product:

```cartesianProduct :: [a] -> [b] -> [(a, b)]
cartesianProduct as bs = as >>= (\a -> bs >>= (\b -> [(a, b)]))```

Thus for every element `a` in `as` we get every element `b` in `bs` and form the singleton list of the pair `(a, b)`, all of which get concatenated to produce the result. For example:

```*Main> cartesianProduct [1..4] [1..4]
[(1,1),(1,2),(1,3),(1,4),(2,1),(2,2),(2,3),(2,4),(3,1),(3,2),(3,3),(3,4),(4,1),(4,2),(4,3),(4,4)]```

This traverses the Cartesian space in row order. That is, if we imagine the square grid of possibilities here (where the element at the ith row and jth column corresponds to element `(i, j)`), then `cartesianProduct` generates the pairs in the following order:

In mathematics, the Cartesian product is defined on sets, so the order of these pairs is irrelevant. Indeed, if we want to examine all pairs, then it may not matter in what order. However, if we learn something as we look at the pairs (i.e., accumulating some state), then the order can be important.

In some recent research, I was building an automated algebra tool to find axioms for some algebraic structure, based on an interpretation. This involved generating all pairs of syntax trees `t` and `t'` of terms over the algebraic structure, up to a particular depth, and evaluating whether `t = t'`. I also had some automated proof search machinery to make sure that this equality wasn’t already derivable from the previously generated axioms. I could have done this derivability check as a filtering afterwards, but I was exploring a very large space, and did not expect to even be able to generate all the possibilities. I just wanted to let the thing run for a few hours and see how far it got. Therefore, I needed Cartesian product to get all pairs, but the order in which I generated the pairs became important for the effectiveness of my approach.  The above ordering (row major order) was not so useful as I was unlikely to find interesting axioms quickly by traversing the span of a (very long) row; I needed to unpack the space in a more balanced way.

My first attempt at a more balanced Cartesian product was the following:

```cartesianProductBalanced :: [a] -> [b] -> ([(a, b)])
cartesianProductBalanced as bs =
concatMap (zip as) (tails bs)
++ concatMap (flip zip bs) (tail (tails as))```

`tails` gives the list of successively applying `tail`, e.g., `tails [1..4]` = `[[1,2,3,4],[2,3,4],[3,4],[4],[]]`

This definition for `cartesianProductBalanced` essentially traverses the diagonal of the space and then the lower triangle of the matrix, progressing away from the diagonal to the bottom-left, then traversing the upper triangle of the matrix (the last line of code), progressing away from the diagonal to the top-right corner. The ordering for a 4×4 space is then:

```*Main> cartesianProductBalanced [1..4] [1..4]
[(1,1),(2,2),(3,3),(4,4),(1,2),(2,3),(3,4),(1,3),(2,4),(1,4),(2,1),(3,2),(4,3),(3,1),(4,2),(4,1)]```

This was more balanced, giving me a better view of the space, but does not scale well: traversing the elements of this Cartesian product linearly means first traversing all the way down the diagonal, which could be very long! So, we’ve ended up with a similar problem to the row-major traversal.

Instead, I finally settled on a “tiling” Cartesian product:

```cartesianProductTiling :: [a] -> [b] -> [(a, b)]
cartesianProductTiling [] _ = []
cartesianProductTiling _ [] = []
cartesianProductTiling [a] [b] = [(a, b)]
cartesianProductTiling as bs =
cartesianProductTiling as1 bs1
++ cartesianProductTiling as2 bs1
++ cartesianProductTiling as1 bs2
++ cartesianProductTiling as2 bs2
where
(as1, as2) = splitAt ((length as) `div` 2) as
(bs1, bs2) = splitAt ((length bs) `div` 2) bs```

This splits the space into four quadrants and recursively applies itself to the upper-left, then lower-left, then upper-right, then lower-right, e.g.,

```*Main> cartesianProductTiling [1..4] [1..4]
[(1,1),(2,1),(1,2),(2,2),(3,1),(4,1),(3,2),(4,2),(1,3),(2,3),(1,4),(2,4),(3,3),(4,3),(3,4),(4,4)]```

Thus, we explore the upper-left portion first, without having to traverse down a full row, column, or diagonal. This turned out to be much better for my purposes of searching the space of possible axioms for an algebraic structure.

