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**Approximations of points and polygonal chains**

Approximation of points by a line: find a line l (if one exists) such that all points are within distance ε from l, under some distance metric. Other definitions possible: find line l that minimizes the maximum distance from the points to l.

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**Distance Metrics Uniform metric (prev. slide). Euclidean metric.**

yi-f(xi) ≤ ε Euclidean metric. L1 and L∞ metrics. Other measures possible (will be introduced latter).

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**Approx. of points by line**

Each vertical line segment maps to a parallel strip in a dual point/line space (duality transform, to be explained shortly). Can define the strip as intersection of two halfplanes. Then solve a 2-D linear programming (LP) problem, which takes O(n) time.

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**Approx. points by min-# chain**

Would a simple Greedy algorithm work? Stab as many as possible, discard them and repeat.

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**Greedy vs. Dynamic Programming**

If objects are convex and disjoint, O(n) time greedy algorithm is possible (vertices are unrestricted, that is, can lie anywhere). For intersecting objects, dynamic programming gives O(n2) and O(n2 log n) time, depending on how visiting order of objects is defined.

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Min-# vs. Min-ε Min-#: given some error bound, minimize number of vertices of approximating path. Min-ε: minimize the error of an approximating path that has at most a specified number of vertices, m (m < n). If finite set of possible approximation errors, and if set can be computed in reasonable time, then can use min-# to solve min-ε.

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Points vs. Chains Define a piecewise linear region associated with the input, as in figure. May ask approximating path stays in that region, for chains. Chains may be monotone (functions) or not.

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**Min-# for Piecewise Linear Function**

Approximation must stay in error region. Approximation sought: monotone polygonal line with min-# of vertices. Use “edge visibility” concept. A point u of polygon P is visible from edge e, if there is point v on e such that line segment uv is contained in P. An edge e’ is visible from e if there is a point of e’ visible from e. Visibility polygon VP(P,e): points of P visible from e.

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Visibility: VP(P,e) VP(P,e) splits the polygon in a few connected regions: Visible region/polygon; Invisible polygons. Consider the invisible polygon P’ that contains end point of chain. Common boundary of P and P’ is a line segment called “window”. Use “convex hull” algorithm to find windows. Results in O(n) time algorithm.

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**Error criteria for chains**

E1: Tolerance zone. E2: Parallel-strip (also called “infinite beam”). E3: Parallel rectangle: half of min. width of rectangle containing pk, i≤ k≤ j, with parallel sides orthogonal on line segment pipj. E4: Rectangle: required only contain pk, i≤ k ≤j.

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**General Approach Using Graphs**

Given: C=p1p2…pn, polygonal chain. Directed graph G(C): Vertex vi corresponds to point pi. Edge (vi,vj) has weight = error of segment pipj. For given error bound ε: G(C,ε): Subgraph of G with edge weights ≤ ε. A path in G(C,ε): corresponds to approximating curve with error at most ε. Approximating curve for min-#: Shortest path from v1 to vn in G(C,ε) (length is # arcs). Use BFS to find it. Min-ε: minimize ε such that in G(C,ε) there is path with at most a specified # of vertices.

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**Complexity If G(C) can be computed in f(n) time:**

Min-# can be solved in O(f(n)+n2) time. Min-ε can be solved in O(f(n)+n2 log n) time. If G(C,ε) can be computed in g(n) time: Min-# can be solved in O(g(n)+n2) time. Most of known solutions use this approach (see solution on handout paper for some exceptions). The “closed” min-# and min-ε problems can eventually be solved by solving n “open” problems (better in some cases; depends on restrictions).

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**Complexities Under error criteria discussed:**

e1: O(n2) and O(n2 log n) time, O(n) space. Constructs G(C,ε). Makes use of binary search with less space. e2: O(n2 log n) time, O(n) and O(n2) space. Constructs G(C). O(n2) possible if some condition holds. Even o(n2) possible if that condition holds. e3: O(n2 log n) time, O(n2) space (O(n) space possible?!). e4: O(n log n) time, O(mn log n log (n/m)) space.

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Min-# under E4 Property: In G(C,ε), if (vi,vj) є G, then (vk,vl) є G, i ≤ k < l ≤ j. Algorithm 1: v=1; i(1)=1; i(v+1)=max{j | (vi,vj) є G(c,ε)}; if i(v+1)==n then done; else v=v+1; return to 2. The algorithm is greedy.

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E4 (cont.) There is a rectangle of width w that covers a set V of points such that it contains an edge of CH(V). W(i,j): min. width of rectangle covering V. For e є CH(V): Antipodal point: point v of CH(V) farthest from line containing e. d(e): distance from v to line containing e. W(i,j)=min{d(e) | e is on CH(i,j)}. Given CH(V), can find antipodal points of all edges in O(|V|) time.

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**E4 (cont.) Algorithm 1’: Step 2 is main step.**

v=1; i(1)=1; i(v+1)=max{j | i(v) < j ≤ n, W(i(v),j) ≤ w}; if i(v+1)==n then done; else v=v+1; go to step 2. Step 2 is main step. Will show for v=1 van be executed in O(i* log i*) where i*=i(2).

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E4 (cont.) W(1,j) nondecreasing in j: use doubling binary search to find i*. Find m, l such that W(1,m) ≤ w, W(1,l) > w and m < l ≤ min{2m,n} Start with m=1, l=2; double each until holds. i* lies between m and l-1. Again binary search to find i*:

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**E4 (cont.) Algorithm 2: k=m; l=l-1; Construct and keep CH(1,k);**

h=(k+l)/2; Construct CH(k+1,h); Construct CH(1,h) from CH(1,k), CH(k+1,h); Compute wh=W(1,h); If wh ≤ w then k=h; else l=h-1; If k==l then i*=k and done; else keep CH(1,k) and go to step 2.

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**E4 (cont.) Algorithm 1’+2 takes O(n log n).**

Compute W(1,2g), g=1,2,..,M+1, M=log m. W(1,j) can be found in O(j log j). O(Σ2glog2g)=O(i*logi*). Note that m=O(i*). Step 2 in Alg. 2 takes O(i*logi*) CH(1,h): O(h) (from CH(1,k) and CH(k+1,h); W(1,h): O(h) time; In g-th iteration CH(k+1,h) is on at most m2-g=2M-g points: O(2M-glog2M-g) time. Executed at most M times (g=1,…M):O(Σ2M-glogM-g)=O(i*logi*).

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**Min-ε under E4 O(mn(log n)(log n/m) time. (output sensitive).**

Algorithm 3 (compute min. width w*). M’=m; i=1; w=0; Find j, by doubling binary search, such that i < j ≤ n, pi,…,pn cannot be covered by m’ or fewer rectangles with width less than W(i,j), pi,…,pn can be covered by m’ or fewer rectangles with width less than W(i,j+1). i=j; m’=m-1; w=max{w,W(i,j)}; If i=n done; else go to step 2. Note: in a sequence of rectangles with width w* found by min-#, the first rectangle cover points p1,…,pj, where j is found in first execution of step 2 above (then induction).

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