Note that this visits the elements in Z-order (also known as the Lebesgue space-filling curve) [Actually, it is a reflected Z-order curve, but that doesn’t matter here].

The only downside to this tiling approach is that it does not work for infinite spaces (as it calculates the length of `as` and `bs`), so we cannot exploit Haskell’s laziness here to help us in the infinite case. I’ll leave it as an exercise for the reader to define a tiling version that works for infinite lists.

Of course, there are many other possibilities depending on your application domain!

Later edit:  In the case of the example I mentioned, I actually do not need the full Cartesian product since the order of pairs was not relevant to me (equality is symmetric) and neither do I need the pairs of identical elements (equality is reflexive). So for my program, computing just the lower triangle of Cartesian product square is much more efficient, i.e,:

```cartesianProductLowerTriangle :: [a] -> [b] -> [(a, b)]
cartesianProductLowerTriangle as bs = concatMap (zip as) (tail (tails bs))```

However, this does not have the nice property of tiling product where we visit the elements in a more balanced fashion. I’ll leave that one for another day (I got what I needed from my program in the end anyway).

# Rearranging equations using a zipper

Whilst experimenting with some ideas for a project, I realised I needed a quick piece of code to rearrange equations (defined in terms of +, *, -, and /) in AST form, e.g., given an AST for the equation x = y + 3, rearrange to get y = x - 3.

I realised that equations can be formulated as zippers over an AST, where operations for navigating the zipper essentially rearrange the equation. I thought this was quite neat, so I thought I would show the technique here. The code is in simple Haskell.

I’ll show the construction for a simple arithmetic calculus with the following AST data type of terms:

```data Term = Add Term Term
| Mul Term Term
| Div Term Term
| Sub Term Term
| Neg Term
| Var String
| Const Integer
```

with some standard pretty printing code:

```instance Show Term where
show (Add t1 t2) = (show' t1) ++ " + " ++ (show' t2)
show (Mul t1 t2) = (show' t1) ++ " * " ++ (show' t2)
show (Sub t1 t2) = (show' t1) ++ " - " ++ (show' t2)
show (Div t1 t2) = (show' t1) ++ " / " ++ (show' t2)
show (Neg t) = "-" ++ (show' t)
show (Var v) = v
show (Const n) = show n
```

where show' is a helper to minimise brackets e.g. pretty printing “-(v)” as “-v”.

```show' :: Term -> String
show' (Var v) = v
show' (Const n) = show n
show' t@(Neg (Var v)) = show t
show' t@(Neg (Const n)) = show t
show' t = "(" ++ show t ++ ")"
```

Equations can be defined as pairs of terms, i.e., ‘T1 = T2’ where T1 and T2 are both represented by values of Term. However, instead, I’m going to represent equations via a zipper.

Zippers (described beautifully in the paper by Huet) represent values that have some subvalue “in focus”. The position of the focus can then be shifted through the value, refocussing on different parts. This is encoded by pairing a focal subvalue with a path to this focus, which records the rest of the value that is not in focus. For equations, the zipper type pairs a focus Term (which we’ll think of as the left-hand side of the equation) with a path (which we’ll think of as the right-hand side of the equation).

```data Equation = Eq Term Path
```

Paths give a sequence of direction markers, essentially providing an address to the term in focus, starting from the root, where each marker is accompanied with the label of the parent node and the subtree of the branch not taken, i.e., a path going left is paired with the right subtree (which is not on the path to the focus).

```data Path = Top (Either Integer String)  -- At top: constant or variable
| Bin Op                -- OR in a binary operation Op,
Dir               --    in either left (L) or right (R) branch
Term              --    with the untaken branch
Path              --    and the rest of the equation
| N Path                -- OR in the unary negation operation

data Dir = L | R
data Op = A | M | D | S | So | Do
```

The Op type gives tags for every operation, as well as additional tags So and Do which represent sub and divide but with arguments flipped. This is used to get an isomorphism between the operations that zip “up” and “down” the equation zipper, refocussing on subterms.

A useful helper maps tags to their operations:

```opToTerm :: Op -> (Term -> Term -> Term)
opToTerm M = Mul
opToTerm D = Div
opToTerm S = Sub
opToTerm So = (\x -> \y -> Sub y x)
opToTerm Do = (\x -> \y -> Div y x)
```

Equations are pretty printed as follows:

```instance Show Path where
show p = show . pathToTerm \$ p

instance Show Equation where
show (Eq t p) = (show t) ++ " = " ++ (show p)
```

where pathToTerm converts paths to terms:

```pathToTerm :: Path -> Term
pathToTerm (Top (Left c)) = Const c
pathToTerm (Top (Right v))= Var v
pathToTerm (Bin op L t p) = (opToTerm op) (pathToTerm p) t
pathToTerm (Bin op R t p) = (opToTerm op) t (pathToTerm p)
pathToTerm (N p)          = Neg (pathToTerm p)
```

Now onto the zipper operations which providing rebalancing of the equation. Equations are zipped-down by left and right, which for a binary operation focus on either the left or right argument respectively, for unary negation focus on the single argument, and for constants or variables does nothing. When going left or right, the equations are rebalanced with their inverse arithmetic operations (show in the comments here):

```left (Eq (Var s) p)     = Eq (Var s) p
left (Eq (Const n) p)   = Eq (Const n) p
left (Eq (Add t1 t2) p) = Eq t1 (Bin S L t2 p)   -- t1 + t2 = p  -> t1 = p - t2
left (Eq (Mul t1 t2) p) = Eq t1 (Bin D L t2 p)   -- t1 * t2 = p  -> t1 = p / t2
left (Eq (Div t1 t2) p) = Eq t1 (Bin M L t2 p)   -- t1 / t2 = p  -> t1 = p * t2
left (Eq (Sub t1 t2) p) = Eq t1 (Bin A L t2 p)   -- t1 - t2 = p  -> t1 = p + t2
left (Eq (Neg t) p)     = Eq t (N p)             -- -t = p       -> t = -p

right (Eq (Var s) p)     = Eq (Var s) p
right (Eq (Const n) p)   = Eq (Const n) p
right (Eq (Add t1 t2) p) = Eq t2 (Bin So R t1 p)  -- t1 + t2 = p -> t2 = p - t1
right (Eq (Mul t1 t2) p) = Eq t2 (Bin Do R t1 p)  -- t1 * t2 = p -> t2 = p / t1
right (Eq (Div t1 t2) p) = Eq t2 (Bin D R t1 p)   -- t1 / t2 = p -> t2 = t1 / p
right (Eq (Sub t1 t2) p) = Eq t2 (Bin S R t1 p)   -- t1 - t2 = p -> t2 = t1 - p
right (Eq (Neg t) p)     = Eq t (N p)
```

In both left and right, Add and Mul become subtraction and dividing, but in right in order for the the zipping-up operation to be the inverse, subtraction and division are represented using the flipped So and Do markers.

Equations are zipped-up by up, which unrolls one step of the path and reforms the term on the left-hand side from that on the right. This is the inverse of left and right:

```up (Eq t1 (Top a))        = Eq t1 (Top a)
up (Eq t1 (Bin A L t2 p)) = Eq (Sub t1 t2) p -- t1 = t2 + p -> t1 - t2 = p
up (Eq t1 (Bin M L t2 p)) = Eq (Div t1 t2) p -- t1 = t2 * p -> t1 / t2 = p
up (Eq t1 (Bin D L t2 p)) = Eq (Mul t1 t2) p -- t1 = p / t2 -> t1 * t2 = p
up (Eq t1 (Bin S L t2 p)) = Eq (Add t1 t2) p -- t1 = p - t2 -> t1 + t2 = p

up (Eq t1 (Bin So R t2 p)) = Eq (Add t2 t1) p -- t1 = p - t2 -> t2 + t1 = p
up (Eq t1 (Bin Do R t2 p)) = Eq (Mul t2 t1) p -- t1 = p / t2 -> t2 * t1 = p
up (Eq t1 (Bin D R t2 p))  = Eq (Div t2 t1) p -- t1 = t2 / p -> t2 / t1 = p
up (Eq t1 (Bin S R t2 p))  = Eq (Sub t2 t1) p -- t1 = t2 - p -> t2 - t1 = p

up (Eq t1 (N p))           = Eq (Neg t1) p    -- t1 = -p     -> -t1 = p
```

And that’s it! Here is an example of its use from GHCi.

```foo = Eq (Sub (Mul (Add (Var "x") (Var "y")) (Add (Var "x")
(Const 1))) (Const 1)) (Top (Left 0))

*Main> foo
((x + y) * (x + 1)) - 1 = 0

*Main> left \$ foo
(x + y) * (x + 1) = 0 + 1

*Main> right . left \$ foo
x + 1 = (0 + 1) / (x + y)

*Main> left . right . left \$ foo
x = ((0 + 1) / (x + y)) - 1

*Main> up . left . right . left \$ foo
x + 1 = (0 + 1) / (x + y)

*Main> up . up . left . right . left \$ foo
(x + y) * (x + 1) = 0 + 1

*Main> up . up . up . left . right . left \$ foo
((x + y) * (x + 1)) - 1 = 0
```

It is straightforward to prove that: up . left \$ x = x (when left x is not equal to x) and up . right \$ x = x(when right x is not equal to x).

Note, I am simply rebalancing the syntax of equations: this technique does not help if you have multiple uses of a variable and you want to solve the question for a particular variable, e.g. y = x + 1/(3x), or quadratics.

Here’s a concluding thought. The navigation operations left, right, and up essentially apply the inverse of the operation in focus to each side of the equation. We could therefore reformulate the navigation operations in terms of any group: given a term L ⊕ R under focus where is the binary operation of a group with inverse operation -1, then navigating left applies ⊕ R-1 to both sides and navigating right applies ⊕ L-1. However, in this blog post there is a slight difference: navigating applies the inverse to both sides and then reduces the term of the left-hand side using the group axioms X ⊕ X-1 = I (where I is the identity element of the group) and X ⊕ I = X such that the term does not grow larger and larger with inverses.

I wonder if there are other applications, which have a group structure (or number of interacting groups), for which the above zipper approach would be useful?

# Subcategories & “Exofunctors” in Haskell

In my previous post I discussed the new constraint kinds extension to GHC, which provides a way to get type-indexed constraint families in GHC/Haskell. The extension provides some very useful expressivity. In this post I’m going to explain a possible use of the extension.

In Haskell the Functor class is misleading named as it actually captures the notion of an endofunctor, not functors in general. This post shows a use of constraint kinds to define a type class of exofunctors; that is, functors that are not necessarily endofunctors. I will explain what all of this means.

This example is just one from a draft note (edit July 2012: draft note subsumed by my TFP 2012 submission) explaining the use of constraint families, via the constraint kinds extension, for describing abstract structures from category theory that are parameterised by subcategories, including non-endofunctors, relative monads, and relative comonads.

I will try to concisely describe any relevant concepts from category theory, through the lens of functional programming, although I’ll elide some details.

The starting point of the idea is that programs in Haskell can be understood as providing definitions within some category, which we will call Hask. Categories comprise a collection of objects and a collection of morphisms which are mappings between objects. Categories come equipped with identity morphisms for every object and an associative composition operation for morphisms (see Wikipedia for a more complete, formal definition). For Hask, the objects are Haskell types, morphisms are functions in Haskell, identity morphisms are provided by the identity function, and composition is the usual function composition operation. For the purpose of this discussion we are not really concerned about the exact properties of Hask, just that Haskell acts as a kind of internal language for category theory, within some arbitrary category Hask (Dan Piponi provides some discussion on this topic).

### Subcategories

Given some category C, a subcategory of C comprises a subcollection of the objects of C and a subcollection of the morphisms of C which map only between objects in the subcollection of this subcategory.

We can define for Hask a singleton subcategory for each type, which has just that one type as an object and functions from that type to itself as morphisms e.g. the Int-subcategory of Hask has one object, the Int type, and has functions of type Int → Int as morphisms. If this subcategory has all the morphisms Int → Int it is called a full subcategory. Is there a way to describe “larger” subcategories with more than just one object?

Via universal quantification we could define the trivial (“non-proper”) subcategory of Hask with objects of type a (implicitly universally quantified) and morphisms a -> b, which is just Hask again. Is there a way to describe “smaller” subcategories with fewer than all the objects, but more than one object? Yes. For this we use type classes.

### Subcategories as type classes

The instances of a single parameter type class can be interpreted as describing the members of a set of types (or a relation on types for multi-parameter type classes). In a type signature, a universally quantified type variable constrained by a type class constraint represents a collection of types that are members of the class. E.g. for the Eq class, the following type signature describes a collection of types for which there are instances of Eq:

Eq a => a

The members of Eq are a subcollection of the objects of Hask. Similarly, the type:

(Eq a, Eq b) => (a -> b)

represents a subcollection of the morphisms of Hask mapping between objects in the subcollection of objects which are members of Eq. Thus, the Eq class defines an Eq-subcategory of Hask with the above subcollections of objects and morphisms.

Type classes can thus be interpreted as describing subcategories in Haskell. In a type signature, a type class constraint on a type variable thus specifies the subcategory which the type variable ranges over the objects of. We will go on to use the constraint kinds extension to define constraint-kinded type families, allowing structures from category theory to be parameterised by subcategories, encoded as type class constraints. We will use functors as the example in this post (more examples here).

In category theory, a functor provides a mapping between categories e.g. F : C → D, mapping the objects and morphisms of C to objects and morphisms of D. Functors preserves identities and composition between the source and target category (see Wikipedia for more). An endofunctor is a functor where C and D are the same category.

The type constructor of a parametric data type in Haskell provides an object mapping from Hask to Hask e.g. given a data type data F a = ... the type constructor F maps objects (types) of Hask to other objects in Hask. A functor in Haskell is defined by a parametric data type, providing an object mapping, and an instance of the well-known Functor class for that data type:

``` class Functor f where fmap :: (a -> b) -> f a -> f b ```

which provides a mapping on morphisms, called fmap. There are many examples of functors in Haskell, for examples lists, where the fmap operation is the usual map operation, or the Maybe type. However, not all parametric data types are functors.

It is well-known that the Set data type in Haskell cannot be made an instance of the Functor class. The Data.Set library provides a map operation of type:

``` Set.map :: (Ord a, Ord b) => (a -> b) -> Set a -> Set b ```

The Ord constraint on the element types is due to the implementation of Set using balanced binary trees, thus elements must be comparable. Whilst the data type is declared polymorphic, the constructors and transformers of Set allow only elements of a type that is an instance of Ord.

Using Set.map to define an instance of the Functor class for Set causes a type error:

``` instance Functor Set where fmap = Data.Set.map ... foo.lhs:4:14: No instances for (Ord b, Ord a) arising from a use of `Data.Set.map' In the expression: Data.Set.map In an equation for `fmap': fmap = Data.Set.map In the instance declaration for `Functor Set' ```

The type error occurs as the signature for fmap has no constraints, or the empty (always true) constraint, whereas Set.map has Ord constraints. A mismatch occurs and a type error is produced.

The type error is however well justified from a mathematical perspective.

### Haskell functors are not functors, but endofunctors

First of all, the name Functor is a misnomer; the class actually describes endofunctors, that is functors which have the same category for their source and target. If we understand type class constraints as specifying a subcategory, then the lack of constraints on fmap means that Functor describes endofunctors HaskHask.

The Set data type is not an endofunctor; it is a functor which maps from the Ord-subcategory of Hask to Hask. Thus Set :: OrdHask. The class constraints on the element types in Set.map declare the subcategory of Set functor to which the morphisms belong.

### Type class of exofunctors

Can we define a type class which captures functors that are not necessarily endofunctors, but may have distinct source and target categories? Yes, using an associated type family of kind Constraint.

The following ExoFunctor type class describes a functor from a subcategory of Hask to Hask:

``` {-# LANGUAGE ConstraintKinds #-} {-# LANGUAGE TypeFamilies #-} class ExoFunctor f where type SubCat f x :: Constraint fmap :: (SubCat f a, SubCat f b) => (a -> b) -> f a -> f b ```

The SubCat family defines the source subcategory for the functor, which depends on f. The target subcategory is just Hask, since f a and f b do not have any constraints.

We can now define the following instance for Set:

``` instance ExoFunctor Set where type SubCat Set x = Ord x fmap = Set.map ```

Endofunctors can also be made an instance of ExoFunctor using the empty constraint e.g.:

``` instance ExoFunctor [] where type SubCat [] a = () fmap = map ```

(Aside: one might be wondering whether we should also have a way to restrict the target subcategory to something other than Hask here. By covariance we can always “cast” a functor C → D, where D is a subcategory of some other category E, to C → E without any problems. Thus, there is nothing to be gained from restricting the target to a subcategory, as it can always be reinterpreted as Hask.)

### Conclusion (implementational restrictions = subcategories)

Subcategory constraints are needed when a data type is restricted in its polymorphism by its operations, usually because of some hidden implementational details that have permeated to the surface. These implementational details have until now been painful for Haskell programmers, and have threatened abstractions such as functors, monads, and comonads. Categorically, these implementational restrictions can be formulated succinctly with subcategories, for which there are corresponding structures of non-endofunctors, relative monads, and relative comonads. Until now there has been no succinct way to describe such structures in Haskell.

Using constraint kinds we can define associated type families, of kind Constraint, which allow abstract categorical structures, described via their operations as a type class, to be parameterised by subcategories on a per-instance basis. We can thus define a class of exofunctors, i.e. functors that are not necessarily endofunctors, which we showed here. The other related structures which are difficult to describe in Haskell without constraint kinds: relative monads and relative comonads, are discussed further in a draft note (edit July 2012: draft note subsumed by my TFP 2012 submission). The note includes examples of a Set monad and an unboxed array comonad, both of which expose their implementational restrictions as type class constraints which can be described as subcategory constraints.
Any feedback on this post or the draft note is greatly appreciated. Thanks.

# The Koch Snowflake

This post has been imported from my old blog.

I mentioned a couple of weeks ago I was going to write a post about a certain fractal. Now I have finally gotten round to writing something it has come at a very appropriate time as Britain has seen an unusual amount of snow recently. Sadly all of that snow has melted now. I am going to briefly introduce the Koch Snowflake (also known as the Koch Island, or Koch Star) which is a fractal with a very simple construction. I will first informally describe a geometrical construction of the fractal and will then show some of its properties, particularly the property that the fractal has an infinite length enclosing a finite area.

Construction
The base case of the construction is an equilateral triangle. For each successive iteration the construction proceeds as follows:

• Divide each edge of the polygon (say of length a) into three equal segments of length a/3.
• Replace the middle segment with an equilateral triangle of side a/3.
• Remove the base edge of the new equilateral triangle to form a continuous curve with other line segments.

Thus each edge is transformed as such:

The first 5 iterations look like:

Properties
Two properties of this fractal may seem paradoxical at first but are easily shown. I will show the derivations here for the interested without skipping too many steps.
For an infinite number of iterations the perimeter of the fractal (the length of all the edges) tends to infinity whilst the area enclosed remains finite.

Infinite Perimeter
First consider the number of edges for an infinite number of iterations. Initially the number of edges is 3, each iteration transforms a single edge into 4 edges (see above), thus the series of the edge count is: 3, 12, 48, …. The general form is:

Starting from an edge length of a each iteration divides the edge length by 3. The following defines the edge length for iteration n and from this calculates the perimeter for iteration n. The limit as n tends to infinity is found, showing that the
perimeter is infinite.

Thus it is easy to see that the perimeter length is infinite through standard results on limits.

Finite Area
The area calculation is a little bit more involved. The result can be seen on Mathworld or Wikipedia but a full derivation isn’t given, so I show my derivation here for the interested.

The area of the n-th iteration of the Koch snowflake is the area of the base triangle + the area of the new, smaller, triangles added to each edge.
From the above length calculations and construction we know that each iteration of the construction divides the length of the edges by three. First consider the relationship between the area of a triangle of edge length a and a triangle of edge length a/3

Unsurprisingly we see that dividing the edge length by 3 divides the area by 9. At each iteration we add new equilateral triangles to each edge, a ninth of the area of the previous iteration’s triangles.
We formulate this as a summation of base case A0 plus the next n-1 iterations where the triangle area at each stage is A0 divided by 32n or 9n (successive divisions of the area by 9). The number of edges is defined by the previous iterations number of edges 3 * 4n-1, thus at the n-th iteration we need to add this number of triangles to the snowflake.

Thus the area is finite at 8/5 times the area of the base equilateral triangle.

Summary
So the Koch snowflake construction has been introduced and it has been shown relatively easily that the area of a Koch snowflake tends to a finite limit of 8/5 times the base case area (5) and that the length of perimeter tends to infinity (4). Next time I will probably talk a bit about the fractal dimension of the Koch snowflake and also show an interesting problem that uses the Koch snowflake.

As a last note, following on from the blog post I made about the Python turtle library, a turtle program of the Koch snowflake for Python can be downloaded here.